| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| This CVE ID has been rejected or withdrawn by its CVE Numbering Authority. |
| In the Linux kernel, the following vulnerability has been resolved:
keys: Protect the type name while describing a key
keyctl_describe_key() must keep the key type's name alive until
kasprintf() has consumed it. The retained key reference keeps the key
and its description alive, but does not pin the module providing its
type. The name is passed to kasprintf() without holding key->sem.
kvasprintf() uses the saved name pointer in two formatting passes, with
a GFP_KERNEL allocation between them. With AF_RXRPC=m, a task that
passes View permission before type retirement can overlap a privileged
module unload in this order:
1. KEYCTL_DESCRIBE looks up the key, saves key->type->name in the
kasprintf() arguments and completes the first formatting pass.
2. While the formatter is delayed before its second pass,
af_rxrpc_exit() calls unregister_key_type(), which removes the type
from the registry and waits for key_gc_keytype().
3. The collector takes key->sem for writing, retypes the retained key
to key_type_dead and completes retirement. Unregistration and module
exit return, allowing free_module() to release the type and name.
4. The formatter resumes its second pass and reads the saved name from
released module storage, which can fault in string().
Hold key->sem for reading across kasprintf(). The collector then cannot
retype the key or complete retirement until both formatting passes have
finished. If retirement wins the lock first, formatting uses the core
key_type_dead name instead. Release the semaphore immediately after
kasprintf(), including on allocation failure, before copying to
userspace.
Oops report as below:
BUG: unable to handle page fault for address: fffffbfff8054794
#PF: supervisor read access in kernel mode
#PF: error_code(0x0000) - not-present page
RIP: 0010:string+0x23f/0x470
Call Trace:
vsnprintf+0x330/0x1110 |
| In the Linux kernel, the following vulnerability has been resolved:
keys: Check the authorization payload in process keyring searches
An authorization payload and its saved credentials must remain valid
while search_process_keyrings_rcu() searches the requestor's keyrings.
The function validates the authorization key, then separately loads
payload.data[0] and unconditionally reads rka->cred. Validation does
not prevent request_key_auth_revoke() from clearing the payload.
A request-key helper can assume authority and fork a descendant that
inherits the authorization key. If the descendant misses in its own
keyrings, it falls back to the saved requestor credentials. Successful
target instantiation invalidates the authorization key without clearing
the payload. After the helper exits, UMH_WAIT_PROC returns to the
requester, which calls complete_request_key() and revokes that key.
With the descendant paused after validation, the paths can interleave
as follows:
1. The descendant enters search_process_keyrings_rcu() under RCU and
key_validate() succeeds.
2. The helper instantiates the target, invalidates the authorization
key, and exits.
3. The requester resumes from UMH_WAIT_PROC, calls
complete_request_key(), and revokes the authorization key.
request_key_auth_revoke() sets its payload pointer to NULL.
4. The descendant loads NULL and dereferences rka->cred.
This NULL dereference can oops the kernel. RCU delays payload disposal
but does not prevent the pointer from being cleared.
Load the payload with dereference_key_rcu() and require a non-NULL
result before validating the authorization key and searching its saved
credentials. A non-NULL payload and its credentials remain alive for
the caller's RCU read-side section; a cleared payload skips the fallback
search and follows the existing error selection.
KASAN report as below:
Oops: general protection fault, probably for non-canonical address 0xdffffc0000000005: 0000 [#1] SMP KASAN NOPTI
KASAN: null-ptr-deref in range [0x0000000000000028-0x000000000000002f]
RIP: 0010:search_process_keyrings_rcu+0x32b/0x790 |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/waitid: prevent cancel reference wrap
io_waitid packs the cancellation state into bit 31 of ->refs and uses
bits 30:0 as a 31-bit reference counter:
#define IO_WAITID_CANCEL_FLAG BIT(31)
#define IO_WAITID_REF_MASK GENMASK(30, 0)
__io_waitid_cancel() marks a request canceled, then increments the low
bits of ->refs to claim completion ownership. When another path owns the
request, the increment remains even though cancellation returns false.
The request remains discoverable by cancellation, so repeated failed
attempts can wrap the 31-bit reference field to zero. A later
cancellation can claim ownership and complete the request a second time.
Use atomic_fetch_or() when setting the cancel flag and inspect the old
flag value. The first cancel attempt continues to the existing ownership
increment, preserving its wake and issuer handoff semantics. Later
attempts see the cancel flag and return without incrementing the
reference field.
This makes repeated failed cancellation attempts idempotent while
leaving the first-cancel state transitions unchanged. |
| In the Linux kernel, the following vulnerability has been resolved:
coredump: hold RCU while releasing parked threads
coredump_finish() releases the parked threads by clearing ->task in
their core_thread entry and calling wake_up_process() on each of them.
It does that without holding a reference on the task and without being
inside rcu.
But calling wake_up_process(task) without rcu here isn't safe. A parked
thread doesn't need that wakeup to leave. A spurious wakeup or a
preemption after the store is enough for the task to go away.
So if the coredump client is preempted between the store and
wake_up_process() the thread can exit and be freed in the meantime and
try_to_wake_up() takes pi_lock in freed memory.
Hold rcu across the loop. |
| In the Linux kernel, the following vulnerability has been resolved:
ftrace: Initialize modname and offset in print_rec()
print_rec() leaves 'modname' and 'offset' uninitialized and
kallsyms_lookup() only sets them on success. When the lookup fails like
when a module is being loaded or unloaded, 'modname' holds a garbage
value, passes the 'if (modname)' check, and is dereferenced by
seq_printf().
Oops: general protection fault
RIP: 0010:string+0x48/0xe0
Call Trace:
vsnprintf+0x1f5/0x650
seq_printf+0x8e/0xb0
print_rec+0x6c/0xc0
t_show+0x4c/0x440
seq_read_iter+0x119/0x480
This can be triggered by reading available_filter_functions while a
module is being loaded or unloaded.
Initialize both variables so a failed lookup skips the module name. |
| In the Linux kernel, the following vulnerability has been resolved:
keys/trusted_keys: return immediately after TPM unseal failure
trusted_tpm_unseal() proceeds to pcrlock() when the TPM unseal operation
fails. If pcrlock() succeeds, its return value overwrites the unseal error,
causing key instantiation to succeed.
Return immediately when unseal fails to preserve the original error. |
| In the Linux kernel, the following vulnerability has been resolved:
keys: make keyring indexes independent of kernel addresses
KEYCTL_READ returns the serial numbers in a keyring in associative-array
traversal order. The first word of the index is a public, deterministic
hash of the type pointer, domain pointer and description. An unprivileged
user can therefore insert keys with known descriptions, predict their
order for each possible KASLR slide, and recover the kernel image address.
On an x86-64 build with 16 MiB alignment, 45 user keys suffice to
distinguish the candidate slides.
Remove the addresses from the whole index, not just its first hash word.
Assign each key type and domain a nonzero 64-bit ID and hash those IDs and
the description with a once-initialized SipHash key. Use the full IDs in
the later index chunks as well. Truncating or hashing the IDs there would
lose the guarantee that distinct type/domain/description tuples have a
differing index bit.
Allocate IDs lazily since private key types need not be registered. Cache
them in the index key so tree walks and copies of an existing index do not
need to dereference a type descriptor that may have been unregistered.
Keep the 32-bit chunk layout and diff_objects() bit offsets consistent,
and retain the separate root slot for keyrings used by recursive search.
The key comparison, permissions and userspace data format are unchanged;
only the address-dependent in-memory ordering is replaced. |
| In the Linux kernel, the following vulnerability has been resolved:
keys: Avoid the owner account dereference in named keyring lookup
Named keyring lookup must keep the storage containing the owner UID
alive through the namespace mapping check. find_keyring_by_name()
reads keyring->user->uid under keyring_name_lock, but that lock protects
the keyring's name entry and allocation, not its separate key_user.
A named session-keyring join can overlap a privileged KEYCTL_CHOWN on
another CPU. If the keyring holds the last reference to its old dynamic
key_user and the ownership transfer succeeds, the following ordering is
possible:
Named join Chown
find_keyring_by_name() keyctl_chown_key()
read_lock(keyring_name_lock) down_write(key->sem)
load old keyring->user
replace key->user and key->uid
up_write(key->sem)
key_put(key)
key_user_put(old user): free
read old user->uid
read_unlock(keyring_name_lock)
Chown neither takes keyring_name_lock nor key_session_mutex, so it can
free the old account between the pointer load and the UID read. The
lookup then reads freed memory even though the keyring itself is alive.
Use keyring->uid for the mapping check. key_alloc() initializes this
inline owner UID and keyctl_chown_key() updates it on chown.
The existing name lock keeps its containing keyring allocated throughout
the check, so lookup no longer depends on the account's lifetime. This
also matches the owner UID used by the permission check.
KASAN report as below:
BUG: KASAN: slab-use-after-free in find_keyring_by_name+0x577/0x5d0
Read of size 4 at addr ffff8881128ae2ec by task keycase/500 |
| IBM InfoSphere Information Server 11.7 vulnerable to SQL injection. A remote attacker could send specially crafted SQL statements, which could allow the attacker to view, add, modify, or delete information in the back-end database. |
| In the Linux kernel, the following vulnerability has been resolved:
dma-buf: heaps: Fix off-by-one in CMA heap fault handler
Until VM_DONTEXPAND was added in commit 1c1914d6e8c6 ("dma-buf: heaps:
Don't track CMA dma-buf pages under RssFile") it was possible to obtain
a mapping larger than the buffer size via mremap and bypass the overflow
check in dma_buf_mmap_internal. When using such a mapping to attempt to
fault past the end of the buffer, the CMA heap fault handler also checks
the fault offset against the buffer size, but gets the boundary wrong by
1. Fix the boundary check so that we don't read off the end of the pages
array and insert an arbitrary page in the mapping. |
| Use after free in Canvas in Google Chrome on Linux, ChromeOS prior to 147.0.7727.138 allowed a remote attacker to execute arbitrary code inside a sandbox via a crafted HTML page. (Chromium security severity: Critical) |
| In the Linux kernel, the following vulnerability has been resolved:
af_unix: Fix garbage collector racing against connect()
Garbage collector does not take into account the risk of embryo getting
enqueued during the garbage collection. If such embryo has a peer that
carries SCM_RIGHTS, two consecutive passes of scan_children() may see a
different set of children. Leading to an incorrectly elevated inflight
count, and then a dangling pointer within the gc_inflight_list.
sockets are AF_UNIX/SOCK_STREAM
S is an unconnected socket
L is a listening in-flight socket bound to addr, not in fdtable
V's fd will be passed via sendmsg(), gets inflight count bumped
connect(S, addr) sendmsg(S, [V]); close(V) __unix_gc()
---------------- ------------------------- -----------
NS = unix_create1()
skb1 = sock_wmalloc(NS)
L = unix_find_other(addr)
unix_state_lock(L)
unix_peer(S) = NS
// V count=1 inflight=0
NS = unix_peer(S)
skb2 = sock_alloc()
skb_queue_tail(NS, skb2[V])
// V became in-flight
// V count=2 inflight=1
close(V)
// V count=1 inflight=1
// GC candidate condition met
for u in gc_inflight_list:
if (total_refs == inflight_refs)
add u to gc_candidates
// gc_candidates={L, V}
for u in gc_candidates:
scan_children(u, dec_inflight)
// embryo (skb1) was not
// reachable from L yet, so V's
// inflight remains unchanged
__skb_queue_tail(L, skb1)
unix_state_unlock(L)
for u in gc_candidates:
if (u.inflight)
scan_children(u, inc_inflight_move_tail)
// V count=1 inflight=2 (!)
If there is a GC-candidate listening socket, lock/unlock its state. This
makes GC wait until the end of any ongoing connect() to that socket. After
flipping the lock, a possibly SCM-laden embryo is already enqueued. And if
there is another embryo coming, it can not possibly carry SCM_RIGHTS. At
this point, unix_inflight() can not happen because unix_gc_lock is already
taken. Inflight graph remains unaffected. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring: don't flood a cancelling task with task_work
io_ring_exit_work() asks each task which has used the ring to remove its
tctx node, by task_work, and repeats that until the list is empty.
io_tctx_exit_cb() doesn't touch the node while the task is in
io_uring_cancel_generic(), which removes all the nodes at its end. So the
work item queues the callback again as soon as it has completed, for
nothing.
That is more than a waste of time. task_work_run() only returns when it
finds the list empty. If the rings are released from inside the
cancellation, because requests of other rings held the last references to
their files, up to 64 work items keep sending callbacks to the one task
which is running them. It never finds the list empty, never gets back to
the cancellation, the nodes stay, and so it goes on: a task with 60 pairs
of rings which poll each other ran 280000 callbacks in 5 seconds, 20
seconds after exit_group(). It ends when all the senders happen to be late
at the same time. Until then the task keeps a CPU and the work items busy.
An exiting task can't be killed meanwhile. execve() cancels the requests in
the same way, so a live task can get stuck there as well, for minutes. If
it takes more than five, the work items run into the
WARN_ON_ONCE(time_after(jiffies, timeout)) of io_ring_exit_work(). No
privilege is needed for any of this.
If the callback has met a task in cancellation, wait a little before it is
sent again. Nobody has to wake the work item up for that: either the node
is gone by then, or it is the turn of the next round. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing: Unregister field variable histograms when the command fails
An action that names a field on another event needs that field as a
variable, so create_field_var_hist() registers a second hist trigger on
that event to provide it, and records it in
target_hist_data->field_var_hists[].
If a later part of the same command fails, create_actions() returns an
error and event_hist_trigger_parse() jumps to out_free. The write fails
with -EINVAL and sched_switch gets no trigger, but:
# echo 'hist:keys=pid:ts0=common_timestamp.usecs' >> \
events/sched/sched_waking/trigger
# echo 'my_synth u64 lat; int a; int b' >> synthetic_events
# echo 'hist:keys=next_pid:wakeup_lat=common_timestamp.usecs-$ts0:\
onmatch(sched.sched_waking).my_synth($wakeup_lat,prio,nosuchfld)'\
>> events/sched/sched_switch/trigger
# cat events/sched/sched_waking/trigger
hist:keys=pid:vals=hitcount:ts0=common_timestamp.usecs:sort=hitcount:size=2048:clock=global [active]
hist:keys=pid:vals=hitcount:synthetic_prio=prio:sort=hitcount:size=2048 [active]
The second histogram is the one created for 'prio'. It is still active and
still counting.
out_free reaches destroy_field_var_hists(), which frees the
field_var_hists[] entries, but never unregister_field_var_hists(), which is
what removes the triggers and runs only from event_hist_trigger_free(). The
entry is freed and the trigger is left with nothing tracking it.
Unregister the field variable histograms in out_free as well, and record
each one as soon as its trigger is registered rather than after the lookup
that follows it, so no exit can leave a registered trigger untracked. The
entry then belongs to field_var_hists[] and is released by
destroy_field_var_hists(), so freeing it on that exit would be a double
free. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/wait: don't spin on a stuck overflow list
As long as IO_CHECK_CQ_OVERFLOW_BIT is set, io_cqring_wait_schedule()
returns right away, before it gets to look at signals or arms the timeout,
for io_cqring_wait() to flush the overflow list. If that doesn't empty the
list, and the CQ ring doesn't hold the number of events the caller wants,
the loop starts over, with the same result. The task spins in the kernel,
and can't be killed.
The assumption is that this can't happen: if CQEs can't be flushed, the
ring is full, and min_events is capped to the size of the ring. But:
- With IORING_SETUP_CQE_MIXED a 32b CQE needs two contiguous entries. It
is stuck on the list if a single one is free, and then the ring holds
less than cq_entries events. A single threaded application that asks
for as many events as the ring has entries never gets out of
io_uring_enter().
- Whether the ring is full is up to the CQ head, which the application
can write to. While one thread waits for events on an empty ring,
another one can move the head back so that the ring looks full and
have a CQE overflow. The waiter can't be stopped any more, not even by
SIGKILL.
Only the application can make room in the CQ ring. Return to it if the
list could not be flushed. |
| In the Linux kernel, the following vulnerability has been resolved:
keys: Protect the read method check against type removal
keyctl_read_key() checks key->type->read without holding key->sem.
A reference keeps the key alive, but does not pin the module containing
its type. A surviving RxRPC key can therefore be read while AF_RXRPC
is being unloaded:
KEYCTL_READ task Module unload and key GC
---------------- ------------------------
lookup_user_key()
permission / possession checks
load key->type
af_rxrpc_exit()
unregister_key_type()
GC: down_write(key->sem)
GC: retype to key_type_dead
GC: up_write(key->sem)
unregister_key_type() returns
free_module()
read saved type's read member
The final access can read freed module storage. The read invocation in
__keyctl_read_key() is already protected by key->sem, but the earlier
method check is outside that critical section.
Move the method check into __keyctl_read_key(), after key_validate() and
under its existing read hold on key->sem. The check and invocation then
use the same protected type for length queries and every read retry.
The GC cannot retype the key until the semaphore is released, and type
unregistration waits for that GC pass before the module can be freed.
Return -EOPNOTSUPP when validation succeeds and the type has no read
method. A temporary buffer may now be allocated before this result;
allocation and copying to userspace remain outside the semaphore. |
| In the Linux kernel, the following vulnerability has been resolved:
asymmetric_keys: reject trust keys without IDs
The PKCS#8 parser deliberately leaves the asym_key_ids payload slot
empty. key_or_keyring_common() unconditionally dereferences that slot
when a PKCS#8 key is supplied to a key_or_keyring restriction. Both the
plain and :chain forms reach this branch, allowing an unprivileged caller
to trigger a NULL pointer dereference through KEYCTL_RESTRICT_KEYRING
followed by add_key().
Oops: general protection fault, probably for non-canonical address 0xdffffc0000000000
KASAN: null-ptr-deref in range [0x0000000000000000-0x0000000000000007]
RIP: 0010:key_or_keyring_common (crypto/asymmetric_keys/restrict.c:205 crypto/asymmetric_keys/restrict.c:279)
Call Trace:
<TASK>
__key_create_or_update (security/keys/key.c:884)
key_create_or_update (security/keys/key.c:1021)
__do_sys_add_key (security/keys/keyctl.c:134)
do_syscall_64 (arch/x86/entry/common.c:52 arch/x86/entry/common.c:83)
entry_SYSCALL_64_after_hwframe (arch/x86/entry/entry_64.S:130)
</TASK>
Kernel panic - not syncing: Fatal exception
Reject an asymmetric restriction key without key IDs. A PKCS#8 private
key cannot identify the signer of a candidate certificate and is not a
valid trust anchor. |
| In the Linux kernel, the following vulnerability has been resolved:
ntfs: fix resident conversion in ntfs_new_attr_flags
When setting sparse/compressed flags on a resident attribute, the
function skipped the resident-to-non-resident conversion and
terminated. |
| In the Linux kernel, the following vulnerability has been resolved:
mm/damon/core: error damos_commit_quota_goal() for zero target_value
Patch series "mm/damon: move zero damos quota target_value handling to the
core layer".
Having zero DAMOS quota target value can cause division by zero. DAMON
API callers are checking the target value parameters to avoid that. It is
easy to make mistakes in some of the multiple API callers. Move that to
the core layer.
Patch 1 adds the corner case handling into the core layer DAMON parameters
validation logic. Patches 2 and 3 remove no more needed DAMON API callers
side handling of the corner case in DAMON_LRU_SORT and DMON_SAMPLIE_MTIER,
respectively.
This patch (of 3):
If a DAMOS scheme has a damos_quota_goal of zero target_value,
damos_quota_goal() could trigger division-by-zero error. Hence each DAMON
API callers should do the zero target_value validation. It is easy to
make mistakes. Actually such bugs in DAMON_LRU_SORT and
DAMON_SAMPLE_MTIER were found and fixed [1].
It is better to handle the corner case only once in the core layer,
instead of multiple places in all DAMON API callers. One straightforward
option is using an alternative denominator for the corner case in the
damos_quota_goal(). However, the zero target_value is meaningless. In
this case, the quota goal is always evaluated as achieved or
over-achieved. The quota will only keep being reduced.
Simply avoid using zero target_value by adding a check in the core layer
DAMOS quota goal parameters validation/commit path,
damos_commit_quota_goal(). Update it to return an error in the case.
Also update its caller to propagate the error. |
