| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
NFSD: check truncate permission under inode lock
nfsd_setattr() checks whether a size update needs NFSD_MAY_TRUNC
before it takes inode_lock(). The comparison uses the file size sampled
by that unlocked read, but the actual ATTR_SIZE update is applied later
under inode_lock() by notify_change().
This leaves a TOCTOU window for append-only files. If a client sends a
SETATTR that does not shrink the file at the time of the unlocked
sample, a concurrent append can extend the file before nfsd_setattr()
takes inode_lock(). notify_change() then applies a real truncation
without the NFSD_MAY_TRUNC check that rejects IS_APPEND(inode). The VFS
truncate syscall paths perform their own append-only checks before
calling notify_change(), so NFSD must make this decision against the
locked size it is about to change.
Split the write-count acquisition from the truncation permission check.
Keep get_write_access() before the locked setattr work, then recheck
whether the requested size is below i_size_read(inode) after inode_lock()
has been acquired and before notify_change(ATTR_SIZE). This also avoids
the plain unlocked inode->i_size load. |
| In updateInternal of MediaProvider.java, there is a possible expose contents of files due to a race condition. This could lead to local information disclosure with no additional execution privileges needed. User interaction is not needed for exploitation. |
| In the Linux kernel, the following vulnerability has been resolved:
pidfd: hold exec_update_lock around namespace ioctl
The PIDFD_GET_*_NAMESPACE ioctls in pidfd_ioctl() perform a filesystem
credentials ptrace access check before handing out a namespace file
descriptor. The accompanying comment states that the code "mirrors nsfs
behavior", but, unlike the corresponding procfs paths, it does so without
holding the target task's exec_update_lock.
proc_ns_get_link() and proc_ns_readlink() both take exec_update_lock for
reading around the ptrace check and the namespace lookup, so that the
credentials used for the access decision match those of the task when its
namespace is read. Without it, a caller can pass the check against the
target's old credentials and then read the namespace after the target has
execve()'d a setuid binary and committed new credentials -- accessing
namespace information it should have been denied.
Hold exec_update_lock for reading around the ptrace check and the
namespace lookup so that pidfd truly mirrors nsfs behavior, as the comment
already claims. open_namespace() itself runs outside the lock: once a
namespace reference is obtained it carries its own refcount and is opened
with the caller's own credentials, so a concurrent execve() on the target
can no longer affect the outcome. |
| In the Linux kernel, the following vulnerability has been resolved:
mm: memcg: stop reclaim when a limit update is superseded
kernfs serializes file operations only per open file, so separate open
files can update the same memory.high or memory.max file concurrently.
Both handlers store the new limit before synchronous reclaim, but continue
to use the writer's local target in the reclaim loop. If another writer
raises or removes the limit, the first writer can continue reclaiming
toward a stale target.
For memory.max, this can leave the writer looping indefinitely once
reclaim retries are exhausted. The OOM path sees sufficient margin under
the current limit and returns true without killing, while the writer still
compares usage against its stale target and records another OOM event.
Check the current limit at the start of each reclaim iteration and stop if
it no longer matches the writer's target.
Reproducer:
Populate a cgroup with anonymous memory and disable swapping. Lower
memory.max from one open file, then restore it to "max" through another
open file after the new limit becomes visible.
Without the patch, the first writer remains blocked and repeatedly
increments the OOM event counter. With the patch, it returns normally.
This was not motivated by a reported production workload. We found it
through automated randomized testing for our cgroup observability work
and reduced it to the reproducer above. |
| In the Linux kernel, the following vulnerability has been resolved:
nfsd: Reset write verifier when async COPY writeback fails
Async COPY captures nn->writeverf at request time and reports it to
the client via CB_OFFLOAD after the worker kthread completes. When
the post-copy vfs_fsync_range() or filemap_check_wb_err() in
_nfsd_copy_file_range() reports an error, the worker correctly
leaves NFSD4_COPY_F_COMMITTED clear so that CB_OFFLOAD encodes
wr_stable_how as NFS_UNSTABLE, but the server's write verifier is
not rotated.
A client that receives NFS_UNSTABLE in CB_OFFLOAD follows up with
COMMIT to make the copied data durable. With the verifier
unchanged, COMMIT returns the same value the client just received
via CB_OFFLOAD, and the client concludes the copy is durable --
silently dropping the data whose writeback in fact failed. This
violates the UNSTABLE+COMMIT durability contract (RFC 7862 section
15.1, RFC 8881 section 18.32) and matches the bug just fixed in
nfsd_vfs_write() and nfsd_commit().
Rotate nn->writeverf at the writeback-failure site. The async COPY
worker has no svc_rqst, so commit_reset_write_verifier() is not
available here; calling nfsd_reset_write_verifier() directly
mirrors the trace-less reset already used by
nfsd_file_check_write_error() for the same purpose. Filter out
-EAGAIN and -ESTALE, matching commit_reset_write_verifier(), since
neither indicates a durable-storage failure. |
| In the Linux kernel, the following vulnerability has been resolved:
nfsd: sample writeback error cursor before async COPY loop
_nfsd_copy_file_range() samples dst->f_wb_err into "since"
after the copy loop, then uses it to detect writeback errors
via filemap_check_wb_err() once vfs_fsync_range() returns.
Because the nfsd_file cache reuses a single struct file
across requests targeting the same inode, a concurrent
COMMIT or stable WRITE on dst advances dst->f_wb_err to the
current mapping->wb_err via file_check_and_advance_wb_err()
during its own vfs_fsync_range(). If that advancement lands
between the writeback error appearing in mapping->wb_err
and the COPY worker sampling "since", the worker captures
the already-advanced cursor, errseq_check() sees cur ==
since and returns zero, and NFSD4_COPY_F_COMMITTED is set
even though writeback failed. CB_OFFLOAD then encodes
wr_stable_how = FILE_SYNC4, the client treats the copied
data as durable, and the failure becomes silent data loss.
Sample since once at the start of the function. The cursor
then reflects state in effect before this COPY issues any
writes, and filemap_check_wb_err() detects any error that
occurs during the copy regardless of which thread first
observes it. This matches the pattern used by
nfsd_vfs_write() and nfsd4_clone_file_range(). |
| In the Linux kernel, the following vulnerability has been resolved:
fsnotify: Fix stale object mask after concurrent mark updates
When a mark gets a new event bit, fanotify and inotify may avoid
recalculating the object mask if the cached aggregate already contains that
bit. This is racy with a recalculation triggered by a concurrent update to
another mark on the same connector.
The concurrent scan can read the mark before the new bit is added, while
the updater reads the old aggregate before that scan publishes its result.
The updater then skips recalculation and the scan publishes a mask without
the bit, leaving the object mask stale after both updates complete.
This can be reproduced with two fanotify groups watching the same inode:
one thread removes FAN_MODIFY from one existing mark while another thread
adds FAN_MODIFY to the other mark. After both fanotify_mark() calls return,
writes can fail to produce FAN_MODIFY for the group whose mark now contains
the bit. This was reproduced on an unmodified v6.12.95 kernel. The
equivalent inotify interleaving loses IN_MODIFY events.
For normal fanotify additions, recalculate whenever the raw mark mask
changes. The normal mask is not cleared asynchronously, so an unchanged
addition cannot introduce missing interest. Always recalculate ignore-mask
updates because FS_MODIFY handling may clear the ignore mask without taking
mark->lock, making snapshot comparisons unreliable.
Always recalculate after updating an existing inotify watch. Its replace
path temporarily sets mark->mask to zero, so a concurrent scan can observe
zero even when the old and final masks are equal. Assigning the replacement
mask directly would avoid the transient zero, but existing-watch updates
are infrequent, so unconditional recalculation is simpler. |
| In the Linux kernel, the following vulnerability has been resolved:
nouveau/gem: reserve the bo in the info ioctl around the vma lookup
In the non-uvmm path, there could be a race between the info lookup
finding the vma, and the gem close path closing the vma leading
to a use-after-free.
Spotted with the help of Opus 4.6. |
| In the Linux kernel, the following vulnerability has been resolved:
sched/core: Handle pick_task() releasing the rq lock
Core scheduling's pick_next_task() breaks when a ->pick_task()
implementation can release the rq lock. The selection state derived on entry
is only valid while the lock is held continuously. Once a pick can drop the
lock, an interleaving selection can invalidate all of it: the single-CPU
fast path can commit an uncookied pick although the core went cookied during
the release, and forceidle committed by the interleaving selection skews the
restarted pass's accounting.
Fix it by restarting the whole selection when a pick returns RETRY_TASK
after releasing the lock: a single restart point above the state derivation
replaces the per-loop restart labels, so a retry picks up state committed by
interleaving selections and accounts and resets forceidle like a fresh
selection would.
need_sync and fi_before latch across retries. Clock validity can't be
re-derived - there is no program-ordered way to tell whether the own and
core rq clocks are still updated after the lock was released, as other
lockers' pin cycles may or may not have invalidated them. When restarting,
clear core_clock_updated so that the sibling loop re-updates the core rq,
and update the own rq clock if invalidated. |
| In the Linux kernel, the following vulnerability has been resolved:
sched_ext: Replace SCX_RQ_BAL_KEEP with a dispatch verdict return
SCX_RQ_BAL_KEEP tells the pick to keep running the previous task, a leftover
from when balancing and picking were separate operations. An rq-level flag
only works while dispatches and picks pair up one to one, which core
scheduling breaks: selections interleave through dispatch's lock drops and a
pick can consume a stale flag, keeping a task that has since been dequeued.
Fixing core scheduling support requires the decision to travel with the
dispatch that made it. Make scx_dispatch_sched() and balance_one() return an
explicit verdict instead and drop the flag's plumbing from the tools autogen
enum headers.
Also factor the pick-side invocation, its follow-up queueing and the
post-dispatch checks out of do_pick_task_scx() into dispatch_pick(). No
functional changes intended.
v2: Drop the SCX_RQ_BAL_KEEP plumbing from the tools autogen enum headers
as well (Andrea). |
| In the Linux kernel, the following vulnerability has been resolved:
RDMA/ucma: Lock the handler in ucma_set_ib_path()
ucma_set_ib_path() calls ucma_event_handler() straight from the write()
path, without the handler lock that keeps ctx->file stable while a uevent
is queued. The handler re-reads ctx->file for every dereference:
mutex_lock(&ctx->file->mut); /* file A */
list_add_tail(&uevent->list, &ctx->file->event_list); /* file B */
mutex_unlock(&ctx->file->mut); /* file B */
wake_up_interruptible(&ctx->file->poll_wait); /* file B */
A concurrent ucma_migrate_id() reassigns ctx->file while the SET_OPTION
caller sleeps in mutex_lock(), so the list_add_tail() lands on file B's
event_list while only file A's mutex is held, racing every other user of
that list:
BUG: KASAN: slab-use-after-free in __list_add_valid_or_report+0x1aa/0x1c0
Read of size 8 at addr ffff888153c6a418 by task poc_corr/486
Call Trace:
__list_add_valid_or_report+0x1aa/0x1c0
ucma_event_handler+0x1be/0xc00
ucma_set_ib_path+0x45e/0x710
ucma_set_option+0x32e/0x590
ucma_write+0x1f9/0x330
Allocated by task 505:
ucma_write_cm_event+0x1a1/0x660
Freed by task 505:
kfree+0x1da/0x4c0
ucma_get_event+0x5d5/0x7e0
The freed object is a ucma_event that another thread dequeued from file B's
list under file B's mutex. File A's mut is left held on top of that,
wedging its next writer in uninterruptible sleep.
This path needs a bound and address-resolved cm_id, so it requires an RDMA
device to be present.
Take the handler lock around the call. |
| In the Linux kernel, the following vulnerability has been resolved:
RDMA/ucma: Lock the handler in ucma_write_cm_event()
ctx->file may only be changed under the handler lock and the xa_lock, which
is what stops uevents being queued for a ctx while ucma_migrate_id() moves
it to another file. The CM core takes that lock before invoking
ucma_event_handler(), but the write() paths that queue uevents themselves
do not.
ucma_write_cm_event() re-reads ctx->file for each of its four dereferences,
so ucma_migrate_id() can swap it mid-sequence:
mutex_lock(&ctx->file->mut); /* file A */
list_add_tail(&uevent->list, &ctx->file->event_list); /* file B */
mutex_unlock(&ctx->file->mut); /* file B */
wake_up_interruptible(&ctx->file->poll_wait); /* file B */
The window is the mutex_lock() itself: the writer sleeps in it while the
migration reassigns ctx->file. The list_add_tail() then runs on file B's
event_list holding only file A's mutex:
list_add corruption. prev->next should be next (ffff888101320f30),
but was ffff88814a08c418. (prev=ffff88814a075c18).
kernel BUG at lib/list_debug.c:32!
Call Trace:
ucma_write_cm_event+0x36e/0x5e0
and file A's mut is left held forever, wedging its next writer in D state.
The uevent is also stranded on a list ucma_cleanup_ctx_events() will not
walk, so it outlives its context. /dev/infiniband/rdma_cm is 0666 and no
RDMA device is involved, so an unprivileged user reaches all of this.
Take the handler lock, as ucma_cleanup_mc_events() does; ctx->cm_id is
pinned by the ucma_get_ctx() reference. |
| In the Linux kernel, the following vulnerability has been resolved:
ring-buffer: Fix subbuf resize race with ring_buffer_alloc_read_page()
ring_buffer_alloc_read_page() is racy with ring_buffer_subbuf_order_set,
it can allocate a reader page with an outdated order. This isn't a big
issue, the user can still re-allocate a new reader page and try again.
However, what is more problematic is if the value of subbuf_order
changes in the middle of ring_buffer_alloc_read_page(). In that case,
bpage->order might not match the actual allocated memory.
Use bpage->order for the allocation to prevent this race. |
| In the Linux kernel, the following vulnerability has been resolved:
timekeeping: Check the return value of tk_get_aux_ts64 in __do_adjtimex()
If the auxiliary clock is disabled during tk_get_aux_ts64() but is enabled
before tks->clock_valid is checked, then uninitialized stackdata will be
used in the calculations and indirectly leaked to userspace.
The same race window also exists after this change and also for the core
timekeeper. But in these cases the only effect would be incorrect
adjustments and this is userspace's responsibility to avoid this. |
| An authorization bypass vulnerability in LXD due to a timing flaw during configuration merging allows an authenticated attacker to bypass target project restrictions during cross-project instance copies. When copying an instance to a target project, LXD performs restriction checks before configuration merging is complete, creating a time-of-check to time-of-use (TOCTOU) condition. An attacker can exploit this flaw to copy instances with disallowed high-privilege configurations into restricted projects, bypassing security controls. |
| Time-of-check time-of-use (toctou) race condition in Windows USB Audio Class driver (usbaudio.sys) allows an authorized attacker to elevate privileges locally. |
| A Time-of-check Time-of-use (TOCTOU) race condition leading to insecure symlink following in Plesk causes local privilege escalation to root via arbitrary file/directory ownership takeover. |
| Heap-based buffer overflow in Windows Program Compatibility Assistant Service allows an authorized attacker to elevate privileges locally. |
| Time-of-check time-of-use (toctou) race condition in Windows MIDI Service Module allows an authorized attacker to elevate privileges locally. |
| Time-of-check time-of-use (toctou) race condition in Windows Kernel allows an authorized attacker to elevate privileges locally. |
