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From: jannh@xxxxxxxxxx
To: kernel-hardening@xxxxxxxxxxxxxxxxxx, linux-hardening@xxxxxxxxxxxxxxx
Subject: An analysis of current and potential security mitigations based on a TIOCSPGRP exploit

I published a blogpost with the title
"How a simple Linux kernel memory corruption bug can lead to complete system
compromise - An analysis of current and potential kernel security mitigations"
on Project Zero's blog at
(The version there looks a bit nicer because it has proper formatting.)

In case folks want to use it as a starting point for discussion, I'm sending
a copy to this list.

Note that it's quite long, so please make sure to trim quotes to the relevant

This blog post describes a straightforward Linux kernel locking bug and how I
exploited it against Debian Buster's 4.19.0-13-amd64 kernel. Based on that, it
explores options for security mitigations that could prevent or hinder
exploitation of issues similar to this one.

I hope that stepping through such an exploit and sharing this compiled knowledge
with the wider security community can help with reasoning about the relative
utility of various mitigation approaches.

A lot of the individual exploitation techniques and mitigation options that I am
describing here aren't novel. However, I believe that there is value in writing
them up together to show how various mitigations interact with a fairly normal
use-after-free exploit.

Our bugtracker entry for this bug, along with the proof of concept, is at

Code snippets in this blog post that are relevant to the exploit are taken from
the upstream 4.19.160 release, since that is what the targeted Debian kernel is
based on; some other code snippets are from mainline Linux.

(In case you're wondering why the bug and the targeted Debian kernel are from
end of last year: I already wrote most of this blogpost around April, but only
recently finished it)

I would like to thank Ryan Hileman (https://twitter.com/lunixbochs) for a
discussion we had a while back about how static analysis might fit into static
prevention of security bugs (but note that Ryan hasn't reviewed this post and
doesn't necessarily agree with any of my opinions). I also want to thank Kees
Cook (https://twitter.com/kees_cook) for providing feedback on an earlier
version of this post (again, without implying that he necessarily agrees with
everything), and my Project Zero colleagues for reviewing this post and
frequent discussions about exploit mitigations.

Background for the bug
On Linux, terminal devices (such as a serial console or a virtual console
(https://en.wikipedia.org/wiki/Virtual_console)) are represented by a
`struct tty_struct`. Among other things, this structure contains fields used
for the job control
features of terminals, which are usually modified using a set of ioctls

struct tty_struct {
        spinlock_t ctrl_lock;
        struct pid *pgrp;               /* Protected by ctrl lock */
        struct pid *session;
        struct tty_struct *link;

The `pgrp` field points to the foreground process group of the terminal
(normally modified from userspace via the `TIOCSPGRP` ioctl); the `session`
field points to the session associated with the terminal. Both of these fields
do not point directly to a process/task, but rather to a `struct pid`. `struct
pid` ties a specific incarnation of a numeric ID to a set of processes that use
that ID as their PID (also known in userspace as TID), TGID (also known in
userspace as PID), PGID, or SID. You can kind of think of it as a weak
reference to a process, although that's not entirely accurate. (There's some
extra nuance around `struct pid` when `execve()` is called by a non-leader
thread, but that's irrelevant here.)

All processes that are running inside a terminal and are subject to its job
control refer to that terminal as their "controlling terminal" (stored in
`->signal->tty` of the process).

A special type of terminal device are
pseudoterminals (https://man7.org/linux/man-pages/man7/pty.7.html), which are
used when you, for example, open a terminal application in a graphical
environment or connect to a remote machine via SSH. While other terminal
devices are connected to some sort of hardware, both ends of a pseudoterminal
are controlled by userspace, and pseudoterminals can be freely created by
(unprivileged) userspace. Every time `/dev/ptmx`
(https://man7.org/linux/man-pages/man4/pts.4.html) (short for
"pseudoterminal multiplexor") is opened, the resulting file descriptor
represents the device side (referred to in documentation and kernel sources as
"the pseudoterminal master" (https://man7.org/linux/man-pages/man7/pty.7.html))
of a new pseudoterminal.
You can read from it to get the data that should be printed on the emulated
screen, and write to it to emulate keyboard inputs. The corresponding terminal
device (to which you'd usually connect a shell) is automatically created by the
kernel under `/dev/pts/<number>`.

One thing that makes pseudoterminals particularly strange is that both ends of
the pseudoterminal have their own `struct tty_struct`, which point to each
other using the `link` member, even though the device side of the
pseudoterminal does not have terminal features like job control - so many of
its members are unused.

Many of the ioctls for terminal management can be used on both ends of the
pseudoterminal; but no matter on which end you call them, they affect the same
state, sometimes with minor differences in behavior. For example, in the ioctl
handler for `TIOCGPGRP`:

 *      tiocgpgrp               -       get process group
 *      @tty: tty passed by user
 *      @real_tty: tty side of the tty passed by the user if a pty else the tty
 *      @p: returned pid
 *      Obtain the process group of the tty. If there is no process group
 *      return an error.
 *      Locking: none. Reference to current->signal->tty is safe.
static int tiocgpgrp(struct tty_struct *tty, struct tty_struct *real_tty, pid_t __user *p)
        struct pid *pid;
        int ret;
         * (tty == real_tty) is a cheap way of
         * testing if the tty is NOT a master pty.
        if (tty == real_tty && current->signal->tty != real_tty)
                return -ENOTTY;
        pid = tty_get_pgrp(real_tty);
        ret =  put_user(pid_vnr(pid), p);
        return ret;

As documented in the comment above, these handlers receive a pointer `real_tty`
that points to the normal terminal device; an additional pointer `tty` is
passed in that can be used to figure out on which end of the terminal the ioctl
was originally called. As this example illustrates, the `tty` pointer is
normally only used for things like pointer comparisons. In this case, it is
used to prevent `TIOCGPGRP` from working when called on the terminal side by a
process which does not have this terminal as its controlling terminal.

Note: If you want to know more about how terminals and job control are intended
to work, the book "The Linux Programming Interface" (https://man7.org/tlpi/)
provides a nice introduction to how these older parts of the userspace API are
supposed to work. It doesn't describe any of the kernel internals though, since
it's written as a reference for userspace programming. And it's from 2010, so
it doesn't have anything in it about new APIs that have showed up over the last

The bug
The bug was in the ioctl handler `tiocspgrp`:

 *      tiocspgrp               -       attempt to set process group
 *      @tty: tty passed by user
 *      @real_tty: tty side device matching tty passed by user
 *      @p: pid pointer
 *      Set the process group of the tty to the session passed. Only
 *      permitted where the tty session is our session.
 *      Locking: RCU, ctrl lock
static int tiocspgrp(struct tty_struct *tty, struct tty_struct *real_tty, pid_t __user *p)
        struct pid *pgrp;
        pid_t pgrp_nr;
        if (get_user(pgrp_nr, p))
                return -EFAULT;
        pgrp = find_vpid(pgrp_nr);
        real_tty->pgrp = get_pid(pgrp);

The `pgrp` member of the terminal side (`real_tty`) is being modified, and the
reference counts of the old and new process group are adjusted accordingly
using `put_pid` and `get_pid`; but the lock is taken on `tty`, which can be
either end of the pseudoterminal pair, depending on which file descriptor we
pass to `ioctl()`. So by simultaneously calling the `TIOCSPGRP` ioctl on both
sides of the pseudoterminal, we can cause data races between concurrent
accesses to the `pgrp` member. This can cause reference counts to become skewed
through the following races:

  ioctl(fd1, TIOCSPGRP, pid_A)        ioctl(fd2, TIOCSPGRP, pid_B)
    spin_lock_irq(...)                  spin_lock_irq(...)
    real_tty->pgrp = get_pid(A)
                                        real_tty->pgrp = get_pid(B)
    spin_unlock_irq(...)                spin_unlock_irq(...)

  ioctl(fd1, TIOCSPGRP, pid_A)        ioctl(fd2, TIOCSPGRP, pid_B)
    spin_lock_irq(...)                  spin_lock_irq(...)
                                        real_tty->pgrp = get_pid(B)
    real_tty->pgrp = get_pid(A)
    spin_unlock_irq(...)                spin_unlock_irq(...)

In both cases, the refcount of the old `struct pid` is decremented by 1 too
much, and either A's or B's is incremented by 1 too much.

Once you understand the issue, [the fix](https://git.kernel.org/linus/54ffccbf053b) seems relatively obvious:
         if (session_of_pgrp(pgrp) != task_session(current))
                 goto out_unlock;
         retval = 0;
-        spin_lock_irq(&tty->ctrl_lock);
+        spin_lock_irq(&real_tty->ctrl_lock);
         real_tty->pgrp = get_pid(pgrp);
-        spin_unlock_irq(&tty->ctrl_lock);
+        spin_unlock_irq(&real_tty->ctrl_lock);
         return retval;

Attack stages
In this section, I will first walk through how my exploit works; afterwards I
will discuss different defensive techniques that target these attack stages.

Attack stage: Freeing the object with multiple dangling references
This bug allows us to probabilistically skew the refcount of a `struct pid`
down, depending on which way the race happens: We can run colliding `TIOCSPGRP`
calls from two threads repeatedly, and from time to time that will mess up the
refcount. But we don't immediately know how many times the refcount skew has
actually happened.

What we'd really want as an attacker is a way to skew the refcount
deterministically. We'll have to somehow compensate for our lack of information
about whether the refcount was skewed successfully. We could try to somehow
make the race deterministic (seems difficult), or after each attempt to skew
the refcount assume that the race worked and run the rest of the exploit (since
if we didn't skew the refcount, the initial memory corruption is gone, and
nothing bad will happen), or we can attempt to find an information leak that
lets us figure out the state of the reference count.

On typical desktop/server distributions, the following approach works
(unreliably, depending on RAM size) for setting up a freed `struct pid` with
multiple dangling references:

1. Allocate a new `struct pid` (by creating a new task).
2. Create a large number of references to it (by sending messages with
   `SCM_CREDENTIALS` to unix domain sockets, and leaving those messages queued
3. Repeatedly trigger the `TIOCSPGRP` race to skew the reference count
   downwards, with the number of attempts chosen such that we expect that the
   resulting refcount skew is bigger than the number of references we need for
   the rest of our attack, but smaller than the number of extra references we
4. Let the task owning the `pid` exit and die, and wait for RCU
   (read-copy-update, a mechanism that involves delaying the freeing of some
   objects) to settle such that the task's reference to the `pid` is gone.
   (Waiting for an RCU grace period from userspace is not a primitive that is
   intentionally exposed through the UAPI, but there are various ways userspace
   can do it - e.g. by testing when a released BPF program's memory is
   subtracted from memory accounting, or by abusing the
   `membarrier(MEMBARRIER_CMD_GLOBAL, ...)` syscall after the kernel version
   where RCU flavors were unified.)
5. Create a new thread, and let that thread attempt to drop all the references
   we created.

Because the refcount is smaller at the start of step 5 than the number of
references we are about to drop, the `pid` will be freed at some point during
step 5; the next attempt to drop a reference will cause a use-after-free:

struct upid {
        int nr;
        struct pid_namespace *ns;

struct pid
        atomic_t count;
        unsigned int level;
        /* lists of tasks that use this pid */
        struct hlist_head tasks[PIDTYPE_MAX];
        struct rcu_head rcu;
        struct upid numbers[1];
void put_pid(struct pid *pid)
        struct pid_namespace *ns;

        if (!pid)

        ns = pid->numbers[pid->level].ns;
        if ((atomic_read(&pid->count) == 1) ||
             atomic_dec_and_test(&pid->count)) {
                kmem_cache_free(ns->pid_cachep, pid);

When the object is freed, the SLUB allocator normally replaces the first 8
bytes (sidenote: a different position is chosen starting in 5.7, see Kees'
of the freed object with an XOR-obfuscated freelist pointer; therefore, the
`count` and `level` fields are now effectively random garbage. This means that
the load from `pid->numbers[pid->level]` will now be at some random offset from
the `pid`, in the range from zero to 64 GiB. As long as the machine doesn't
have tons of RAM, this will likely cause a kernel segmentation fault. (Yes, I
know, that's an absolutely gross and unreliable way to exploit this. It mostly
works though, and I only noticed this issue when I already had the whole thing
written, so I didn't really want to go back and change it... plus, did I
mention that it mostly works?)

Linux in its default configuration, and the configuration shipped by most
general-purpose distributions, attempts to fix up unexpected kernel page faults
and other types of "oopses" by killing only the crashing thread. Therefore,
this kernel page fault is actually useful for us as a signal: Once the thread
has died, we know that the object has been freed, and can continue with the
rest of the exploit.

If this code looked a bit differently and we were actually reaching a
double-free, the SLUB allocator would also detect that and trigger a kernel
oops (see `set_freepointer()` for the `CONFIG_SLAB_FREELIST_HARDENED` case).

Discarded attack idea: Directly exploiting the UAF at the SLUB level
-------------------------------------------------------------------- On the
Debian kernel I was looking at, a `struct pid` in the initial namespace is
allocated from the same `kmem_cache` as `struct seq_file` and `struct epitem` -
these three slabs have been merged into one by `find_mergeable()` to reduce
memory fragmentation, since their object sizes, alignment requirements, and
flags match:

root@deb10:/sys/kernel/slab# ls -l pid
lrwxrwxrwx 1 root root 0 Feb  6 00:09 pid -> :A-0000128
root@deb10:/sys/kernel/slab# ls -l | grep :A-0000128
drwxr-xr-x 2 root root 0 Feb  6 00:09 :A-0000128
lrwxrwxrwx 1 root root 0 Feb  6 00:09 eventpoll_epi -> :A-0000128
lrwxrwxrwx 1 root root 0 Feb  6 00:09 pid -> :A-0000128
lrwxrwxrwx 1 root root 0 Feb  6 00:09 seq_file -> :A-0000128

A straightforward way to exploit a dangling reference to a SLUB object is to
reallocate the object through the same `kmem_cache` it came from, without ever
letting the page reach the page allocator. To figure out whether it's easy to
exploit this bug this way, I made a table listing which fields appear at each
offset in these three data structures (using `pahole -E --hex -C <typename>
<path to vmlinux debug info>`):

| offset | pid                             | eventpoll_epi / epitem **(RCU-freed)**  | seq_file                 |
| 0x00   | count.counter (4) **(CONTROL)** | rbn.__rb_parent_color (8) **(TARGET?)** | buf (8) **(TARGET?)**    |
| 0x04   | level (4)                       |                                         |                          |
| 0x08   | tasks[PIDTYPE_PID] (8)          | rbn.rb_right (8) / rcu.func (8)         | size (8)                 |
| 0x10   | tasks[PIDTYPE_TGID] (8)         | rbn.rb_left (8)                         | from (8)                 |
| 0x18   | tasks[PIDTYPE_PGID] (8)         | rdllink.next (8)                        | count (8)                |
| 0x20   | tasks[PIDTYPE_SID] (8)          | rdllink.prev (8)                        | pad_until (8)            |
| 0x28   | rcu.next (8)                    | next (8)                                | index (8)                |
| 0x30   | rcu.func (8)                    | ffd.file (8)                            | read_pos (8)             |
| 0x38   | numbers[0].nr (4)               | ffd.fd (4)                              | version (8)              |
| 0x3c   | [hole] (4)                      | nwait (4)                               |                          |
| 0x40   | numbers[0].ns (8)               | pwqlist.next (8)                        | lock (0x20): counter (8) |
| 0x48   | ---                             | pwqlist.prev (8)                        |                          |
| 0x50   | ---                             | ep (8)                                  |                          |
| 0x58   | ---                             | fllink.next (8)                         |                          |
| 0x60   | ---                             | fllink.prev (8)                         | op (8)                   |
| 0x68   | ---                             | ws (8)                                  | poll_event (4)           |
| 0x6c   | ---                             |                                         | [hole] (4)               |
| 0x70   | ---                             | event.events (4)                        | file (8)                 |
| 0x74   | ---                             | event.data (8) **(CONTROL)**            |                          |
| 0x78   | ---                             |                                         | private (8) **(TARGET?)**|
| 0x7c   | ---                             | ---                                     |                          |
| 0x80   | ---                             | ---                                     | ---                      |

In this case, reallocating the object as one of those three types didn't seem
to me like a nice way forward (although it should be possible to exploit this
somehow with some effort, e.g. by using `count.counter` to corrupt the `buf`
field of `seq_file`). Also, some systems might be using the `slab_nomerge`
kernel command line flag, which disables this merging behavior.

Another approach that I didn't look into here would have been to try to corrupt
the obfuscated SLUB freelist pointer (obfuscation is implemented in
`freelist_ptr()`); but since that stores the pointer in big-endian,
`count.counter` would only effectively let us corrupt the more significant half
of the pointer, which would probably be a pain to exploit.

Attack stage: Freeing the object's page to the page allocator
------------------------------------------------------------- This section will
refer to some internals of the SLUB allocator; if you aren't familiar with
those, you may want to at least look at slides 2-4 and 13-14 of Christoph
Lameter's slab allocator overview talk from 2014
(Note that that talk covers three different allocators; the SLUB allocator is
what most systems use nowadays.)

The alternative to exploiting the UAF at the SLUB allocator level is to flush
the page out to the page allocator (also called the buddy allocator), which is
the last level of dynamic memory allocation on Linux (once the system is far
enough into the boot process that the memblock allocator is no longer used).
>From there, the page can theoretically end up in pretty much any context. We
can flush the page out to the page allocator with the following steps:

1. Instruct the kernel to pin our task to a single CPU. Both SLUB and the page
   allocator use per-cpu structures; so if the kernel migrates us to a different
   CPU in the middle, we would fail.
2. Before allocating the victim `struct pid` whose refcount will be corrupted,
   allocate a large number of objects to drain partially-free slab pages of all
   their unallocated objects. If the victim object (which will be allocated in
   step 5 below) landed in a page that is already partially used at this point,
   we wouldn't be able to free that page.
3. Allocate around `objs_per_slab * (1+cpu_partial)` objects - in other words,
   a set of objects that completely fill at least `cpu_partial` pages, where
   `cpu_partial` is the maximum length of the "percpu partial list". Those newly
   allocated pages that are completely filled with objects are not referenced by
   SLUB's freelists at this point because SLUB only tracks pages with free
   objects on its freelists.
4. Fill `objs_per_slab-1` more objects, such that at the end of this step, the
   "CPU slab" (the page from which allocations will be served first) will not
   contain anything other than free space and fresh allocations (created in this
5. Allocate the victim object (a `struct pid`). The victim page (the page from
   which the victim object came) will usually be the CPU slab from step 4, but
   if step 4 completely filled the CPU slab, the victim page might also be a
   new, freshly allocated CPU slab.
6. Trigger the bug on the victim object to create an uncounted reference, and
   free the object.
7. Allocate `objs_per_slab+1` more objects. After this, the victim page will be
   completely filled with allocations from steps 4 and 7, and it won't be the
   CPU slab anymore (because the last allocation can not have fit into the
   victim page).
8. Free all allocations from steps 4 and 7. This causes the victim page to
   become empty, but does *not* free the page; the victim page is placed on the
   percpu partial list once a single object from that page has been freed, and
   then stays on that list.
9. Free one object per page from the allocations from step 3. This adds all
   these pages to the percpu partial list until it reaches the limit
   `cpu_partial`, at which point it will be flushed: Pages containing some
   in-use objects are placed on SLUB's per-NUMA-node partial list, and pages
   that are completely empty are freed back to the page allocator. (We don't
   free _all_ allocations from step 3 because we only want the victim page to be
   freed to the page allocator.) Note that this step requires that every
   `objs_per_slab`-th object the allocator gave us in step 3 is on a different

When the page is given to the page allocator, we benefit from the page being
order-0 (4 KiB, native page size): For order-0 pages, the page allocator has
special freelists, one per CPU+zone+migratetype combination. Pages on these
freelists are not normally accessed from other CPUs, and they don't immediately
get combined with adjacent free pages to form higher-order free pages.

At this point we are able to perform use-after-free accesses to some offset
inside the free victim page, using codepaths that interpret part of the victim
page as a `struct pid`. Note that at this point, we still don't know exactly at
which offset inside the victim page the victim object is located.

Attack stage: Reallocating the victim page as a pagetable
At the point where the victim page has reached the page allocator's freelist,
it's essentially game over - at this point, the page can be reused as anything
in the system, giving us a broad range of options for exploitation. In my
opinion, most defences that act after we've reached this point are fairly

One type of allocation that is directly served from the page allocator and has
nice properties for exploitation are page tables (which have also been used to
exploit Rowhammer, see
One way to abuse the ability to modify a page table would be to enable the
read/write bit in a page table entry (PTE) that maps a file page to which we
are only supposed to have read access - for example, this could be used to gain
write access to part of a setuid binary's `.text` segment and overwrite it with
malicious code.

We don't know at which offset inside the victim page the victim object is
located; but since a page table is effectively an array of 8-byte-aligned
elements of size 8 and the victim object's alignment is a multiple of that, as
long as we spray all elements of the victim array, we don't need to know the
victim object's offset.

To allocate a page table full of PTEs mapping the same file page, we have to:

 - prepare by setting up a 2MiB-aligned memory region (because each last-level
   page table describes 2MiB of virtual memory) containing single-page `mmap()`
   mappings of the same file page (meaning each mapping corresponds to one PTE);
 - trigger allocation of the page table and fill it with PTEs by reading from
   each mapping

`struct pid` has the same alignment as a PTE, and it starts with a 32-bit
refcount, so that refcount is guaranteed to overlap the first half of a PTE,
which is 64-bit. Because X86 CPUs are little-endian, incrementing the refcount
field in the freed `struct pid` increments the least significant half of the
PTE - so it effectively increments the PTE. (Except for the edge case where the
least significant half is `0xffffffff`, but that's not the case here.)

struct pid: count | level |   tasks[0]  |   tasks[1]  |   tasks[2]  | ... 
pagetable:       PTE      |     PTE     |     PTE     |     PTE     | ...

Therefore we can increment one of the PTEs by repeatedly triggering
`get_pid()`, which tries to increment the refcount of the freed object. This
can be turned into the ability to write to the file page as follows:

 - Increment the PTE by 0x42 to set the Read/Write bit and the Dirty bit. (If we
   didn't set the Dirty bit, the CPU would do it by itself when we write to the
   corresponding virtual address, so we could also just increment by 0x2 here.)
 - For each mapping, attempt to overwrite its contents with malicious data and
   ignore page faults.
   - This might throw spurious errors because of outdated TLB entries, but
     taking a page fault will automatically evict such TLB entries, so if we
     just attempt the write twice, this can't happen on the second write (modulo
     CPU migration, as mentioned above).
   - One easy way to ignore page faults is to let the kernel perform the memory
     write using `pread()`, which will return `-EFAULT` on fault.

If the kernel notices the Dirty bit later on, that might trigger writeback,
which could crash the kernel if the mapping isn't set up for writing.
Therefore, we have to reset the Dirty bit. We can't reliably decrement the PTE
because `put_pid()` inefficiently accesses `pid->numbers[pid->level]` even when
the refcount isn't dropping to zero, but we can increment it by an additional
0x80-0x42=0x3e, which means the final value of the PTE, compared to the initial
value, will just have the additional bit 0x80 set, which the kernel ignores.

Afterwards, we launch the setuid executable (which, in the version in the
pagecache, now contains the code we injected), and gain root privileges:

user@deb10:~/tiocspgrp$ make
as -o rootshell.o rootshell.S
ld -o rootshell rootshell.o --nmagic
gcc -Wall -o poc poc.c
user@deb10:~/tiocspgrp$ ./poc
starting up...
executing in first level child process, setting up session and PTY pair...
setting up unix sockets for ucreds spam...
draining pcpu and node partial pages
preparing for flushing pcpu partial pages
launching child process
child is 1448
ucreds spam done, struct pid refcount should be lifted. starting to skew refcount...
refcount should now be skewed, child exiting
child exited cleanly
waiting for RCU call...
bpf load with rlim 0x0: -1 (Operation not permitted)
bpf load with rlim 0x1000: 452 (Success)
bpf load success with rlim 0x1000: got fd 452
RCU callbacks executed
gonna try to free the pid...
double-free child died with signal 9 after dropping 9990 references (99%)
hopefully reallocated as an L1 pagetable now
PTE forcibly marked WRITE | DIRTY (hopefully)
clobber via corrupted PTE succeeded in page 0, 128-byte-allocation index 3, returned 856
clobber via corrupted PTE succeeded in page 0, 128-byte-allocation index 3, returned 856
bash: cannot set terminal process group (1447): Inappropriate ioctl for device
bash: no job control in this shell
root@deb10:/home/user/tiocspgrp# id
uid=0(root) gid=1000(user) groups=1000(user),24(cdrom),25(floppy),27(sudo),29(audio),30(dip),44(video),46(plugdev),108(netdev),112(lpadmin),113(scanner),120(wireshark)

Note that nothing in this whole exploit requires us to leak any kernel-virtual
or physical addresses, partly because we have an increment primitive instead of
a plain write; and it also doesn't involve directly influencing the instruction

This section describes different ways in which this exploit could perhaps have
been prevented from working. To assist the reader, the titles of some of the
subsections refer back to specific exploit stages from the section above.

Against bugs being reachable: Attack surface reduction
A potential first line of defense against many kernel security issues is to
only make kernel subsystems available to code that needs access to them. If an
attacker does not have direct access to a vulnerable subsystem _and_ doesn't
have sufficient influence over a system component with access to make it
trigger the issue, the issue is effectively unexploitable from the attacker's
security context.

Pseudoterminals are (more or less) only necessary for interactively serving
users who have shell access (or something resembling that), including:

 - terminal emulators inside graphical user sessions
 - SSH servers
 - `screen` sessions started from various types of terminals

Things like webservers or phone apps won't normally need access to such
devices; but there are exceptions. For example:

 - a web server is used to provide a remote root shell for system administration
 - a phone app's purpose is to make a shell available to the user
 - a shell script uses `expect` to interact with a binary that requires a
   terminal for input/output

In my opinion, the biggest limits on attack surface reduction as a defensive
strategy are:

1. It exposes a workaround to _an implementation concern_ of the kernel
   (potential memory safety issues) in user-facing API, which can lead to
   compatibility issues and maintenance overhead - for example, from a security
   standpoint, I think it might be a good idea to require phone apps and systemd
   services to declare their intention to use the PTY subsystem at install time,
   but that would be an API change requiring some sort of action from
   application authors, creating friction that wouldn't be necessary if we were
   confident that the kernel is working properly. This might get especially
   messy in the case of software that invokes external binaries depending on
   configuration, e.g. a web server that needs PTY access when it is used for
   server administration. (This is somewhat less complicated when a
   benign-but-potentially-exploitable application actively applies restrictions
   to itself; but not every application author is necessarily willing to design
   a fine-grained sandbox for their code, and even then, there may be
   compatibility issues caused by libraries outside the application author's
   control - see e.g. https://lwn.net/Articles/738694/.)
2. It can't protect a subsystem from a context that fundamentally needs access
   to it.  (E.g. Android's `/dev/binder` is directly accessible by Chrome
   renderers on Android because they have Android code running inside them.)

3. It means that decisions that ought to not influence the security of a system
   (making an API that does not grant extra privileges available to some
   potentially-untrusted context) essentially involve a security tradeoff.

Still, in practice, I believe that attack surface reduction mechanisms
(especially seccomp) are currently some of the most important defense
mechanisms on Linux.

Against bugs in source code: Compile-time locking validation

The bug in `TIOCSPGRP` was a fairly straightforward violation of a
straightforward locking rule: While a `tty_struct` is live, accessing its
`pgrp` member is forbidden unless the `ctrl_lock` of the same `tty_struct` is
held. This rule is sufficiently simple that it wouldn't be entirely
unreasonable to expect the compiler to be able to verify it - as long as you
somehow inform the compiler about this rule, because figuring out the intended
locking rules just from looking at a piece of code can often be hard even for
humans (especially when some of the code is incorrect).

When you are starting a new project from scratch, the overall best way to
approach this is to use a memory-safe language (see
- in other words, a language that has explicitly been designed such that the
programmer has to provide the compiler with enough information about intended
memory safety semantics that the compiler can automatically verify them. But
for existing codebases, it might be worth looking into how much of this can be

Clang's Thread Safety Analysis
(https://clang.llvm.org/docs/ThreadSafetyAnalysis.html) feature does something
vaguely like what we'd need to verify the locking in this situation:

$ nl -ba -s' ' thread-safety-test.cpp | sed 's|^   ||'
  1 struct __attribute__((capability("mutex"))) mutex {
  2 };
  4 void lock_mutex(struct mutex *p) __attribute__((acquire_capability(*p)));
  5 void unlock_mutex(struct mutex *p) __attribute__((release_capability(*p)));
  7 struct foo {
  8     int a __attribute__((guarded_by(mutex)));
  9     struct mutex mutex;
 10 };
 12 int good(struct foo *p1, struct foo *p2) {
 13     lock_mutex(&p1->mutex);
 14     int result = p1->a;
 15     unlock_mutex(&p1->mutex);
 16     return result;
 17 }
 19 int bogus(struct foo *p1, struct foo *p2) {
 20     lock_mutex(&p1->mutex);
 21     int result = p2->a;
 22     unlock_mutex(&p1->mutex);
 23     return result;
 24 }
$ clang++ -c -o thread-safety-test.o thread-safety-test.cpp -Wall -Wthread-safety
thread-safety-test.cpp:21:22: warning: reading variable 'a' requires holding mutex 'p2->mutex' [-Wthread-safety-precise]
    int result = p2->a;
thread-safety-test.cpp:21:22: note: found near match 'p1->mutex'
1 warning generated.

However, this does not currently work when compiling as C code because the
`guarded_by` attribute can't find the other struct member; it seems to have
been designed mostly for use in C++ code. A more fundamental problem is that it
also doesn't appear to have built-in support for distinguishing the different
rules for accessing a struct member depending on the lifetime state of the
object. For example, almost all objects with locked members will have
initialization/destruction functions that have exclusive access to the entire
object and can access members without locking. (The lock might not even be
initialized in those states.) 

Some objects also have more lifetime states; in particular, for many objects
with RCU-managed lifetime, only a subset of the members may be accessed through
an RCU reference without having upgraded the reference to a refcounted one
beforehand. Perhaps this could be addressed by introducing a new type attribute
that can be used to mark pointers to structs in special lifetime states? (For
C++ code, Clang's Thread Safety Analysis simply disables all checks in all
constructor/destructor functions.)

I am hopeful that, with some extensions, something vaguely like Clang's Thread
Safety Analysis could be used to retrofit some level of compile-time safety
against unintended data races. This will require adding a lot of annotations,
in particular to headers, to document intended locking semantics; but such
annotations are probably anyway necessary to enable productive work on a
complex codebase. In my experience, when there are no detailed
comments/annotations on locking rules, every attempt to change a piece of code
you're not intimately familiar with (without introducing horrible memory safety
bugs) turns into a foray into the thicket of the surrounding call graphs,
trying to unravel the intentions behind the code.

The one big downside is that this requires getting the development community
for the codebase on board with the idea of backfilling and maintaining such
annotations. And someone has to write the analysis tooling that can verify the

At the moment, the Linux kernel does have some very coarse locking validation
via `sparse`; but this infrastructure is not capable of detecting situations
where the wrong lock is used or validating that a struct member is protected by
a lock. It also can't properly deal with things like conditional locking, which
makes it hard to use for anything other than spinlocks/RCU. The kernel's
runtime locking validation via `LOCKDEP` is more advanced, but mostly with a
focus on locking correctness of RCU pointers as well as deadlock detection (the
main focus); again, there is no mechanism to, for example,automatically
validate that a given struct member is only accessed under a specific lock
(which would probably also be quite costly to implement with runtime
validation). Also, as a runtime validation mechanism, it can't discover errors
in code that isn't executed during testing (although it can combine separately
observed behavior into race scenarios without ever actually observing the

Against bugs in source code: Global static locking analysis
An alternative approach to checking memory safety rules at compile time is to
do it either after the entire codebase has been compiled, or with an external
tool that analyzes the entire codebase. This allows the analysis tooling to
perform analysis across compilation units, reducing the amount of information
that needs to be made explicit in headers. This may be a more viable approach
if peppering annotations everywhere across headers isn't viable; but it also
reduces the utility to human readers of the code, unless the inferred semantics
are made visible to them through some special code viewer. It might also be
less ergonomic in the long run if changes to one part of the kernel could make
the verification of other parts fail - especially if those failures only show
up in some configurations.

I think global static analysis is probably a good tool for finding some subsets
of bugs, and it might also help with finding the worst-case depth of kernel
stacks or proving the absence of deadlocks, but it's probably less suited for
proving memory safety correctness?

Against exploit primitives: Attack primitive reduction via syscall restrictions
(Yes, I made up that name because I thought that capturing this under "Attack
surface reduction" is too muddy.)

Because allocator fastpaths (both in SLUB and in the page allocator) are
implemented using per-CPU data structures, the ease and reliability of exploits
that want to coax the kernel's memory allocators into reallocating memory in
specific ways can be improved if the attacker has fine-grained control over the
assignment of exploit threads to CPU cores. I'm calling such a capability,
which provides a way to facilitate exploitation by influencing relevant system
state/behavior, an "attack primitive" here. Luckily for us, Linux allows tasks
to pin themselves to specific CPU cores without requiring any privilege using
the `sched_setaffinity()` syscall.

(As a different example, one primitive that can provide an attacker with fairly
powerful capabilities is being able to indefinitely stall kernel faults on
userspace addresses via
FUSE (see
or userfaultfd.)

Just like in the section "Attack surface reduction" above, an attacker's
ability to use these primitives can be reduced by filtering syscalls; but while
the mechanism and the compatibility concerns are similar, the rest is fairly

Attack primitive reduction does not normally reliably prevent a bug from being
exploited; and an attacker will sometimes even be able to obtain a similar but
shoddier (more complicated, less reliable, less generic, ...) primitive
indirectly, for example:

 - Instead of `sched_setaffinity()`, an attacker could attempt to launch several
   threads, let them poll `getcpu()` to figure out which cores they're running
   on, and then dispatch work to the threads as appropriate.
 - Instead of delaying page faults with FUSE or userfaultfd, an attacker may be
   able to abuse discontiguous file mappings and scheduler behavior - see
   https://static.sched.com/hosted_files/lsseu2019/04/LSSEU2019%20-%20Exploiting%20race%20conditions%20on%20Linux.pdf, page 30).

Attack surface reduction is about limiting access to code that is suspected to
contain exploitable bugs; in a codebase written in a memory-unsafe language,
that tends to apply to pretty much the entire codebase. Attack surface
reduction is often fairly opportunistic: You permit the things you need, and
deny the rest by default.

Attack primitive reduction limits access to code that is suspected or known to
provide (sometimes very specific) exploitation primitives. For example, one
might decide to specifically forbid access to FUSE and userfaultfd for most
code because of their utility for kernel exploitation, and, if one of those
interfaces is truly needed, design a workaround that avoids exposing the attack
primitive to userspace. This is different from attack surface reduction, where
it often makes sense to permit access to any feature that a legitimate workload
wants to use.

A nice example of an attack primitive reduction is the sysctl
`vm.unprivileged_userfaultfd`, which was first introduced
(https://git.kernel.org/linus/cefdca0a86be) so that userfaultfd can
be made completely inaccessible to normal users and was then later
adjusted(https://git.kernel.org/linus/d0d4730ac2e4) so that users can be
granted access to part of its functionality without gaining the dangerous
attack primitive.  (But if you can create unprivileged user namespaces, you
can still use FUSE to get an equivalent effect.)

When maintaining lists of allowed syscalls for a sandboxed system component, or
something along those lines, it may be a good idea to explicitly track which
syscalls are explicitly forbidden for attack primitive reduction reasons, or
similarly strong reasons - otherwise one might accidentally end up permitting
them in the future. (I guess that's kind of similar to issues that one can run
into when maintaining ACLs...)

But like in the previous section, attack primitive reduction also tends to rely
on making some functionality unavailable, and so it might not be viable in all
situations. For example, newer versions of Android deliberately indirectly give
apps access to FUSE through the AppFuse mechanism
(That API doesn't actually give an app direct access to `/dev/fuse`, but it
does forward read/write requests to the app.)

Against oops-based oracles: Lockout or panic on crash

The ability to recover from kernel oopses in an exploit can help an attacker
compensate for a lack of information about system state. Under some
circumstances, it can even serve as a binary oracle that can be used to more or
less perform a binary search for a value, or something like that.

(It used to be even worse on some distributions (see
where `dmesg` was accessible for unprivileged users; so if you managed to
trigger an oops or `WARN`, you could then grab the register states at all IRET
frames in the kernel stack, which could be used to leak things like kernel
pointers. Luckily nowadays most distributions, including Ubuntu 20.10
restrict `dmesg` access.)

Android and Chrome OS nowadays set the kernel's `panic_on_oops` flag, meaning
the machine will immediately restart when a kernel oops happens. This makes it
hard to use oopsing as part of an exploit, and arguably also makes more sense
from a reliability standpoint - the system will be down for a bit, and it will
lose its existing state, but it will also reset into a known-good state instead
of continuing in a potentially half-broken state, especially if the crashing
thread was holding mutexes that can never again be released, or things like
that. On the other hand, if some service crashes on a desktop system, perhaps
that shouldn't cause the whole system to immediately go down and make you lose
unsaved state - so `panic_on_oops` might be too drastic there.

A good solution to this might require a more fine-grained approach. (For
example, grsecurity has for a long time had the ability to lock out specific
UIDs that have caused crashes.) Perhaps it would make sense to allow the `init`
daemon to use different policies for crashes in different

Against UAF access: Deterministic UAF mitigation
One defense that would reliably stop an exploit for this issue would be a
deterministic use-after-free mitigation. Such a mitigation would reliably
protect the memory formerly occupied by the object from accesses through
dangling pointers to the object, at least once the memory has been reused for a
different purpose (including reuse to store heap metadata). For write
operations, this probably requires either atomicity of the access check and the
actual write or an RCU-like delayed freeing mechanism. For simple read
operations, it can also be implemented by ordering the access check after the
read, but before the read value is used.

A big downside of this approach on its own is that extra checks on every memory
access will probably come with an extremely high efficiency penalty, especially
if the mitigation can not make any assumptions about what kinds of parallel
accesses might be happening to an object, or what semantics pointers have. (The
proof-of-concept implementation I presented at LSSNA 2020 (slides at
recording at https://www.youtube.com/watch?v=uE1w0Mxldwk) had CPU overhead
roughly in the range 60%-159% in kernel-heavy benchmarks, and ~8% for a very
userspace-heavy benchmark.)

Unfortunately, even a deterministic use-after-free mitigation often won't be
enough to deterministically limit the blast radius of something like a
refcounting mistake to the object in which it occurred. Consider a case where
two codepaths concurrently operate on the same object: Codepath A assumes that
the object is live and subject to normal locking rules. Codepath B knows that
the reference count reached zero, assumes that it therefore has exclusive
access to the object (meaning all members are mutable without any locking
requirements), and is trying to tear down the object. Codepath B might then
start dropping references the object was holding on other objects while
codepath A is following the same references. This could then lead to
use-after-frees on pointed-to objects. If all data structures are subject to
the same mitigation, this might not be too much of a problem; but if some data
structures (like `struct page`) are not protected, it might permit a mitigation

Similar issues apply to data structures with `union` members that are used in
different object states; for example, here's some random kernel data structure
with an `rcu_head` in a `union` (just a random example, there isn't anything
wrong with this code as far as I know):

struct allowedips_node {
        struct wg_peer __rcu *peer;
        struct allowedips_node __rcu *bit[2];
        /* While it may seem scandalous that we waste space for v4,
         * we're alloc'ing to the nearest power of 2 anyway, so this
         * doesn't actually make a difference.
        u8 bits[16] __aligned(__alignof(u64));
        u8 cidr, bit_at_a, bit_at_b, bitlen;

        /* Keep rarely used list at bottom to be beyond cache line. */
        union {
                struct list_head peer_list;
                struct rcu_head rcu;

As long as everything is working properly, the `peer_list` member is only used
while the object is live, and the `rcu` member is only used after the object
has been scheduled for delayed freeing; so this code is completely fine. But
_if_ a bug somehow caused the `peer_list` to be read after the `rcu` member has
been initialized, type confusion would result.

In my opinion, this demonstrates that while UAF mitigations do have a lot of
value (and would have reliably prevented exploitation of this specific bug),
**a use-after-free is just one possible consequence of the symptom class
"object state confusion"** (which may or may not be the same as the bug class
of the root cause). It would be even better to enforce rules on object states,
and ensure that an object e.g. can't be accessed through a "refcounted"
reference anymore after the refcount has reached zero and has logically
transitioned into a state like "non-RCU members are exclusively owned by thread
performing teardown" or "RCU callback pending, non-RCU members are
uninitialized" or "exclusive access to RCU-protected members granted to thread
performing teardown, other members are uninitialized". Of course, doing this as
a runtime mitigation would be even costlier and messier than a reliable UAF
mitigation; this level of protection is probably only realistic with at least
some level of annotations and static validation.

Against UAF access: Probabilistic UAF mitigation; pointer leaks
**Summary: Some types of probabilistic UAF mitigation break if the attacker can
leak information about pointer values; and information about pointer values
easily leaks to userspace, e.g. through pointer comparisons in map/set-like

If a deterministic UAF mitigation is too costly, an alternative is to do it
probabilistically; for example, by tagging pointers with a small number of bits
that are checked against object metadata on access, and then changing that
object metadata when objects are freed.

The downside of this approach is that information leaks can be used to break
the protection. One example of a type of information leak that I'd like to
highlight (without any judgment on the relative importance of this compared to
other types of information leaks) are intentional pointer comparisons, which
have quite a few facets.

A relatively straightforward example where this could be an issue is the
[`kcmp()`](https://man7.org/linux/man-pages/man2/kcmp.2.html) syscall. This
syscall compares two kernel objects using an arithmetic comparison of their
permuted pointers (using a per-boot randomized permutation, see
`kptr_obfuscate()`) and returns the result of the comparison (smaller, equal or
greater). This gives userspace a way to order handles to kernel objects (e.g.
file descriptors) based on the identities of those kernel objects (e.g. `struct
file` instances), which in turn allows userspace to group a set of such handles
by backing kernel object in `O(n*log(n))` time using a standard sorting

This syscall can be abused for improving the reliability of use-after-free
exploits against some struct types because it checks whether two pointers to
kernel objects are equal without accessing those objects: An attacker can
allocate an object, somehow create a reference to the object that is not
counted properly, free the object, reallocate it, and then verify whether the
reallocation indeed reused the same address by comparing the dangling reference
and a reference to the new object with `kcmp()`. If `kcmp()` includes the
pointer's tag bits in the comparison, this would likely also permit breaking
probabilistic UAF mitigations.

Essentially the same concern applies when a kernel pointer is encrypted and
then given to userspace in `fuse_lock_owner_id()`, which encrypts the pointer
to a `files_struct` with an open-coded version of
XTEA (https://en.wikipedia.org/wiki/XTEA) before passing it to a FUSE daemon.

In both these cases, explicitly stripping tag bits would be an acceptable
workaround because a pointer without tag bits still uniquely identifies a
memory location; and given that these are very special interfaces that
intentionally expose some degree of information about kernel pointers to
userspace, it would be reasonable to adjust this code manually.

A somewhat more interesting example is the behavior of this piece of userspace
#define _GNU_SOURCE
#include <sys/epoll.h>
#include <sys/eventfd.h>
#include <sys/resource.h>
#include <err.h>
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
#include <sched.h>

#define SYSCHK(x) ({          \
  typeof(x) __res = (x);      \
  if (__res == (typeof(x))-1) \
    err(1, "SYSCHK(" #x ")"); \
  __res;                      \

int main(void) {
  struct rlimit rlim;
  SYSCHK(getrlimit(RLIMIT_NOFILE, &rlim));
  rlim.rlim_cur = rlim.rlim_max;
  SYSCHK(setrlimit(RLIMIT_NOFILE, &rlim));

  cpu_set_t cpuset;
  CPU_SET(0, &cpuset);
  SYSCHK(sched_setaffinity(0, sizeof(cpuset), &cpuset));

  int epfd = SYSCHK(epoll_create1(0));
  for (int i=0; i<1000; i++)
    SYSCHK(eventfd(0, 0));
  for (int i=0; i<192; i++) {
    int fd = SYSCHK(eventfd(0, 0));
    struct epoll_event event = {
      .events = EPOLLIN,
      .data = { .u64 = i }
    SYSCHK(epoll_ctl(epfd, EPOLL_CTL_ADD, fd, &event));

  char cmd[100];
  sprintf(cmd, "cat /proc/%d/fdinfo/%d", getpid(), epfd);

It first creates a ton of eventfds that aren't used. Then it creates a bunch
more eventfds and creates epoll watches for them, in creation order, with a
monotonically incrementing counter in the "data" field. Afterwards, it asks the
kernel to print the current state of the epoll instance, which comes with a
list of all registered epoll watches, including the value of the `data` member
(in hex). But how is this list sorted? Here's the result of running that code
in a Ubuntu 20.10 VM (truncated, because it's a bit long):

user@ubuntuvm:~/epoll_fdinfo$ ./epoll_fdinfo 
pos:        0
flags:        02
mnt_id:        14
tfd:     1040 events:       19 data:               24  pos:0 ino:2f9a sdev:d
tfd:     1050 events:       19 data:               2e  pos:0 ino:2f9a sdev:d
tfd:     1024 events:       19 data:               14  pos:0 ino:2f9a sdev:d
tfd:     1029 events:       19 data:               19  pos:0 ino:2f9a sdev:d
tfd:     1048 events:       19 data:               2c  pos:0 ino:2f9a sdev:d
tfd:     1042 events:       19 data:               26  pos:0 ino:2f9a sdev:d
tfd:     1026 events:       19 data:               16  pos:0 ino:2f9a sdev:d
tfd:     1033 events:       19 data:               1d  pos:0 ino:2f9a sdev:d

The `data:` field here is the loop index we stored in the `.data` member,
formatted as hex. Here is the complete list of the `data` values in decimal:

36, 46, 20, 25, 44, 38, 22, 29, 30, 45, 33, 28, 41, 31, 23, 37, 24, 50, 32, 26, 21, 43, 35, 48, 27, 39, 40, 47, 42, 34, 49, 19, 95, 105, 111, 84, 103, 97, 113, 88, 89, 104, 92, 87, 100, 90, 114, 96, 83, 109, 91, 85, 112, 102, 94, 107, 86, 98, 99, 106, 101, 93, 108, 110, 12, 1, 14, 5, 6, 9, 4, 17, 7, 13, 0, 8, 2, 11, 3, 15, 16, 18, 10, 135, 145, 119, 124, 143, 137, 121, 128, 129, 144, 132, 127, 140, 130, 122, 136, 123, 117, 131, 125, 120, 142, 134, 115, 126, 138, 139, 146, 141, 133, 116, 118, 66, 76, 82, 55, 74, 68, 52, 59, 60, 75, 63, 58, 71, 61, 53, 67, 54, 80, 62, 56, 51, 73, 65, 78, 57, 69, 70, 77, 72, 64, 79, 81, 177, 155, 161, 166, 153, 147, 163, 170, 171, 154, 174, 169, 150, 172, 164, 178, 165, 159, 173, 167, 162, 152, 176, 157, 168, 148, 149, 156, 151, 175, 158, 160, 186, 188, 179, 180, 183, 191, 181, 187, 182, 185, 189, 190, 184

While these look sort of random, you can see that the list can be split into
blocks of length 32 that consist of shuffled contiguous sequences of numbers:

Block 1 (32 values in range 19-50):
36, 46, 20, 25, 44, 38, 22, 29, 30, 45, 33, 28, 41, 31, 23, 37, 24, 50, 32, 26, 21, 43, 35, 48, 27, 39, 40, 47, 42, 34, 49, 19

Block 2 (32 values in range 83-114):
95, 105, 111, 84, 103, 97, 113, 88, 89, 104, 92, 87, 100, 90, 114, 96, 83, 109, 91, 85, 112, 102, 94, 107, 86, 98, 99, 106, 101, 93, 108, 110

Block 3 (19 values in range 0-18):
12, 1, 14, 5, 6, 9, 4, 17, 7, 13, 0, 8, 2, 11, 3, 15, 16, 18, 10

Block 4 (32 values in range 115-146):
135, 145, 119, 124, 143, 137, 121, 128, 129, 144, 132, 127, 140, 130, 122, 136, 123, 117, 131, 125, 120, 142, 134, 115, 126, 138, 139, 146, 141, 133, 116, 118

Block 5 (32 values in range 51-82):
66, 76, 82, 55, 74, 68, 52, 59, 60, 75, 63, 58, 71, 61, 53, 67, 54, 80, 62, 56, 51, 73, 65, 78, 57, 69, 70, 77, 72, 64, 79, 81

Block 6 (32 values in range 147-178):
177, 155, 161, 166, 153, 147, 163, 170, 171, 154, 174, 169, 150, 172, 164, 178, 165, 159, 173, 167, 162, 152, 176, 157, 168, 148, 149, 156, 151, 175, 158, 160

Block 7 (13 values in range 179-191):
186, 188, 179, 180, 183, 191, 181, 187, 182, 185, 189, 190, 184

What's going on here becomes clear when you look at the data structures `epoll`
uses internally. `ep_insert` calls `ep_rbtree_insert` to insert a `struct
epitem` into a red-black tree (a type of sorted binary tree); and this
red-black tree is sorted using a tuple of a `struct file *` and a file
descriptor number:

/* Compare RB tree keys */
static inline int ep_cmp_ffd(struct epoll_filefd *p1,
                             struct epoll_filefd *p2)
        return (p1->file > p2->file ? +1:
                (p1->file < p2->file ? -1 : p1->fd - p2->fd));

So the values we're seeing have been ordered based on the virtual address of
the corresponding `struct file`; and SLUB allocates `struct file` from order-1
pages (i.e. pages of size 8 KiB), which can hold 32 objects each:

root@ubuntuvm:/sys/kernel/slab/filp# cat order 
root@ubuntuvm:/sys/kernel/slab/filp# cat objs_per_slab 

This explains the grouping of the numbers we saw: Each block of 32 contiguous
values corresponds to an order-1 page that was previously empty and is used by
SLUB to allocate objects until it becomes full.

With that knowledge, we can transform those numbers a bit, to show the order in
which objects were allocated inside each page (excluding pages for which we
haven't seen all allocations):

$ cat slub_demo.py 
#!/usr/bin/env python3
blocks = [
  [ 36, 46, 20, 25, 44, 38, 22, 29, 30, 45, 33, 28, 41, 31, 23, 37, 24, 50, 32, 26, 21, 43, 35, 48, 27, 39, 40, 47, 42, 34, 49, 19 ],
  [ 95, 105, 111, 84, 103, 97, 113, 88, 89, 104, 92, 87, 100, 90, 114, 96, 83, 109, 91, 85, 112, 102, 94, 107, 86, 98, 99, 106, 101, 93, 108, 110 ],
  [ 12, 1, 14, 5, 6, 9, 4, 17, 7, 13, 0, 8, 2, 11, 3, 15, 16, 18, 10 ],
  [ 135, 145, 119, 124, 143, 137, 121, 128, 129, 144, 132, 127, 140, 130, 122, 136, 123, 117, 131, 125, 120, 142, 134, 115, 126, 138, 139, 146, 141, 133, 116, 118 ],
  [ 66, 76, 82, 55, 74, 68, 52, 59, 60, 75, 63, 58, 71, 61, 53, 67, 54, 80, 62, 56, 51, 73, 65, 78, 57, 69, 70, 77, 72, 64, 79, 81 ],
  [ 177, 155, 161, 166, 153, 147, 163, 170, 171, 154, 174, 169, 150, 172, 164, 178, 165, 159, 173, 167, 162, 152, 176, 157, 168, 148, 149, 156, 151, 175, 158, 160 ],
  [ 186, 188, 179, 180, 183, 191, 181, 187, 182, 185, 189, 190, 184 ]

for alloc_indices in blocks:
  if len(alloc_indices) != 32:
  # indices of allocations ('data'), sorted by memory location, shifted to be relative to the block
  alloc_indices_relative = [position - min(alloc_indices) for position in alloc_indices]
  # reverse mapping: memory locations of allocations,
  # sorted by index of allocation ('data').
  # if we've observed all allocations in a page,
  # these will really be indices into the page.
  memory_location_by_index = [alloc_indices_relative.index(idx) for idx in range(0, len(alloc_indices))]
$ ./slub_demo.py 
[31, 2, 20, 6, 14, 16, 3, 19, 24, 11, 7, 8, 13, 18, 10, 29, 22, 0, 15, 5, 25, 26, 12, 28, 21, 4, 9, 1, 27, 23, 30, 17]
[16, 3, 19, 24, 11, 7, 8, 13, 18, 10, 29, 22, 0, 15, 5, 25, 26, 12, 28, 21, 4, 9, 1, 27, 23, 30, 17, 31, 2, 20, 6, 14]
[23, 30, 17, 31, 2, 20, 6, 14, 16, 3, 19, 24, 11, 7, 8, 13, 18, 10, 29, 22, 0, 15, 5, 25, 26, 12, 28, 21, 4, 9, 1, 27]
[20, 6, 14, 16, 3, 19, 24, 11, 7, 8, 13, 18, 10, 29, 22, 0, 15, 5, 25, 26, 12, 28, 21, 4, 9, 1, 27, 23, 30, 17, 31, 2]
[5, 25, 26, 12, 28, 21, 4, 9, 1, 27, 23, 30, 17, 31, 2, 20, 6, 14, 16, 3, 19, 24, 11, 7, 8, 13, 18, 10, 29, 22, 0, 15]

And these sequences are almost the same, except that they have been rotated
around by different amounts. This is exactly the SLUB freelist randomization
scheme, as introduced in commit 210e7a43fa905

When a SLUB `kmem_cache` is created (an instance of the SLUB allocator for a
specific size class and potentially other specific attributes, usually
initialized at boot time), `init_cache_random_seq` and
`cache_random_seq_create` fill an array `->random_seq` with randomly-ordered
object indices via Fisher-Yates shuffle, with the array length equal to the
number of objects that fit into a page. Then, whenever SLUB grabs a new page
from the lower-level page allocator, it initializes the page freelist using the
indices from `->random_seq`, starting at a random index in the array (and
wrapping around when the end is reached). (I'm ignoring the low-order
allocation fallback here.)

So in summary, we can bypass SLUB randomization for the slab from which `struct
file` is allocated because someone used it as a lookup key in a specific type
of data structure. This is already fairly undesirable if SLUB randomization is
supposed to provide protection against some types of local attacks for all

The heap-randomization-weakening effect of such data structures is not
necessarily limited to cases where elements of the data structure can be listed
in-order by userspace: If there was a codepath that iterated through the tree
in-order and freed all tree nodes, that could have a similar effect, because
the objects would be placed on the allocator's freelist sorted by address,
cancelling out the randomization. In addition, you might be able to leak
information about iteration order through cache side channels or such.

If we introduce a probabilistic use-after-free mitigation that relies on
attackers not being able to learn whether the uppermost bits of an object's
address changed after it was reallocated, this data structure could also break
that. This case is messier than things like `kcmp()` because here the address
ordering leak stems from a standard data structure.

You may have noticed that some of the examples I'm using here would be more or
less limited to cases where an attacker is reallocating memory _with the same
type as the old allocation_, while a typical use-after-free attack ends up
replacing an object with a differently-typed one to cause type confusion. As an
example of a bug that can be exploited for privilege escalation without type
confusion at the C structure level, see entry 808 in our
bugtracker (https://bugs.chromium.org/p/project-zero/issues/detail?id=808).
My exploit for that bug first starts a `writev()` operation on a writable file,
lets the kernel validate that the file is indeed writable, then replaces the
`struct file` with a read-only `file` pointing to `/etc/crontab`, and lets
`writev()` continue. This allows gaining root privileges through a
use-after-free bug without having to mess around with kernel pointers, data
structure layouts, ROP, or anything like that. Of course that approach doesn't
work with every use-after-free though.

(By the way: For an example of pointer leaks through container data structures
in a JavaScript engine, see this bug I reported to Firefox back in 2016, when I
wasn't a Google employee:
https://thejh.net/misc/firefox-cve-2016-9904-and-cve-2017-5378-bugreport. It
leaks the low 32 bits of a pointer by timing operations on pessimal hash tables
- basically turning the HashDoS attack into an infoleak. Of course, nowadays, a
side-channel-based pointer leak in a JS engine would probably not be worth
treating as a security bug anymore, since you can probably get the same result
with Spectre...)

Against freeing SLUB pages: Preventing virtual address reuse beyond the slab
(Also discussed a little bit on the kernel-hardening list in this thread:

A weaker but less CPU-intensive alternative to trying to provide complete
use-after-free protection for individual objects would be to ensure that
_virtual_ addresses that have been used for slab memory are never reused
outside the slab, but that physical pages can still be reused. This would be
the same basic approach as used by PartitionAlloc
and others. In kernel terms, that would essentially mean serving SLUB
allocations from vmalloc space.

Some challenges I can think of with this approach are:

 - SLUB allocations are currently served from the linear mapping, which normally
   uses hugepages; if vmalloc mappings with 4K PTEs were used instead, TLB
   pressure might increase, which might lead to some performance degradation.
 - To be able to use SLUB allocations in contexts that operate directly on
   physical memory, it is sometimes necessary for SLUB pages to be physically
   contiguous. That's not really a problem, but it is different from default
   vmalloc behavior. (Sidenote: DMA buffers don't always have to be physically
   contiguous - if you have an IOMMU, you can use that to map discontiguous
   pages to a contiguous DMA address range, just like how normal page tables
   create virtually-contiguous memory. See the kernel-internal API introduced in
   https://git.kernel.org/linus/7d5b5738d151 for an example that makes use of
   this, and Fuchsia's documentation
   (https://fuchsia.dev/fuchsia-src/concepts/drivers/driver_development/dma) for
   a high-level overview of how all this works in general.)
 - Some parts of the kernel convert back and forth between virtual addresses,
   `struct page` pointers, and (for interaction with hardware) physical
   addresses. This is a relatively straightforward mapping for addresses in the
   linear mapping, but would become a bit more complicated for vmalloc
   addresses. In particular, `page_to_virt()` and `phys_to_virt()` would have to
   be adjusted.
   - This is probably also going to be an issue for things like Memory Tagging,
     since pointer tags will have to be reconstructed when converting back to a
     virtual address. Perhaps it would make sense to forbid these helpers
     outside low-level memory management, and change existing users to instead
     keep a normal pointer to the allocation around? Or maybe you could let
     pointers to `struct page` carry the tag bits for the corresponding virtual
     address in unused/ignored address bits?

The probability that this defense can prevent UAFs from leading to exploitable
type confusion depends somewhat on the granularity of slabs; if specific struct
types have their own slabs, it provides more protection than if objects are
only grouped by size. So to improve the utility of virtually-backed slab
memory, it would be necessary to replace the generic kmalloc slabs (which
contain various objects, grouped only by size) with ones that are segregated by
type and/or allocation site. (The grsecurity/PaX folks have vaguely alluded to
doing something roughly along these lines using compiler instrumentation.)

After reallocation as pagetable: Structure layout randomization
Memory safety issues are often exploited in a way that involves creating a type
confusion; e.g. exploiting a use-after-free by replacing the freed object with
a new object of a different type.

A defense that first appeared in grsecurity/PaX is to shuffle the order of
struct members at build time to make it harder to exploit type confusions
involving structs; the upstream Linux version of this is in

How effective this is depends partly on whether the attacker is forced to
exploit the issue as a confusion between two structs, or whether the attacker
can instead exploit it as a confusion between a struct and an array (e.g.
containing characters, pointers or PTEs). Especially if only a single struct
member is accessed, a struct-array confusion might still be viable by spraying
the entire array with identical elements. Against the type confusion described
in this blogpost (between `struct pid` and page table entries), structure
layout randomization could still be somewhat effective, since the reference
count is half the size of a PTE and therefore can randomly be placed to overlap
either the lower or the upper half of a PTE. (Except that the upstream Linux
version of randstruct only randomizes explicitly-marked structs or structs
containing only function pointers, and `struct pid` has no such marking.)

Of course, drawing a clear distinction between structs and arrays
oversimplifies things a bit; for example, there might be struct types that have
a large number of pointers of the same type or attacker-controlled values, not
unlike an array.

*If* the attacker can not completely sidestep structure layout randomization by
spraying the entire struct, the level of protection depends on how kernel
builds are distributed:

 - If the builds are created centrally by one vendor and distributed to a large
   number of users, an attacker who wants to be able to compromise users of this
   vendor would have to rework their exploit to use a different type confusion
   for each release, which may force the attacker to rewrite significant chunks
   of the exploit.
 - If the kernel is individually built per machine (or similar), and the kernel
   image is kept secret, an attacker who wants to reliably exploit a target
   system may be forced to somehow leak information about some structure layouts
   and either prepare exploits for many different possible struct layouts in
   advance or write parts of the exploit interactively after leaking information
   from the target system.

To maximize the benefit of structure layout randomization in an environment
where kernels are built centrally by a distribution/vendor, it would be
necessary to make randomization a boot-time process by making structure offsets
relocatable. (Or install-time, but that would break code signing.) Doing this
cleanly (for example, such that 8-bit and 16-bit immediate displacements can
still be used for struct member access where possible) would probably require a
lot of fiddling with compiler internals, from the C frontend all the way to the
emission of relocations. A somewhat hacky version of this approach already
exists for C->BPF compilation as BPF CO-RE]
using the clang builtin `__builtin_preserve_access_index`
but that relies on debuginfo, which probably isn't a very clean approach.

Potential issues with structure layout randomization are:

 - If structures are hand-crafted to be particularly cache-efficient, fully
   randomizing structure layout could worsen cache behavior. The existing
   randstruct implementation optionally avoids this by trying to randomize only
   within a cache line.
 - Unless the randomization is applied in a way that is reflected in DWARF debug
   info and such (which it isn't in the existing GCC-based implementation, see
   https://gcc.gnu.org/bugzilla/show_bug.cgi?id=84052), it can make debugging
   and introspection harder.
 - It can break code that makes assumptions about structure layout; but such
   code is gross and should be cleaned up anyway (and Gustavo Silva has been
   working (https://github.com/KSPP/linux/issues/109) on fixing some of those

While structure layout randomization by itself is limited in its effectiveness
by struct-array confusions, it might be more reliable in combination with
limited heap partitioning: If the heap is partitioned such that only
struct-struct confusion is possible, and structure layout randomization makes
struct-struct confusion difficult to exploit, and no struct in the same heap
partition has array-like properties, then it would probably become much harder
to directly exploit a UAF as type confusion. On the other hand, if the heap is
already partitioned like that, it might make more sense to go all the way with
heap partitioning and create one partition per type instead of dealing with all
the hassle of structure layout randomization.

(By the way, if structure layouts are randomized, padding should probably also
be randomized explicitly instead of always being on the same side to maximally
randomize structure members with low alignment; see my list post on this topic
for details.)

Control Flow Integrity
**I want to explicitly point out that kernel Control Flow Integrity would have
had no impact at all on this exploit strategy**. By using a data-only strategy,
we avoid having to leak addresses, avoid having to find ROP gadgets for a
specific kernel build, and are completely unaffected by any defenses that
attempt to protect kernel code or kernel control flow. Things like getting
access to arbitrary files, increasing the privileges of a process, and so on
don't require kernel instruction pointer control.

Like in my last blogpost on Linux kernel exploitation
(which was about a buggy subsystem that an Android vendor added to their
downstream kernel), to me, a data-only approach to exploitation feels very
natural and seems less messy than trying to hijack control flow anyway.

Maybe things are different for userspace code; but for attacks by userspace
against the kernel, I don't currently see a lot of utility in CFI because it
typically only affects one of many possible methods for exploiting a bug.
(Although of course there could be specific cases where a bug can only be
exploited by hijacking control flow, e.g. if a type confusion only permits
overwriting a function pointer and none of the permitted callees make
assumptions about input types or privileges that could be broken by changing
the function pointer.)

Making important data readonly

A defense idea that has shown up in a bunch of places (including Samsung phone
kernels and XNU kernels for iOS) is to make data that is crucial to kernel
security read-only except when it is intentionally being written to - the idea
being that even if an attacker has an arbitrary memory write, they should not
be able to directly overwrite specific pieces of data that are of exceptionally
high importance to system security, such as credential structures, page tables,
or (on iOS, using PPL) userspace code pages:

The problem I see with this approach is that a large portion of the things a
kernel does are, in some way, critical to the correct functioning of the system
and system security. MMU state management, task scheduling, memory allocation,
page cache, IPC, ... - if any one of these parts of the kernel is corrupted
sufficiently badly, an attacker will probably be able to gain access to all
user data on the system, or use that corruption to feed bogus inputs into one
of the subsystems whose own data structures are read-only.

In my view, instead of trying to split out the most critical parts of the
kernel and run them in a context with higher privileges, it might be more
productive to go in the opposite direction and try to approximate something
like a proper microkernel: Split out drivers that don't strictly need to be in
the kernel and run them in a lower-privileged context that interacts with the
core kernel through proper APIs. Of course that's easier said than done! But
Linux does already have APIs for safely accessing PCI devices (VFIO) and USB
devices from userspace, although userspace drivers aren't exactly its main

(One might also consider making page tables read-only not because of their
importance to system integrity, but because the structure of page table entries
makes them nicer to work with in exploits that are constrained in what
modifications they can make to memory. I dislike this approach because I think
it has no clear conclusion and it is highly invasive regarding how data
structures can be laid out.)

This was essentially a boring locking bug in some random kernel subsystem that,
if it wasn't for memory unsafety, shouldn't really have much of a relevance to
system security. I wrote a fairly straightforward, unexciting (and admittedly
unreliable) exploit against this bug; and probably the biggest challenge I
encountered when trying to exploit it on Debian was to properly understand how
the SLUB allocator works.

My intent in describing the exploit stages, and how different mitigations might
affect them, **is to highlight that the further a memory corruption exploit
progresses, the more options an attacker gains; and so as a general rule, the
earlier an exploit is stopped, the more reliable the defense is. Therefore,
even if defenses that stop an exploit at an earlier point have higher overhead,
they might still be more useful**.

I think that the current situation of software security could be dramatically
improved - in a world where a little bug in some random kernel subsystem can
lead to a full system compromise, the kernel can't provide reliable security
isolation. Security engineers should be able to focus on things like buggy
permission checks and core memory management correctness, and not have to spend
their time dealing with issues in code that ought to not have any relevance to
system security.

In the short term, there are some band-aid mitigations that could be used to
improve the situation - like heap partitioning or fine-grained UAF mitigation.
These might come with some performance cost, and that might make them look
unattractive; but I still think that they're a better place to invest
development time than things like CFI, which attempts to protect against much
later stages of exploitation.

In the long term, I think something has to change about the programming
language - plain C is simply too error-prone. Maybe the answer is Rust; or
maybe the answer is to introduce enough annotations to C (along the lines of
Microsoft's Checked C project
(https://www.microsoft.com/en-us/research/project/checked-c/), although as far
as I can see they mostly focus on things like array bounds rather than temporal
issues) to allow Rust-equivalent build-time verification of locking rules,
object states, refcounting, void pointer casts, and so on. Or maybe another
completely different memory-safe language will become popular in the end,
neither C nor Rust?

My hope is that perhaps in the mid-term future, we could have a statically
verified, high-performance core of kernel code working together with
instrumented, runtime-verified, non-performance-critical legacy code, such that
developers can make a tradeoff between investing time into backfilling correct
annotations and run-time instrumentation slowdown without compromising on
security either way.

**memory corruption is a big problem because small bugs even outside
security-related code can lead to a complete system compromise; and to address
that, it is important that we:**

 - in the short to medium term:
   - **design new memory safety mitigations**:
     - ideally, that can stop attacks at an early point where attackers don't
       have a lot of alternate options yet
       - maybe at the memory allocator level (i.e. SLUB)
     - that can't be broken using address tag leaks (or we try to prevent tag
       leaks, but that's really hard)
   - **continue using attack surface reduction**
     - in particular seccomp
   - **explicitly prevent untrusted code from gaining important attack
     - like FUSE, and potentially consider fine-grained scheduler control

 - in the long term:
   - **statically verify correctness of most performance-critical code**
     - this will require determining how to retrofit annotations for object
       state and locking onto legacy C code
     - consider designing runtime verification just for gaps in static

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