| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
audit: fix exe mark UAF in kill_rules()
kill_rules() removes mixed AUDIT_DIR and AUDIT_EXE rules when an audit
tree is pruned. It drops entry->rule.exe before removing the rule from
the RCU-visible filter lists.
After a rule has been installed with AUDIT_ADD_RULE, which requires
CAP_AUDIT_CONTROL, removing the watched directory can race with another
task that is still evaluating the rule. In that case, fsnotify can free
the executable mark before the reader reaches audit_mark_compare(),
causing a use-after-free.
KASAN reports:
BUG: KASAN: slab-use-after-free in audit_mark_compare+0x8d/0xa0
Unlink the published rules from the RCU-visible lists and retain them
on tree->rules for cleanup. If any rule has an executable mark, wait for
a single RCU grace period before removing the marks and scheduling the
entries for freeing. Otherwise, call_rcu() already provides the required
deferred freeing without a synchronous wait. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: x86: Re-pend GET_NESTED_STATE_PAGES if getting said pages fails
Re-pend GET_NESTED_STATE_PAGES before exiting to userspace if getting the
nested pages fails in the KVM_RUN path. If userspace re-runs the vCPU, and
vmcs02 holds valid PFNs from the *previous* run of L2, then KVM could
re-enter L2 with stale, unpinned PFNs mapped into e.g. the vAPIC page.
Note, both SVM and VMX (as of commit 11722439fb20 ("KVM: nVMX: Ensure
KVM_REQ_GET_NESTED_STATE_PAGES is cleared on VM-Exit") ensure the request
is cleared on VM-Exit (including the "forced" case), i.e. there is no risk
of double-mapping due to emulated VMLAUNCH/VMRESUME/VMRUN *and* the request
trying to map the nested pages. |
| In the Linux kernel, the following vulnerability has been resolved:
sched_ext: Wait for SCX_OPSS_DISPATCHING before reenqueueing a task
ebf1ccff79c4 ("sched_ext: Fix ops.dequeue() semantics") moved the final
ops_state store in scx_dispatch_enqueue() after the DSQ unlock so that the
custody update and ops.dequeue() precede it. A task can thus be found on a
DSQ while still SCX_OPSS_DISPATCHING.
The dequeue and core-sched pick paths wait for the state to clear in
ops_dequeue() but the reenqueue paths don't. A reenqueue in that window runs
ops.enqueue() and sets SCX_OPSS_QUEUED before the dispatch has completed.
The dispatcher's final store then overwrites it with SCX_OPSS_NONE and
finish_dispatch() drops every later dispatch of the task.
Wait for SCX_OPSS_DISPATCHING to clear before dequeueing a task for
reenqueue, the same way ops_dequeue() does. |
| In the Linux kernel, the following vulnerability has been resolved:
dm-pcache: bound the logical key offset from persistent memory
cache_key_decode() takes a key's logical off from the cache device and
later indexes req_key_tree->subtrees[] by it in get_subtree(). An off
past the device forms a subtree pointer outside the array, which
rb_insert() writes through during replay.
Reject a key of zero length, or whose off+len (computed in 64 bits)
exceeds the device size, before it is used. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Infer zext_dst based on static register liveness analysis
As reported in the thread [1], the verifier's 32-bit operations zero
extension logic is broken. This logic is responsible for correct
semantics of 32-bit operations on s390 architecture.
According to BPF semantics, operation `w1 += 1` is supposed to zero
extend the upper half of the register `r1`. On s390 the JIT relies on
the verifier emitting explicit zero extension before such operations.
The verifier attempts to minimize the amount of zero extensions
inserted by tracking whether upper halves of the 64-bit registers are
ever used. Previously such tracking worked as follows:
- bpf_reg_state->subreg_def field was set by do_check_insn()
for each operation defining lower but not the upper halves
of the register.
- Whenever an operation reading the whole register was verified,
the verifier checked register's subreg_def and set
bpf_insn_aux_data->zext_dst flag as true via a call to
mark_insn_zext() function.
- After the verification was complete, a special pass
bpf_opt_subreg_zext_lo32_rnd_hi32() extended 32-bit operations
with bpf_insn_aux_data->zext_dst set as true by adding
explicit zero extension.
Note that the logic above relies on bpf_reg_state->subreg_def,
which is a property of a current verifier state.
Before the commit [2] two additional steps happened:
- The verifier tracked upper and lower register halves' liveness as
flags REG_LIVE_READ{32,64} in bpf_reg_state->live.
- The function propagate_liveness() called mark_insn_zext()
in order to transfer the knowledge about which registers have
their upper halves alive (and thus might require zero extension).
The commit [2] removed the two steps described above,
hence making possible a situation like below:
- The register's upper half is set and is used on some verification
path P1 and the register happens not to be marked as precise.
- The checkpoint C is created while processing some instruction
between register initialization and usage.
- On some other verification path P2 the register's upper half is not
initialized and that path ends hitting the checkpoint C.
- In such a case the register's initialization on path P2 would lack
zext_dst mark, making it possible for the program to inject
an arbitrary value in the register's upper half.
This commit replaces subreg_def based logic with computing zext_dst
statically, as a part of the bpf_compute_live_registers() analysis:
- The analysis now tracks usage of upper and lower halves of the
registers separately.
- If some instruction defines a 32-bit subregister, but not the whole
register, *and* the upper half of the register is alive after that
instruction, the instruction is marked as zext_dst.
There is one notable drop in precision: whenever a BPF subprogram is
called, all 64 bits of parameter registers are presumed to be used.
The assumption is that such a drop in precision would not inflict
a noticeable performance penalty.
[1] https://lore.kernel.org/bpf/CAGKGUv=sOuqQtA1Ub-5JXfA4FPosJFYKAQE4B79cK+P1erxqtg@mail.gmail.com/
[2] commit 107e16979905 ("bpf: disable and remove registers chain based liveness") |
| Heap-based buffer overflow in Microsoft Office Excel allows an unauthorized attacker to execute code over a network. |
| Use after free in Windows Hyper-V allows an authorized attacker to elevate privileges locally. |
| Out-of-bounds read in SQL Server allows an authorized attacker to disclose information over a network. |
| A flaw was found in cifs-utils. This flaw allows an unprivileged local user to achieve privilege escalation. When a privileged service mounts a Common Internet File System (CIFS) share using a user-controlled credentials file, input values are not properly sanitized before being assembled into mount options. A local attacker can exploit this vulnerability to inject unauthorized mount options, allowing them to bypass permission checks and execute files with administrative (root) privileges. |
| In the Linux kernel, the following vulnerability has been resolved:
nfsd: fix XDR padding calculation in ff_encode_getdeviceinfo
nfsd4_ff_encode_getdeviceinfo() computes the da_addr_body reservation
as 16 + netid_len + addr_len, but the subsequent xdr_encode_opaque()
calls emit 8 + round_up(netid_len, 4) + round_up(addr_len, 4) bytes.
The mismatch means the declared da_addr_body length exceeds the actual
encoded data by 2-8 bytes on every flexfile GETDEVICEINFO reply,
leaking stale reply-page content to the client and mis-aligning the
subsequent version list decode.
Use xdr_align_size() for each string length to match what
xdr_encode_opaque() actually writes. |
| In the Linux kernel, the following vulnerability has been resolved:
wireguard: noise: reject response consumption after intermediate initiation
Two threads begin processing the identical response message, received
twice. The first thread, A, runs. While it's running, the second one,
B, gets partway through, and during that slow calculation, or even while
blocking on down_write(), A completes and then also a handshake
initiation that's already been queued up runs in thread C, which itself
takes that same down_write(). The handshake initiation creation
succeeds, and sets the state back to waiting-for-response, and calls
up_write(), at which point thread B resumes, because either its finished
its calculations or was finally allowed to acquire down_write(). Thread
B then copies the state back to the peer, and begins a new session,
using that state, which is the same session as the one made in thread A.
Thread A Thread B Thread C
down_read()
sA = handshake->state
memcpy(cA, handshake->crypto)
up_read()
if (sA != 1)
goto fail
slow_crypto(cA)
down_read()
sB = handshake->state
memcpy(cB, handshake->crypto)
up_read()
if (sB != 1)
goto fail
slow_crypto(cB)
down_write()
if (sA != handshake->state)
goto fail
memcpy(handshake->crypto, cA)
handshake->state = 2
up_write()
down_write()
if (handshake->state != 2)
goto fail
derive_session(handshake->crypto)
up_write()
down_write()
slow_crypto(handshake->crypto)
handshake->state = 1
up_write()
down_write()
if (sB != handshake->state)
goto fail
memcpy(handshake->crypto, cB)
handshake->state = 2
up_write()
down_write()
if (handshake->state != 2)
goto fail
derive_session(handshake->crypto)
up_write()
This seems basically impossible to hit in a meaningful way in practice,
but ensure that it absolutely cannot happen by comparing the ephemeral
private key that's on the stack with the latest one that the peer's
handshake state has. |
| In the Linux kernel, the following vulnerability has been resolved:
fsverity: RCU-delay the freeing of struct fsverity_info
__fsverity_get_info() calls rhashtable_lookup_fast(), which just uses
rcu_read_lock() and doesn't directly synchronize with
fsverity_remove_info().
For the same inode this isn't a problem: its fsverity_info is removed
only at inode eviction time or upon failure to enable verity, when the
inode no longer needs its fsverity_info and it will no longer be
accessed via that inode.
However, this is broken and can cause a use-after-free for concurrent
__fsverity_get_info() for *different* inodes. Those rely on following
fsverity_info::rhash_head in the rhashtable under rcu_read_lock() only.
They also use fsverity_info::inode to do the key comparison.
Fix this by RCU-delaying the freeing of 'struct fsverity_info' after
it's been removed from the rhashtable. |
| In the Linux kernel, the following vulnerability has been resolved:
gve: fix XSK buffer leak on error descriptor
When the error bit is set in the RX completion descriptor, the buf_state
and its attached buffer should be freed. In the case of AF_XDP ZC, the
XSK buffer was not freed, leading to a leak.
This issue was caught by LLM. |
| In the Linux kernel, the following vulnerability has been resolved:
net/smc: protect clcsock lifetime in smc_getname
smc_getname() dereferences smc->clcsock without holding
clcsock_release_lock. Link-group termination can release the CLC socket
through smc_close_active_abort() while the SMC socket is still open,
for example after shutdown(SHUT_WR).
A getsockname() caller can load smc->clcsock, then the termination worker
can clear the pointer and call sock_release() before the caller accesses
clcsock->ops or invokes getname(). This causes a use-after-free; if the
worker clears the pointer before the load, it causes a NULL dereference.
The syscall's file reference keeps the SMC socket alive but does not
prevent asynchronous release of its CLC socket.
KASAN reported the following with test-only timing instrumentation:
BUG: KASAN: slab-use-after-free in smc_getname+0x19e/0x1b0
Read of size 8 at addr ffff888109abb4e0 by task poc/103
Call Trace:
smc_getname+0x19e/0x1b0
do_getsockname+0xe5/0x170
__sys_getsockname+0x8c/0x100
Allocated by task 95:
sock_alloc_inode+0x1e/0x280
sock_alloc+0x3d/0x240
__sock_create+0x7e/0x430
smc_create+0x121/0x240
Freed by task 0:
kmem_cache_free+0xcc/0x340
rcu_core+0x50a/0x1850
Last potentially related work creation:
evict+0x446/0x6c0
smc_clcsock_release+0xa8/0xd0
smc_close_active_abort+0x26a/0x3a0
__smc_lgr_terminate.part.0+0x137/0x2e0
Hold clcsock_release_lock across the pointer check and the getname()
callback to serialize with smc_clcsock_release(). Return -EBADF if the
CLC socket has already been released, preserving the existing peer
state check and the callback's return value otherwise. |
| In the Linux kernel, the following vulnerability has been resolved:
net: skbuff: don't leave stale bytes in skb_copy_and_csum_bits()
When skb_copy_and_csum_bits() reaches unreadable frags it returns 0
after copying only the linear part, and the rest of the caller's buffer
is left as it was. The callers copy into a buffer that is about to go
out on the wire: an ICMP error quoting the offending packet, or a
driver's TX bounce buffer in skb_copy_and_csum_dev(). Neither buffer
is zeroed beforehand, so whatever was in memory there gets sent.
Zero the part of the buffer we didn't fill. The checksum usually
won't match the data any more, so the receiver will usually drop the
packet, but either way it no longer carries anything it shouldn't.
Only zero for a positive @len, a negative one from a broken caller must
not turn into a huge memset(). |
| In the Linux kernel, the following vulnerability has been resolved:
riscv: errata: select RISCV_ALTERNATIVE_EARLY for THEAD GHOSTWRITE
errata_probe_ghostwrite() only marks the CPU as vulnerable when it is
called at the RISCV_ALTERNATIVES_EARLY_BOOT stage, and the mitigation
in ghostwrite_enable_mitigation() only takes effect if the vulnerable
state was recorded before the boot alternatives are patched.
The early boot stage is only executed when CONFIG_RISCV_ALTERNATIVE_EARLY
is enabled, which currently is only selected by ERRATA_THEAD_MAE. In a
kernel built with ERRATA_THEAD_GHOSTWRITE=y but ERRATA_THEAD_MAE=n, the
early alternatives never run, so ghostwrite_state is never set to
VULNERABLE: ghostwrite_enable_mitigation() returns false, xtheadvector
remains enabled, and sysfs reports "Not affected" on a vulnerable
T-Head C9xx CPU.
Select RISCV_ALTERNATIVE_EARLY from ERRATA_THEAD_GHOSTWRITE so that the
vulnerability is detected in the early boot stage and the default
mitigation is applied. |
| In the Linux kernel, the following vulnerability has been resolved:
Bluetooth: hci_sync: Fix command skb lifetime during scan setup
During PA sync scan setup, hci_le_set_ext_scan_param_sync() gets the
address used for the PA_LINK lookup from hci_sent_cmd_data(). The
returned pointer borrows storage from sent_cmd or req_skb, without
holding a reference to either skb.
The scan worker runs on req_workqueue while hci_cmd_work() runs on the
separate device workqueue. After the scan worker loads sent_cmd, the
command worker can free and replace it before hci_sent_cmd_data()
dereferences the skb. The command payload can also be freed between
returning from the helper and comparing the address. The connection
lookup's RCU critical section does not protect the command skb.
KASAN reported:
BUG: KASAN: slab-use-after-free in hci_sent_cmd_data+0x27c/0x2f0
Workqueue: hci0 hci_cmd_sync_work
Call Trace:
hci_sent_cmd_data+0x27c/0x2f0
hci_le_set_scan_param_sync+0x43a/0x850
hci_passive_scan_sync+0xb09/0x1460
hci_update_passive_scan_sync+0x47c/0x6e0
hci_le_pa_create_sync+0x200/0xa70
hci_cmd_sync_work+0x13c/0x290
Allocated by task 95:
skb_clone+0x13f/0x340
hci_cmd_work+0x2a9/0x7e0
Freed by task 95:
kmem_cache_free+0xba/0x3a0
hci_cmd_work+0x29c/0x7e0
Use hdev->lock to serialize the address snapshot with sent_cmd
replacement. Protect req_skb replacement and completion cleanup with
the same lock because the helper falls back to that skb. Copy the
address before releasing the lock and use the copy for the existing
RCU-protected connection lookup. Release the mutex before sending or
waiting for commands, preserving the lookup and completion ordering. |
| In the Linux kernel, the following vulnerability has been resolved:
namespace: never expire a locked mount
Locked mounts are special. They protect the underlying files and
directories from being revealed. do_umount() refuses to unmount locked
mounts but shrink_submounts() doesn't.
A shrinkable mount can become locked once the owner of a user namespace
puts it beneath a locked mount with MOVE_MOUNT_BENEATH. The locked
property now moves to the mount at the bottom making it possible to
unmount the top mount.
So umount() of an unlocked ancestor now expires the bottom mount first
and the covered directory is revealed:
move_mount(c -> x/hidden, MOVE_MOUNT_BENEATH) = 0
the cover after the lock moved: umount2(x/hidden) = 0
the holder of the lock: umount2(x/hidden) = EINVAL
the root of the copy, busy: umount2(x) = EBUSY
reads x/hidden/secret: "covered-by-root"
So leave a locked mount alone as it dies together with its parent. A
plain umount() of an unlocked mount with a locked one below it is EBUSY
from now on. It's the same for any other locked child. A lazy umount
still takes the whole tree.
mark_mounts_for_expiry() never sees a locked mount. lock_mnt_tree()
leaves a mount on an expiry list alone and nothing else puts a locked
one on such a list. Add an assert for this. |
| In the Linux kernel, the following vulnerability has been resolved:
amt: send the relay's General Query directly from the receive path
An skb queued in a qdisc can outlive the tunnel it references
through a raw pointer in skb->cb. For example, with igmp_qrv set
to 1 on the relay a tunnel lives for 135s, so a netem delay of
180s on the amt device outlives it; when the tunnel expires and is
freed, the subsequent dequeue triggers a use-after-free in
amt_dev_xmit().
BUG: KASAN: slab-use-after-free in amt_dev_xmit+0x2763/0x2e20
Call Trace:
amt_dev_xmit+0x2763/0x2e20 [drivers/net/amt.c:1262]
dev_hard_start_xmit+0x22f/0x620
sch_direct_xmit+0x12e/0xac0
netem_dequeue+0x333/0xc50
net_tx_action+0x35c/0xa60
amt_send_igmp_gq() and amt_send_mld_gq() are only called from
amt_request_handler(), inside the rcu_read_lock_bh() section of
amt_rcv(). amt_request_handler() already has the tunnel the query is
for: it found or created it inside that section. Queuing the query with
dev_queue_xmit() only leads back into amt_dev_xmit(), which strips the
Ethernet header and calls amt_send_membership_query() for that tunnel.
Make that call directly from the two senders instead, the same way
amt_send_advertisement() transmits from the receive path. The query
never waits in a qdisc, the tunnel is only dereferenced inside the
RCU section that found or created it, and nothing is stored in
skb->cb, so no lookup or refcount is needed. Remove the query branch
of amt_dev_xmit(), amt_skb_cb() and struct amt_skb_cb, which have no
users left.
Behaviour changes:
- The relay's own General Queries no longer pass through the amt
device's egress path: its qdisc, tc egress (clsact/tcx), the
netfilter egress hook and packet taps. They are still visible as
UDP on the underlay.
- A query that is sent successfully is no longer counted as
tx_dropped. The old query branch left through the unlock label,
which counted every sent query as dropped.
- A query that reaches amt_dev_xmit() on a relay from elsewhere,
such as a userspace querier, is now dropped at the IGMP/MLD type
switch. Before, it trusted whatever skb->cb held, and a NULL
tunnel hit the WARN_ON(1). |
| In the Linux kernel, the following vulnerability has been resolved:
netfilter: ipset: mark the rcu locked areas properly
When we bump the uref counter, there's no need to keep
the rcu lock because the referred hash table can't
disappear. Also, from the same reason in mtype_gc we
need the rcu lock and not a spinlock. |
