| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| IBM i 7.6, 7.5, 7.4, and 7.3 could allow a remote authenticated attacker to obtain sensitive information and compromise system integrity due to an XML injection flaw. |
| IBM i 7.6, 7.5, 7.4, and 7.3 could allow a remote authenticated attacker to obtain sensitive information due to improper validation of user-supplied path input. |
| In the Linux kernel, the following vulnerability has been resolved:
coresight: platform: defer connection counter increment until alloc succeeds
coresight_add_out_conn() increments nr_outconns before calling
devm_krealloc_array() and again before devm_kmalloc(). If either
allocation fails, the counter is already bumped while the corresponding
array entry is NULL or uninitialized garbage.
coresight_add_in_conn() has the same problem with nr_inconns and
devm_krealloc_array().
In both cases the probe returns -ENOMEM, which causes
coresight_get_platform_data() to call coresight_release_platform_data()
for cleanup. That function iterates up to nr_outconns (or nr_inconns)
entries and dereferences each pointer unconditionally, hitting the NULL
or garbage entry and panicking instead of failing gracefully.
Fix by moving the counter increments to after all allocations succeed,
so the struct is always consistent on any error path. |
| In the Linux kernel, the following vulnerability has been resolved:
staging: nvec: fix use-after-free in nvec_rx_completed()
In nvec_rx_completed(), when an incomplete RX transfer is detected,
nvec_msg_free() is called to return the message back to the pool by
clearing its 'used' atomic flag. Immediately after this, the code
accesses nvec->rx->data[0] to check the message type.
Since nvec_msg_free() marks the pool slot as available via atomic_set(),
any concurrent or subsequent call to nvec_msg_alloc() could claim that
same slot and overwrite its data[] array. Reading nvec->rx->data[0] after
freeing the message is therefore a use-after-free.
Fix this by saving the message type byte before calling nvec_msg_free(),
then using the saved value for the battery quirk check. |
| In the Linux kernel, the following vulnerability has been resolved:
net: serialize netif_running() check in enqueue_to_backlog()
Syzbot reported a KASAN slab-use-after-free in fib_rules_lookup().
The root cause is a race condition where packets can escape the backlog
flushing during device unregistration (e.g., during netns exit).
Commit e9e4dd3267d0 ("net: do not process device backlog during unregistration")
introduced a lockless netif_running() check in enqueue_to_backlog() to
prevent queuing packets to an unregistering device.
However, this creates a TOCTOU race window.
A lockless transmitter (like veth_xmit) can pass
the check before dev_close() clears IFF_UP. If the transmitter is then
delayed, flush_all_backlogs() can run and finish before the transmitter
grabs the backlog lock and queues the packet. The packet then escapes
the flush and triggers UAF later when processed.
Fix this by moving the netif_running() check inside the backlog lock.
This serializes the check with the flush work (which also grabs the lock).
We then either queue the packet before the flush runs (so it gets flushed),
or check netif_running() after the flush/close completes (so it gets dropped). |
| In the Linux kernel, the following vulnerability has been resolved:
tipc: fix UAF in tipc_l2_send_msg()
Syzbot reported a slab-use-after-free in ipvlan_hard_header() when
called from tipc_l2_send_msg().
The root cause is that tipc_disable_l2_media() calls synchronize_net()
while b->media_ptr is still valid. This allows concurrent RCU readers
to obtain the device pointer after synchronize_net() has finished.
The pointer is cleared later in bearer_disable(), but without any
subsequent synchronization, allowing the device to be freed while
still in use by readers.
Fix this by clearing b->media_ptr in tipc_disable_l2_media() before
calling synchronize_net().
This is safe to do now because the call order in bearer_disable()
was reversed in 0d051bf93c06 ("tipc: make bearer packet filtering generic")
to call tipc_node_delete_links() (which needs the pointer) before
disable_media().
https: //lore.kernel.org/netdev/6a2c1007.428ffe26.258b27.015d.GAE@google.com/T/#u |
| In the Linux kernel, the following vulnerability has been resolved:
kcm: use WRITE_ONCE() when changing lower socket callbacks
kcm_attach() replaces a live lower TCP socket's sk_data_ready and
sk_write_space callbacks with KCM handlers, and kcm_unattach() restores
them later. Those callback-pointer updates are still plain stores even
though the same fields can be read and invoked concurrently on other
CPUs.
If another CPU observes an older callback snapshot after the live field
has already been restored, callback execution can run with a mismatched
target and sk_user_data state, leading to stale or misdirected wakeups.
Use WRITE_ONCE() for the callback replacement and restore operations so
these shared callback fields follow the same visibility contract already
established by the earlier 4022 fixes. |
| In the Linux kernel, the following vulnerability has been resolved:
cxl/region: Fix out-of-bounds access in cxl_cancel_auto_attach()
In cxl_cancel_auto_attach(), it assumes cxled->pos is a valid index for
accessing p->targets[]. However, cxled->pos can be set to negative errno
in cxl_region_sort_targets() if cxl_calc_interleave_pos() fails. This
causes the driver to use a negative index to access p->targets[],
resulting in out-of-bounds access.
Fix it by walking p->targets[] instead of using cxled->pos directly. |
| In the Linux kernel, the following vulnerability has been resolved:
RDMA/mlx5: Fix undefined shift of user RQ WQE size
set_rq_size() computes the RQ WQE size as "1 << rq_wqe_shift" based on
the user-provided rq_wqe_shift, which is only checked to be greater than
32, so shifts of 32 are still accepted. A shift of 31 also overflows a
signed integer, leading to undefined behavior.
Use check_shl_overflow() to compute the RQ WQE size and reject any
invalid values. |
| In the Linux kernel, the following vulnerability has been resolved:
crypto: ccp - Check for page allocation failure correctly in TIO
Sashiko notes:
> if __snp_alloc_firmware_pages() returns NULL under memory pressure, is it
> safe to pass it directly to page_address()?
>
> On architectures without HASHED_PAGE_VIRTUAL, page_address(NULL) might
> compute a deterministic but invalid, non-zero virtual address. The
> subsequent if (tio_status) check would then evaluate to true, and
> sev_tsm_init_locked() would dereference the invalid pointer.
Indeed, page_address(NULL) will return non-NULL garbage here. Fix this by
checking the page allocation itself for NULL, not the resulting virtual
address. |
| In the Linux kernel, the following vulnerability has been resolved:
wifi: ath11k: cancel SSR work items during PCI shutdown
A reboot can crash the kernel if it overlaps with WLAN firmware crash
recovery (SSR). The crash is a NULL pointer dereference in the MHI teardown
path while freeing DMA-backed MHI contexts.
Simplified trace:
dma_free_attrs
mhi_deinit_dev_ctxt [mhi]
ath11k_pci_power_down [ath11k_pci]
ath11k_pci_shutdown [ath11k_pci]
device_shutdown
kernel_restart
On the host side, SSR is driven by the MHI RDDM callback, which queues
reset_work to perform device recovery. reset_work power-cycles the device
by calling ath11k_hif_power_down() followed by ath11k_hif_power_up(). The
power-down phase deinitializes MHI and frees DMA resources.
Shutdown/reboot runs fully asynchronously with this RDDM-driven SSR
recovery flow. As a result, the shutdown path
(ath11k_pci_shutdown() -> ath11k_pci_power_down()) can race with the SSR
recovery sequence.
Fix this by canceling SSR-related work items during PCI shutdown, marking
the device as unregistering, and serializing the RDDM callback path that
checks and queues reset_work. This ensures that no new SSR recovery work
can be queued once teardown has started, and that any in-flight recovery
work is fully synchronized before device power-down, preventing MHI
teardown and DMA resource freeing from running more than once.
Note: This issue only affects PCI/MHI-based devices. AHB-based ath11k
devices do not queue reset_work in normal SSR flows.
Tested-on: WCN6855 hw2.1 PCI WLAN.HSP.1.1-04866.5-QCAHSPSWPL_V1_V2_SILICONZ_IOE-1 |
| In the Linux kernel, the following vulnerability has been resolved:
rxrpc: Fix UAF in rxgk_issue_challenge()
Fix rxgk_issue_challenge() to free the page containing the challenge
content after invoking the tracepoint as the whdr passed to the tracepoint
points into the page just freed. |
| In the Linux kernel, the following vulnerability has been resolved:
rxrpc: Don't move a peeked OOB message onto the pending queue
rxrpc_recvmsg_oob() takes a received oob message off recvmsg_oobq and,
if a response is needed, moves it onto the pending_oobq tree. However,
only the unlink from recvmsg_oobq is guarded by MSG_PEEK; the move onto
pending_oobq always runs.
As a result, reading a challenge with MSG_PEEK leaves the skb on
recvmsg_oobq while also adding it to pending_oobq. Since struct
sk_buff's rbnode shares storage with its next and prev pointers,
rb_insert_color() overwrites the list linkage, and the skb, which holds
a single reference, becomes reachable from both queues at once.
When the socket is closed both queues are drained in turn. While
draining recvmsg_oobq, __skb_unlink() follows the next and prev
pointers that rbnode has overwritten and writes to a bad address. Also,
as the skb holds a single reference but is freed from each queue, both
the skb and the connection reference it holds are released twice. This
leads to memory corruption and to a use-after-free caused by the
connection refcount underflow.
MSG_PEEK does not consume the message from the queue, so only unlink it
from recvmsg_oobq and then move it onto pending_oobq or free it when
the message is actually consumed. |
| In the Linux kernel, the following vulnerability has been resolved:
drm/vc4: Supply the overflow slot size in BPOS, not the whole bin BO size
vc4_overflow_mem_work() points BPOA at a 512KB slot inside the 16MB
binner BO, but writes the size of the whole BO to BPOS. On every binner
out-of-memory event the PTB is therefore authorized to write tile lists
across all the other slots (which may hold the tile state, tile alloc and
overflow memory of in-flight jobs) and, for any slot but the first, past
the end of the binner BO into unrelated CMA memory.
Since CMA pages are recycled into page cache and user allocations, this
is arbitrary memory corruption by GPU DMA. In practice it shows up as GPU
hangs with corrupted control list pointers, userspace heap corruption, a
GPU that stays permanently wedged after the first hang, and occasional
full system crashes, whenever a job overflows the initial binner slot.
The bug dates back to the conversion from a dedicated overflow BO (where
writing the full BO size was correct) to the slotted binner BO. |
| In the Linux kernel, the following vulnerability has been resolved:
sctp: prevent peer transport count overflow
sctp_assoc_add_peer() increments the association's 16-bit transport_count
for every new unique peer. Adding the 65,536th transport wraps the count to
zero.
SCTP sock_diag uses transport_count to reserve the INET_DIAG_PEERS payload,
then copies one sockaddr_storage for every entry in transport_addr_list.
After the wrap, a diagnostic dump reserves an empty payload and writes
8 MiB of peer addresses past the skb tail.
Reject a new unique peer when transport_count has reached U16_MAX. Perform
the check after the existing-peer lookup so a duplicate address continues
to return its existing transport at the limit. |
| In the Linux kernel, the following vulnerability has been resolved:
mm/huge_memory: unlock i_mmap_rwsem before releasing after-split folios
__folio_split() keeps dereferencing the mapping after the split:
shmem_uncharge(mapping->host) and remap_page() while the folios are still
frozen/locked, and i_mmap_unlock_read(mapping) at the very end, after the
after-split folios have been unlocked and freed.
Nothing holds an inode reference across that. The split relies on @folio
-- which the beyond-EOF drop loop never removes, as it starts at
folio_next(folio) -- staying locked and in the page cache to hold off
eviction. But the unlock loop unlocks @folio before i_mmap_unlock_read()
runs. If the caller's @lock_at is a tail beyond EOF, as memory_failure()
passes when splitting a poisoned tail of a shmem THP that reaches past
i_size during truncation, it too is gone from the page cache; so once
@folio is unlocked no locked, in-cache folio pins the inode, and a
concurrent final iput() can evict and RCU-free it before
i_mmap_unlock_read() touches i_mmap_rwsem:
BUG: KASAN: slab-use-after-free in __up_read+0x634/0x790
i_mmap_unlock_read include/linux/fs.h:537 [inline]
__folio_split+0x732/0x1640 mm/huge_memory.c:4100
try_to_split_thp_page+0xab/0x390 mm/memory-failure.c:1675
memory_failure+0x1394/0x26e0 mm/memory-failure.c:2470
Freed by task 4601:
shmem_free_in_core_inode+0x54/0xb0 mm/shmem.c:5177
evict+0x57f/0xac0 fs/inode.c:870
Do every mapping dereference while @folio still pins the inode: drop
i_mmap_rwsem right after remap_page(), before the loop that unlocks and
frees the after-split folios, and clear @mapping so the exit path does not
unlock it again. shmem_uncharge() and remap_page() already run before
that point, so after this nothing past the unlock loop touches the inode
or the mapping.
This is now a rule the split depends on, alongside keeping @folio frozen
until the page cache is updated: no inode or mapping dereference once the
after-split folios start being unlocked. |
| In the Linux kernel, the following vulnerability has been resolved:
wifi: mac80211: fix tid_tx use-after-free on BA session stop
ieee80211_stop_tx_ba_cb() hands tid_tx to kfree_rcu() through
ieee80211_remove_tid_tx(), and then reads tid_tx->ndp after dropping
sta->lock:
ieee80211_remove_tid_tx(sta, tid); /* kfree_rcu(tid_tx, rcu_head) */
...
spin_unlock_bh(&sta->lock);
if (start_txq)
ieee80211_agg_start_txq(sta, tid, false);
if (send_delba)
ieee80211_send_delba(..., tid_tx->ndp);
That read is not covered by an RCU read-side critical section, and it runs
in preemptible process context: both callers hold the wiphy mutex, reaching
it either from the ieee80211_ba_session_work() wiphy work or from
ieee80211_sta_tear_down_BA_sessions() during station teardown.
Softirqs can run in that window too, both from the local_bh_enable() that
ends ieee80211_agg_start_txq() and from any interrupt exit, so the RCU
callback can free tid_tx before the read.
Driving the function from a test module with the grace period forced into
that window, KASAN reports the read, and the free arrives on the ordinary
RCU softirq path:
BUG: KASAN: slab-use-after-free in ieee80211_stop_tx_ba_cb+0x3cd/0x400
Read of size 1 at addr ffff888002b9f52e by task kworker/0:1/10
[...]
Freed by task 57:
__kasan_slab_free+0x47/0x70
__rcu_free_sheaf_prepare+0x70/0x250
rcu_free_sheaf_nobarn+0x18/0x40
rcu_core+0x426/0x1310
handle_softirqs+0x144/0x590
__irq_exit_rcu+0xea/0x150
irq_exit_rcu+0x9/0x20
sysvec_apic_timer_interrupt+0x6b/0x80
asm_sysvec_apic_timer_interrupt+0x1a/0x20
send_delba is only set when tx_stop is set, which happens for
AGG_STOP_LOCAL_REQUEST alone, so this is reached on local teardown -
session idle timeout, PTK rekey, suspend, HW reconfig - and not from a
peer's DELBA.
Read ndp into a local before the session is freed, while sta->lock is still
held. tid_tx->ndp has a single writer, in
ieee80211_tx_ba_session_handle_start(), which cannot run concurrently here:
both paths are serialised by the wiphy mutex, and the session is already
marked HT_AGG_STATE_STOPPING at this point. tid_tx->ndp is also the only
tid_tx dereference left after ieee80211_remove_tid_tx() in this function.
[move/change the comment a bit to be more general not just on ndp,
initialize ndp directly] |
| In the Linux kernel, the following vulnerability has been resolved:
netfilter: nf_conntrack_sip: widen NAT rewrite delta to s32 in sip_help_tcp()
sip_help_tcp() stores the size change of each NAT-rewritten SIP message
in s16 diff and accumulates it in s16 tdiff, but a single message can
grow by more than S16_MAX while the packet stays under the 65535
enlarge_skb() limit: nf_nat_sip() rewrites every matching URI, and a long
Contact list expands the message by tens of kilobytes. diff then wraps,
and "datalen = datalen + diff - msglen" yields a huge unsigned datalen,
so the next iteration's ct_sip_get_header() reads past the linearized skb
tail.
Widen diff, tdiff and the seq_adjust hook to s32. Both are bounded by the
65535 byte packet limit, and the seqadj core is already s32
(nf_ct_seqadj_set() takes s32), so no previously accepted input is
rejected.
BUG: KASAN: use-after-free in ct_sip_get_header (net/netfilter/nf_conntrack_sip.c:464)
Read of size 1 at addr ffff888010800000 by task ksoftirqd/1/25
ct_sip_get_header (net/netfilter/nf_conntrack_sip.c:464)
sip_help_tcp (net/netfilter/nf_conntrack_sip.c:1694)
nf_confirm (net/netfilter/nf_conntrack_proto.c:183)
nf_hook_slow (net/netfilter/core.c:619)
ip6_output (net/ipv6/ip6_output.c:246)
ip6_forward (net/ipv6/ip6_output.c:690)
ipv6_rcv (net/ipv6/ip6_input.c:351)
__netif_receive_skb_one_core (net/core/dev.c:6212)
process_backlog (net/core/dev.c:6676)
__napi_poll (net/core/dev.c:7735)
net_rx_action (net/core/dev.c:7955)
handle_softirqs (kernel/softirq.c:622)
run_ksoftirqd (kernel/softirq.c:1076)
... |
| In the Linux kernel, the following vulnerability has been resolved:
dmaengine: idxd: fix fdev setup failure cleanup in idxd_cdev_open()
The failed_dev_add and failed_dev_name paths drop the file-device
reference while wq->wq_lock is still held. If put_device(fdev) drops the
last reference, idxd_file_dev_release() runs synchronously and tries to
take wq->wq_lock again, deadlocking.
Those paths also fall through into the later ctx cleanup labels even
though idxd_file_dev_release() owns that cleanup and frees ctx. This can
make idxd_xa_pasid_remove(ctx) and kfree(ctx) operate on a freed context.
Move idxd_wq_get() before file-device setup can fail, since the release
callback always calls idxd_wq_put(). Then unlock wq->wq_lock before
put_device(fdev) and return directly from the file-device setup failure
path, leaving ctx cleanup to the release callback. |
| In the Linux kernel, the following vulnerability has been resolved:
mm/slab: prevent unbounded recursion in free path with new kmalloc type
Commit 280ea9c3154b ("mm/slab: avoid allocating slabobj_ext array from
its own slab") avoided recursive allocation of obj_exts from kmalloc
caches of the same size, by bumping the obj_exts array's allocation
size whenever the array size equals the size of the object being
allocated.
However, as reported by Danielle Costantino and Shakeel Butt,
even slabs from kmalloc caches of different sizes can form a cycle
by allocating obj_exts arrays from each other [1]:
What happened: a KMALLOC_NORMAL slab's obj_exts array (used by
allocation profiling / memcg accounting) is itself kmalloc()'d from a
KMALLOC_NORMAL cache, so the "slab holds another slab's obj_exts array"
relation can form cycles. With sizeof(struct slabobj_ext) == 16 and
the host's geometry:
- kmalloc-512 has 64 objects/slab -> array is 64*16 == 1024 bytes,
served from kmalloc-1k;
- kmalloc-1k has 32 objects/slab -> array is 32*16 == 512 bytes,
served from kmalloc-512.
A kmalloc-512 slab and a kmalloc-1k slab therefore hold each other's
obj_exts array. Discarding one frees the other's array, which empties
and discards that slab, which frees the first's array, and so on:
__free_slab() -> free_slab_obj_exts() -> kfree() -> discard_slab() ->
__free_slab() recurses along the cycle until the stack is exhausted.
With memory allocation profiling, this allows unbounded recursion
in the free path and led to a stack overflow on a production host in
the Meta fleet [1]:
BUG: TASK stack guard page was hit
Oops: stack guard page
RIP: 0010:kfree+0x8/0x5d0
Call Trace:
__free_slab+0x66/0xc0
kfree+0x3f0/0x5d0
... ( ~125x __free_slab <-> kfree ) ...
<kernel driver freeing a resource>
do_syscall_64
It is proposed [1] to resolve this issue by always serving the obj_exts
array allocation from kmalloc caches (or large kmalloc) of sizes larger
than the object size. However, as pointed out by Vlastimil Babka [2],
this can waste an excessive amount of memory as slabs from large
kmalloc sizes (e.g. kmalloc-8k) generally need obj_exts arrays much
smaller than the object size.
Therefore, rather than bumping the size, let us take a different
approach; disallow formation of cycles between kmalloc types when
allocating obj_exts arrays. Currently, all obj_exts arrays are served
from normal kmalloc caches. Cycles cannot be created if obj_exts arrays
of normal kmalloc caches are served from a special kmalloc type that can
never have obj_exts arrays.
To achieve this, create a new kmalloc type called KMALLOC_NO_OBJ_EXT.
KMALLOC_NO_OBJ_EXT caches are created with SLAB_NO_OBJ_EXT flag when
either 1) memory allocation profiling is not permanently disabled,
or 2) kmalloc types with a priority higher than KMALLOC_CGROUP are
aliased with KMALLOC_NORMAL.
Sheaf bootstrapping for KMALLOC_NO_OBJ_EXT caches now must be deferred
because allocation of a barn can trigger obj_exts array allocation of
normal kmalloc caches when the KMALLOC_NO_OBJ_EXT cache for that size
is not ready yet. For simplicity, perform bootstrapping of sheaves for
all kmalloc caches later.
Introduce a new slab alloc flag, SLAB_ALLOC_NO_OBJ_EXT, to prevent
allocation of obj_exts arrays, and let kmalloc_slab() override the type
to KMALLOC_NO_OBJ_EXT when specified. Note that kmalloc_type() remains
unchanged because kmalloc_flags() bypasses the kmalloc fastpath.
Do not pass SLAB_ALLOC_NO_RECURSE to kmalloc_flags() in
alloc_slab_obj_exts() and instead use SLAB_ALLOC_NO_OBJ_EXT only when
the objects are allocated from normal kmalloc caches. While this
prevents unbounded recursive allocation of obj_exts, it allows
KMALLOC_NO_OBJ_EXT caches to have sheaves.
Since sheaf allocations specify SLAB_ALLOC_NO_RECURSE that prevents
allocation of both sheaves and obj_exts arrays, the recursion depth
is bounded.
obj_exts arrays for non-
---truncated--- |
