| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
af_unix: Fix UAF read of tail->len in unix_stream_data_wait()
unix_stream_data_wait() does skb_peek_tail(&sk->sk_receive_queue) without
holding any lock that prevents SKBs on that queue from being dequeued and
freed.
This has been the case since commit 79f632c71bea ("unix/stream: fix
peeking with an offset larger than data in queue").
The first consequence of this is that the pointer comparison
`tail != last` can be false even if `last` semantically refers to an
already-freed SKB while `tail` is a new SKB allocated at the same address;
which can cause unix_stream_data_wait() to wrongly keep blocking after new
data has arrived, but only in a weird scenario where a peeking recv() and
a normal recv() on the same socket are racing, which is probably not a
real problem.
But since commit 2b514574f7e8 ("net: af_unix: implement splice for stream
af_unix sockets"), `tail` is actually dereferenced, which can cause UAF in
the following race scenario (where test_setup() runs single-threaded,
and afterwards, test_thread1() and test_thread2() run concurrently in
two threads:
```
static int socks[2];
void test_setup(void) {
socketpair(AF_UNIX, SOCK_STREAM, 0, socks);
send(socks[1], "A", 1, 0);
int peekoff = 1;
setsockopt(socks[0], SOL_SOCKET, SO_PEEK_OFF, &peekoff, sizeof(peekoff));
}
void test_thread1(void) {
char dummy;
recv(socks[0], &dummy, 1, MSG_PEEK);
}
void test_thread2(void) {
char dummy;
recv(socks[0], &dummy, 1, 0);
shutdown(socks[1], SHUT_WR);
}
```
when racing like this:
```
thread1 thread2
unix_stream_read_generic
mutex_lock(&u->iolock)
skb_peek(&sk->sk_receive_queue)
skb_peek_next(skb, &sk->sk_receive_queue)
mutex_unlock(&u->iolock)
unix_stream_read_generic
unix_state_lock(sk)
skb_peek(&sk->sk_receive_queue)
unix_state_unlock(sk)
unix_stream_data_wait
unix_state_lock(sk)
tail = skb_peek_tail(&sk->sk_receive_queue)
spin_lock(&sk->sk_receive_queue.lock)
__skb_unlink(skb, &sk->sk_receive_queue)
spin_unlock(&sk->sk_receive_queue.lock)
consume_skb(skb) [frees the SKB]
`tail != last`: false
`tail`: true
`tail->len != last_len` ***UAF***
```
Fix the UAF by removing the read of tail->len; checking tail->len would
only make sense if SKBs in the receive queue of a UNIX socket could grow,
which can no longer happen.
Kuniyuki explained:
> When commit 869e7c62486e ("net: af_unix: implement stream sendpage
> support") added sendpage() support, data could be appended to the last
> skb in the receiver's queue.
>
> That's why we needed to check if the length of the last skb was changed
> while waiting for new data in unix_stream_data_wait().
>
> However, commit a0dbf5f818f9 ("af_unix: Support MSG_SPLICE_PAGES") and
> commit 57d44a354a43 ("unix: Convert unix_stream_sendpage() to use
> MSG_SPLICE_PAGES") refactored sendmsg(), and now data is always added
> to a new skb.
That means this fix is not suitable for kernels before 6.5. |
| In the Linux kernel, the following vulnerability has been resolved:
batman-adv: bla: avoid double decrement of bla.num_requests
The bla.num_requests is increased when no request_sent was in progress. And
it is decremented in various places (announcement was received, backbone is
purged, periodic work). But the check if the request_sent is actually set
to a specific state and the atomic_dec/_inc are not safe because they are
not atomic (TOCTOU) and multiple such code portions can run concurrently.
At the same time, it is necessary to modify request_sent (state) and
bla.num_requests atomically. Otherwise batadv_bla_send_request() might set
request_sent to 1 and is interrupted. batadv_handle_announce() can then
set request_sent back to 0 and decrement num_requests before
batadv_bla_send_request() incremented it.
The two operations must therefore be locked. And since state (request_sent)
and wait_periods are only accessed inside this lock, they can be converted
to simpler datatypes. And to avoid that the bla.num_requests is touched by
a parallel running context with a valid backbone_gw reference after
batadv_bla_purge_backbone_gw() ran, a third state "stopped" is required to
correctly signal that a backbone_gw is in the state of being cleaned up. |
| In the Linux kernel, the following vulnerability has been resolved:
batman-adv: tt: fix TOCTOU race for reported vlans
The local TT based TVLV is generated by first checking the number of VLANs
which have at least one TT entry. A new buffer with the correct size for
the VLANs is then allocated. Only then, the list of VLANs s used to fill
the VLAN entries in the buffer. During this time, the meshif_vlan_list_lock
is held. But the actual number of TT entries of each VLAN can still
increase during this time - just not the number of VLANs in the list.
But the prefilter used in the buffer size calculation might still cause an
increase of the number of VLANs which need to be stored. Simply because a
VLAN might now suddenly have at least one entry when it had none in the
pre-alloc check - and then needs to occupy space which was not allocated.
It is better to overestimate the buffer size at the beginning and then fill
the buffer only with the VLANs which are not empty. |
| In the Linux kernel, the following vulnerability has been resolved:
irq_work: Fix use-after-free in irq_work_single() on PREEMPT_RT
On PREEMPT_RT, non-HARD irq_work runs in per-CPU kthreads via
run_irq_workd(), so irq_work_sync() uses rcuwait() to wait for BUSY==0.
After irq_work_single() clears BUSY via atomic_cmpxchg(), it still
dereferences @work for irq_work_is_hard() and rcuwait_wake_up().
An irq_work_sync() caller on another CPU that enters after BUSY is cleared
can observe BUSY==0 immediately, return, and free the work before those
accesses complete — causing a use-after-free.
Fix this by wrapping run_irq_workd() in guard(rcu)() so that the entire
irq_work_single() execution is within an RCU read-side critical
section. Then add synchronize_rcu() in irq_work_sync() after
rcuwait_wait_event() to ensure the caller waits for the RCU grace period
before returning, preventing premature frees. |
| In the Linux kernel, the following vulnerability has been resolved:
netfs: Fix missing barriers when accessing stream->subrequests locklessly
The list of subrequests attached to stream->subrequests is accessed without
locks by netfs_collect_read_results() and netfs_collect_write_results(),
and then they access subreq->flags without taking a barrier after getting
the subreq pointer from the list. Relatedly, the functions that build the
list don't use any sort of write barrier when constructing the list to make
sure that the NETFS_SREQ_IN_PROGRESS flag is perceived to be set first if
no lock is taken.
Fix this by:
(1) Add a new list_add_tail_release() function that uses a release barrier
to set the pointer to the new member of the list.
(2) Add a new list_first_entry_or_null_acquire() function that uses an
acquire barrier to read the pointer to the first member in a list (or
return NULL).
(3) Use list_add_tail_release() when adding a subreq to ->subrequests.
(4) Use list_first_entry_or_null_acquire() when initially accessing the
front of the list (when an item is removed, the pointer to the new
front iterm is obtained under the same lock). |
| In the Linux kernel, the following vulnerability has been resolved:
netfs: Fix netfs_read_folio() to wait on writeback
Fix netfs_read_folio() to wait for an ongoing writeback to complete so that
it can trust the dirty flag and whatever is attached to folio->private
(folio->private may get cleaned up by the collector before it clears the
writeback flag). |
| In the Linux kernel, the following vulnerability has been resolved:
afs: Fix the locking used by afs_get_link()
The afs filesystem in the kernel doesn't do locking correctly for symbolic
links. There are a number of problems:
(1) It doesn't do any locking around afs_read_single() to prevent races
between multiple ->get_link() calls, thereby allowing the possibility
of leaks.
(2) It doesn't use RCU barriering when accessing the buffer pointers
during RCU pathwalk.
(3) It can race with another thread updating the contents of the symlink
if a third party updated it on the server.
Fix this by the following means:
(0) Move symlink handling into its own file as this makes it more
complicated.
(1) Take the validate_lock around afs_read_single() to prevent races
between multiple ->get_link() calls.
(2) Keep a separate copy of the symlink contents with an rcu_head. This
is always going to be a lot smaller than a page, so it can be
kmalloc'd and save quite a bit of memory. It also needs a refcount
for non-RCU pathwalk.
(3) Split the symlink read and write-to-cache routines in afs from those
for directories.
(4) Discard the I/O buffer as soon as the write-to-cache completes as this
is a full page (plus a folio_queue).
(5) If there's no cache, discard the I/O buffer immediately after reading
and copying if there is no cache. |
| In the Linux kernel, the following vulnerability has been resolved:
ovpn: tcp - use cached peer pointer in ovpn_tcp_close()
ovpn_tcp_close() loads the ovpn_socket via rcu_dereference_sk_user_data()
under rcu_read_lock(), takes a reference on sock->peer, caches the peer
pointer in a local, and drops the read lock. It then passes sock->peer
(rather than the cached local) to ovpn_peer_del(), re-dereferencing the
ovpn_socket after the RCU read section has ended.
Unlike ovpn_tcp_sendmsg(), which uses the same "load under RCU, use
after unlock" pattern but is protected by lock_sock() held across the
function, ovpn_tcp_close() runs without the socket lock: inet_release()
invokes sk_prot->close() without taking lock_sock first.
ovpn_socket_release() can therefore complete its kref_put -> detach ->
synchronize_rcu -> kfree(sock) sequence concurrently, in the window
after ovpn_tcp_close() drops rcu_read_lock() but before it dereferences
sock->peer. The synchronize_rcu() in ovpn_socket_release() protects
readers that use the dereferenced pointer inside the RCU read section,
not those that escape the pointer to a local and use it afterwards.
A reproducer follows the pattern of commit 94560267d6c4 ("ovpn: tcp -
don't deref NULL sk_socket member after tcp_close()"): trigger a peer
removal (keepalive expiration or netlink OVPN_CMD_DEL_PEER) at the same
moment userspace closes the TCP fd. That commit fixed the detach-side
of the same race window; this one fixes the close-side at a different
victim.
Tighten the entry block to read sock->peer exactly once into the cached
peer local, and route all subsequent uses (the hold check, the
ovpn_peer_del() call, and the prot->close() invocation) through that
local. sock->peer is only ever written once in ovpn_socket_new() under
lock_sock(), before rcu_assign_sk_user_data() publishes the ovpn_socket,
and is never reassigned afterwards - but the previous multi-read pattern
made that invariant implicit rather than explicit. The same multi-read
shape exists in ovpn_tcp_recvmsg(), ovpn_tcp_sendmsg(),
ovpn_tcp_data_ready() and ovpn_tcp_write_space(); those will be cleaned
up via a dedicated helper in a follow-up net-next series. |
| In the Linux kernel, the following vulnerability has been resolved:
Revert "f2fs: remove non-uptodate folio from the page cache in move_data_block"
This reverts commit 9609dd704725a40cd63d915f2ab6c44248a44598.
The kernel panics are keeping to be reported especially when the f2fs
partition get almost full. By investigation, we find that the reason is
one f2fs page got freed to buddy without being deleted from LRU and the
root cause is the race happened in [2] which is enrolled by this commit.
There are 3 race processes in this scenario, please find below for their
main activities.
The changed code in move_data_block() lets the GC path evict the tail-end
folio from the page cache through folio_end_dropbehind(). Once
folio_unmap_invalidate() removes the folio from mapping->i_pages, the
page-cache references for all pages in the folio are dropped. The folio
is then kept alive only by temporary external references, which allows a
later split to operate on a folio whose subpages are no longer protected
by page-cache references.
After the page-cache references are gone, split_folio_to_order() can
split the big folio into individual pages and put the resulting subpages
back on the LRU. For tail pages beyond EOF, split removes them from the
page cache and drops their page-cache references. A tail page can then
remain on the LRU with PG_lru set while holding only the split caller's
temporary reference. When free_folio_and_swap_cache() drops that final
reference, the page enters the final folio_put() release path.
In parallel, folio_isolate_lru() can observe the same tail page with a
non-zero refcount and PG_lru set. It clears PG_lru before taking its own
reference. If this races with the final folio_put() from the split path,
__folio_put() sees PG_lru already cleared and skips lruvec_del_folio().
The page is then freed back to the allocator while its lru links are
still present in the LRU list. A later LRU operation on a neighboring
page detects the stale link and reports list corruption.
[1]
[ 22.486082] list_del corruption. next->prev should be fffffffec10e0ac8, but was dead000000000122. (next=fffffffec10e0a88)
[ 22.486130] ------------[ cut here ]------------
[ 22.486134] kernel BUG at lib/list_debug.c:67!
[ 22.486141] Internal error: Oops - BUG: 00000000f2000800 [#1] SMP
[ 22.488502] Tainted: [W]=WARN, [O]=OOT_MODULE
[ 22.488506] Hardware name: Spreadtrum UMS9230 1H10 SoC (DT)
[ 22.488511] pstate: 604000c5 (nZCv daIF +PAN -UAO -TCO -DIT -SSBS BTYPE=--)
[ 22.488517] pc : __list_del_entry_valid_or_report+0x14c/0x154
[ 22.488531] lr : __list_del_entry_valid_or_report+0x14c/0x154
[ 22.488539] sp : ffffffc08006b830
[ 22.488542] x29: ffffffc08006b868 x28: 0000000000003020 x27: 0000000000000000
[ 22.488553] x26: 0000000000000000 x25: 0000000000000004 x24: fffffffec10e0ac0
[ 22.488564] x23: 00000000000000e8 x22: 0000000000000024 x21: dead000000000122
[ 22.488574] x20: fffffffec10e0a88 x19: fffffffec10e0ac8 x18: ffffffc080061060
[ 22.488585] x17: 20747562202c3863 x16: 6130653031636566 x15: 0000000000000058
[ 22.488595] x14: 0000000000000004 x13: ffffff80f91e0000 x12: 0000000000000003
[ 22.488605] x11: 0000000000000003 x10: 0000000000000001 x9 : ffe85721f0e25f00
[ 22.488615] x8 : ffe85721f0e25f00 x7 : 0000000000000000 x6 : 6c65645f7473696c
[ 22.488625] x5 : ffffffed39b23026 x4 : 0000000000000000 x3 : 0000000000000010
[ 22.488636] x2 : 0000000000000000 x1 : 0000000000000000 x0 : 000000000000006d
[ 22.488647] Call trace:
[ 22.488651] __list_del_entry_valid_or_report+0x14c/0x154 (P)
[ 22.488661] __folio_put+0x2bc/0x434
[ 22.488670] folio_put+0x28/0x58
[ 22.488678] do_garbage_collect+0x1a34/0x2584
[ 22.488689] f2fs_gc+0x230/0x9b4
[ 22.488697] f2fs_fallocate+0xb90/0xdf4
[ 22.488706] vfs_fallocate+0x1b4/0x2bc
[ 22.488716] __arm64_sys_fallocate+0x44/0x78
[ 22.488725] invoke_syscall+0x58/0xe4
[ 22.488732] do_el0_svc+0x48/0xdc
[ 22.488739] el0
---truncated--- |
| In the Linux kernel, the following vulnerability has been resolved:
ntfs: serialize volume label accesses
Protect vol->volume_label with a mutex and snaphost the label before
copy_to_user. This prevent a use-after-free when FS_IOC_SETFSLABEL
replaces the vol->volume_label and FS_IOC_GETTSLABEL reads it
concurrently. |
| In the Linux kernel, the following vulnerability has been resolved:
f2fs: fix fsck inconsistency caused by incorrect nat_entry flag usage
f2fs_need_dentry_mark() reads nat_entry flags without mutual exclusion
with the checkpoint path, which can result in an incorrect inode block
marking state. The scenario is as follows:
create & write & fsync 'file A' write checkpoint
- f2fs_do_sync_file // inline inode
- f2fs_write_inode // inode folio is dirty
- f2fs_write_checkpoint
- f2fs_flush_merged_writes
- f2fs_sync_node_pages
- f2fs_fsync_node_pages // no dirty node
- f2fs_need_inode_block_update // return true
- f2fs_fsync_node_pages // inode dirtied
- f2fs_need_dentry_mark //return true
- f2fs_flush_nat_entries
- f2fs_write_checkpoint end
- __write_node_folio // inode with DENT_BIT_SHIFT set
SPO, "fsck --dry-run" find inode has already checkpointed but still
with DENT_BIT_SHIFT set
The state observed by f2fs_need_dentry_mark() can differ from the state
observed in __write_node_folio() after acquiring sbi->node_write. The
root cause is that the semantics of IS_CHECKPOINTED and
HAS_FSYNCED_INODE are only guaranteed after the checkpoint write has
fully completed.
This patch moves set_dentry_mark() into __write_node_folio() and
protects it with the sbi->node_write lock. |
| IBM Cognos Analytics 12.1.3 GA Version with build number through 12.1.3-2606251736 could allow an attacker to obtain incorrect report summary results or cause report-processing failures due to a race condition in the Agentic AI assistant's concurrent request-handling logic when multiple authenticated users submit report-related tasks simultaneously. |
| In the Linux kernel, the following vulnerability has been resolved:
netfs, afs: Fix write skipping in dir/link writepages
Fix netfs_write_single() and afs_single_writepages() to better handle a
write that would be skipped due to lock contention and WB_SYNC_NONE by
returning 1 from netfs_write_single() if it skipped and making
afs_single_writepages() skip also. If a skip occurs, the inode must be
re-marked as the VFS may have cleared the mark.
This is really only theoretical for directories in netfs_write_single() as
the only path to that is through afs_single_writepages() that takes the
->validate_lock around it, thereby serialising it. |
| In the Linux kernel, the following vulnerability has been resolved:
netfs: Fix streaming write being overwritten
In order to avoid reading whilst writing, netfslib will allow "streaming
writes" in which dirty data is stored directly into folios without reading
them first. Such folios are marked dirty but may not be marked uptodate.
If a folio is entirely written by a streaming write, uptodate will be set,
otherwise it will have a netfs_folio struct attached to ->private recording
the dirty region.
In the event that a partially written streaming write page is to be
overwritten entirely by a single write(), netfs_perform_write() will try to
copy over it, but doesn't discard the netfs_folio if it succeeds; further,
it doesn't correctly handle a partial copy that overwrites some of the
dirty data.
Fix this by the following:
(1) If the folio is successfully overwritten, free the netfs_folio struct
before marking the page uptodate.
(2) If the copy to the folio partially fails, but short of the dirty data,
just ignore the copy.
(3) If the copy partially fails and overwrites some of the dirty data,
accept the copy, update the netfs_folio struct to record the new data.
If the folio is now filled, free the netfs_folio and set uptodate,
otherwise return a partial write.
Found with:
fsx -q -N 1000000 -p 10000 -o 128000 -l 600000 \
/xfstest.test/junk --replay-ops=junk.fsxops
using the following as junk.fsxops:
truncate 0x0 0 0x927c0
write 0x63fb8 0x53c8 0
copy_range 0xb704 0x19b9 0x24429 0x79380
write 0x2402b 0x144a2 0x90660 *
write 0x204d5 0x140a0 0x927c0 *
copy_range 0x1f72c 0x137d0 0x7a906 0x927c0 *
read 0x00000 0x20000 0x9157c
read 0x20000 0x20000 0x9157c
read 0x40000 0x20000 0x9157c
read 0x60000 0x20000 0x9157c
read 0x7e1a0 0xcfb9 0x9157c
on cifs with the default cache option.
It shows folio 0x24 misbehaving if the FMODE_READ check is commented out in
netfs_perform_write():
if (//(file->f_mode & FMODE_READ) ||
netfs_is_cache_enabled(ctx)) {
and no fscache. This was initially found with the generic/522 xfstest. |
| In the Linux kernel, the following vulnerability has been resolved:
batman-adv: tp_meter: avoid divide-by-zero for dec_cwnd
The cwnd is always MSS <= cwnd <= 0x20000000. But the calculation in
batadv_tp_update_cwnd() assumes unsigned 32 bit arithmetics.
((mss * 8) ** 2) / (cwnd * 8)
In case cwnd is actually 0x20000000, it will be shifted by 3 bit to the
left end up at 0x100000000 or U32_MAX + 1. It will therefore wrap around
and be 0 - resulting in:
((mss * 8) ** 2) / 0
This is of course invalid and cannot be calculated. The calculation should
must be simplified to avoid this overflow:
(mss ** 2) * 8 / cwnd
It will keep the precision enhancement from the scaling (by 8) but avoid
the overflow in the divisor.
In theory, there could still be an overflow in the dividend. It is at the
moment fixed to BATADV_TP_PLEN in batadv_tp_recv_ack() - so it is not an
imminent problem. But allowing it to use the whole u32 bit range, would
mean that it can still use up to 67 bits. To keep this calculation safe for
32 bit arithmetic, mss must never use more than floor((32 - 3) / 2) bits -
or in other words: must never be larger than 16383. |
| In the Linux kernel, the following vulnerability has been resolved:
iommu/vt-d: Fix race condition during PASID entry replacement
The Intel VT-d PASID table entry is 512 bits (64 bytes). When replacing
an active PASID entry (e.g., during domain replacement), the current
implementation calculates a new entry on the stack and copies it to the
table using a single structure assignment.
struct pasid_entry *pte, new_pte;
pte = intel_pasid_get_entry(dev, pasid);
pasid_pte_config_first_level(iommu, &new_pte, ...);
*pte = new_pte;
Because the hardware may fetch the 512-bit PASID entry in multiple
128-bit chunks, updating the entire entry while it is active (Present
bit set) risks a "torn" read. In this scenario, the IOMMU hardware
could observe an inconsistent state — partially new data and partially
old data — leading to unpredictable behavior or spurious faults.
Fix this by removing the unsafe "replace" helpers and following the
"clear-then-update" flow, which ensures the Present bit is cleared and
the required invalidation handshake is completed before the new
configuration is applied. |
| In the Linux kernel, the following vulnerability has been resolved:
af_unix: Set gc_in_progress to true in unix_gc().
Igor Ushakov reported that unix_gc() could run with gc_in_progress
being false if the work is scheduled while running:
Thread 1 Thread 2 Thread 3
-------- -------- --------
unix_schedule_gc() unix_schedule_gc()
`- if (!gc_in_progress) `- if (!gc_in_progress)
|- gc_in_progress = true |
`- queue_work() |
unix_gc() <----------------/ |
| |- gc_in_progress = true
... `- queue_work()
| |
`- gc_in_progress = false |
|
unix_gc() <---------------------------------------------'
|
... /* gc_in_progress == false */
|
`- gc_in_progress = false
unix_peek_fpl() relies on gc_in_progress not to confuse GC
by MSG_PEEK.
Let's set gc_in_progress to true in unix_gc(). |
| In the Linux kernel, the following vulnerability has been resolved:
Bluetooth: fix UAF in l2cap_sock_cleanup_listen() vs l2cap_conn_del()
bt_accept_dequeue() unlinks a not-yet-accepted child from the parent
accept queue and release_sock()s it before returning, so the returned
sk has no caller reference and is unlocked.
l2cap_sock_cleanup_listen() walks these children on listening-socket
close. A concurrent HCI disconnect drives hci_rx_work ->
l2cap_conn_del() which runs l2cap_chan_del() + l2cap_sock_kill() and
frees the child sk and its l2cap_chan; cleanup_listen() then uses both:
BUG: KASAN: slab-use-after-free in l2cap_sock_kill
l2cap_sock_kill / l2cap_sock_cleanup_listen / __x64_sys_close
Freed by: l2cap_conn_del -> l2cap_sock_close_cb -> l2cap_sock_kill
This is distinct from the two fixes already in this area: commit
e83f5e24da741 ("Bluetooth: serialize accept_q access") serialises the
accept_q list/poll and takes temporary refs inside bt_accept_dequeue(),
and CVE-2025-39860 serialises the userspace close()/accept() race by
calling cleanup_listen() under lock_sock() in l2cap_sock_release().
Neither covers l2cap_conn_del() running from hci_rx_work, so this UAF
still reproduces on current bluetooth/master.
Take the reference at the source: bt_accept_dequeue() does sock_hold()
while sk is still locked, before release_sock(); callers sock_put().
cleanup_listen() pins the chan with l2cap_chan_hold_unless_zero() under
a brief child sk lock (serialising vs l2cap_sock_teardown_cb()), drops
it before l2cap_chan_lock(), and skips a duplicate l2cap_sock_kill() on
SOCK_DEAD. conn->lock is not taken here: cleanup_listen() runs under
the parent sk lock and that would invert
conn->lock -> chan->lock -> sk_lock (lockdep).
KASAN/SMP: an unprivileged listen/close vs HCI-disconnect race produced
12 use-after-free reports per run before this change; 0, and no lockdep
report, over 1600+ raced iterations after it on bluetooth/master. |
| In the Linux kernel, the following vulnerability has been resolved:
fhandle: fix UAF due to unlocked ->mnt_ns read in may_decode_fh()
may_decode_fh() accesses mount::mnt_ns without holding any locks; that
means the mount can concurrently be unmounted, and the mnt_namespace can
concurrently be freed after an RCU grace period.
This race can happens as follows, assuming that the mount point was
created by open_tree(..., OPEN_TREE_CLONE):
thread 1 thread 2 RCU
__do_sys_open_by_handle_at
do_handle_open
handle_to_path
may_decode_fh
is_mounted
[mount::mnt_ns access]
[mount::mnt_ns access]
__do_sys_close
fput_close_sync
__fput
dissolve_on_fput
umount_tree
class_namespace_excl_destructor
namespace_unlock
free_mnt_ns
mnt_ns_tree_remove
call_rcu(mnt_ns_release_rcu)
mnt_ns_release_rcu
mnt_ns_release
kfree
[mnt_namespace::user_ns access] **UAF**
Fix it by taking rcu_read_lock() around the mount::mnt_ns access, like
in __prepend_path().
Additionally, document the semantics of mount::mnt_ns, and use WRITE_ONCE()
for writers that can race with lockless readers.
This bug is unreachable unless one of the following is set:
- CONFIG_PREEMPTION
- CONFIG_RCU_STRICT_GRACE_PERIOD
because it requires an RCU grace period to happen during a syscall without
an explicit preemption.
This doesn't seem to have interesting security impact; worst-case, it could
leak the result of an integer comparison to userspace (from the level
check in cap_capable()), cause an endless loop, or crash the kernel by
dereferencing an invalid address. |
| OpenClaw before 2026.5.28 contains a race condition in the MS Teams safeFetch DNS rebinding check. When the affected feature is enabled and reachable, a lower-trust caller or configured input path could win a timing window between the DNS validation check and use, allowing actions that should have required a stronger authorization or policy check. Practical impact depends on the operator's configuration and whether lower-trust input can reach that path. |
