| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
wifi: mac80211: shut down RX BA session timer on teardown
Tearing down an RX BA session can leave the session timer queued on a
freed tid_rx. Later timer-wheel processing and the timer callback itself
then run against freed memory.
An RX BAR handler can dereference tid_rx under RCU before teardown
unpublishes it and stay in its read-side critical section while teardown
deletes the session timer. The stale handler can then rearm the timer
after timer_delete_sync() returns and before the RCU callback frees
tid_rx.
Use timer_shutdown_sync() so rearm attempts become ineffective once
teardown starts. Keep reorder_timer unchanged: teardown marks the session
removed under reorder_lock and its release path checks that state before
rearming. session_timer has no equivalent guard because
ieee80211_rx_h_ctrl() rearms it before taking reorder_lock. |
| In the Linux kernel, the following vulnerability has been resolved:
wifi: mac80211: keep paired old chanctx alive
An in-place replacement keeps reciprocal replace_ctx pointers between
the old WILL_BE_REPLACED chanctx and the new REPLACES_OTHER chanctx.
However, the early free paths only consider assigned and reserved link
users when deciding whether to free the old chanctx.
That lets the old chanctx be freed too early while the replacement
partner still points back to it, either when the last assigned link is
released or when that link successfully reassigns to another existing
context. Later unreserve and switch-finalization paths can then follow a
stale replace_ctx pointer.
Keep a paired old chanctx alive until the replacement pair has been torn
down. Apply the same check to both __ieee80211_link_release_channel()
and ieee80211_link_use_reserved_reassign(). If the replacement is later
abandoned, ieee80211_link_unreserve_chanctx() already tears down the
pairing before freeing the old chanctx once no users remain. |
| In the Linux kernel, the following vulnerability has been resolved:
wifi: mac80211: count only matching reservations in reserved switch
ieee80211_vif_use_reserved_switch() validates a replacement by counting
assigned links on ctx->replace_ctx and treating reservations too
broadly. A link can still be assigned to the old channel context while
holding a reservation that belongs to a different replacement
operation.
That can make the current replacement appear complete too early. If
another link finalizes first, the function can free the
WILL_BE_REPLACED chanctx while the unmatched link and driver still use
it, leading to a use-after-free in later beacon generation.
Require an in-place reservation to point at the current replacement
context, and require the reserved and assigned contexts to be the
matching replacement pair, before counting or moving the link. |
| In the Linux kernel, the following vulnerability has been resolved:
wifi: cw1200: fix TX padding OOB read leaking heap data to device
The TX path rounds the transfer size up to the bus alignment
(align_size(), up to 511 bytes for SDIO, 1 byte for SPI) but sends
the padded length straight from the frame buffer, which has no room
reserved for the padding. cw1200_data_write() therefore reads past
the end of the frame and the bus transfer carries those bytes to the
device. A bogus device can collect up to 511 bytes of kernel heap
data per TX frame.
The RX path does not have this problem: it allocates the skb with
the aligned length (alloc_len) before reading.
Compute the exact padding with the bus align_size() callback, make
sure queued TX frames have tailroom for it, and zero the padding
bytes before the transfer so nothing past the frame is exposed.
This is not expected to change driver behavior on most cases: frame
contents and transfer sizes are unchanged, only the content of the
padding bytes (previously undefined stale memory, now zero), which
the device never reads since the WSM header carries the real frame
length. The skb expansion can only fail under memory exhaustion,
in which case the frame is dropped like on the existing error paths
of cw1200_tx().
Discovered by Atuin - Automated Vulnerability Discovery Engine. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: arm64: vgic: Take a refcount on IRQs referenced by last_lr_irq
Referencing the last interrupt inserted in an LR is rather fragile, as
this interrupt can vanish if a concurrently unmapped LPI.
Solve this by bumping up the refcount on the interrupt when populating
last_lr_irq, and drop it at vgic_prune_ap_list() time, when LPIs are
being reclaimed. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: arm64: Apply the fine-grained UNDEFs without FEAT_FGT
FGUs don't actually require hardware support, so don't gate fgt trap
config insertion on FGT. This is the only action required to allow FGUs
always, as kvm_calculate_traps() already computes the FGU bits with or
without FEAT_FGT.
This also fixes a spurious WARN when a TLBI OS is run on a guest in vEL1
without FEAT_TLBIOS on non FEAT_FGT hardware. In this configuration the
course-grained HCR_EL2.TTLBOS trap reaches the handler handle_tlbi_el1(),
which expects an EL1 TLBI only from vEL2. With FGUs enabled the UNDEF
will be injected without needing the specific handler.
[oliver: Take Wei-Lin's suggested changelog] |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: SVM: Preserve TLB control (i.e. pending TLB flush) on failed VMRUN
Don't reset the VMCB's TLB control back to "do nothing" on a failed VMRUN,
as empirical testing shows that the CPU performs the requested TLB flush if
and only if VMRUN is successful, i.e. clearing TLB control on a failed
VMRUN effectively drops a TLB flush.
Explicitly track the need to flush all ASIDs on a per-CPU basis, as the
ASID reuse condition is tied to the pCPU, not to the vCPU. As a bonus,
this also obviates the need to avoid clobbering FLUSH_ALL_ASID with
TLB_CONTROL_FLUSH_ASID, e.g. in svm_flush_tlb_asid().
Deliberately don't bother saving/restoring the "old" tlb_ctl on failure,
in quotes because it's not exactly the old tlb_ctl, it's the tlb_ctl from
after pre_svm_run(), but before updating tlb_ctl for flush_all_asids. If
VMRUN fails and TLB_CONTROL_FLUSH_ALL_ASID is forced, then the next
successful run of the VMCB *may* unnecessarily flush all ASIDs, which
strictly speaking could result in noisy neighbor issues. However, the
fact that new_asid() is already guest-triggerable, because of KVM's flawed
behavior of clearing the ASID on emulated INIT, means that a guest can
already trigger a flush of all ASIDs at roughly the same rate. And once
KVM stops clobbering the ASID on emulated INIT, *or* assigns a static ASID
to each vCPU, this flaw goes away. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: SVM: Use the active VMCB's MSR bitmap when checking if MSR is intercepted
Use the MSR permission bitmap of the active VMCB instead of assuming that
KVM is always using vmcb02's bitmap when L2 is active, as KVM uses msrpm02
if and only if L1 wants to intercept MSR accesses, i.e. if and only if KVM
needs to merge msprm01 with msrpm12.
Don't bother tracking the virtual address of the bitmap that's being used,
as __va() is cheap on x86, and caching the virtual address would introduce
yet another source of potentially stale information. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: x86/mmu: Bail from shadow walks if the root is invalid or a dummy
When walking shadow page tables, immediately terminate the walk if the root
is a "dummy" root, i.e. a root whose top-level page table is backed by the
zero page, but otherwise doesn't exist. If memslot creation races with a
stage-2 page fault (EPT violation or #NPF) from L2, then if the stars align,
KVM will attempt to walk shadow page tables using the zero page and hit a
NULL pointer deref.
BUG: kernel NULL pointer dereference, address: 0000000000000021
#PF: supervisor read access in kernel mode
#PF: error_code(0x0000) - not-present page
PGD 0 P4D 0
Oops: Oops: 0000 [#1] SMP
CPU: 30 UID: 1000 PID: 941 Comm: qemu Not tainted 7.2.0-rc2-1b731e5ded48-next-vm #1741 PREEMPTLAZY
Hardware name: QEMU Standard PC (Q35 + ICH9, 2009), BIOS 0.0.0 02/06/2015
RIP: 0010:__kvm_mmu_invalidate_addr+0xea/0x210 [kvm]
Call Trace:
<TASK>
kvm_mmu_invalidate_addr+0x92/0xe0 [kvm]
__kvm_inject_emulated_page_fault+0x67/0x80 [kvm]
ept_page_fault+0x160/0x850 [kvm]
kvm_mmu_do_page_fault+0x102/0x1f0 [kvm]
kvm_mmu_page_fault+0x8e/0x6b0 [kvm]
vmx_handle_exit+0x163/0x640 [kvm_intel]
kvm_arch_vcpu_ioctl_run+0x960/0x2120 [kvm]
kvm_vcpu_ioctl+0x2c7/0x970 [kvm]
__x64_sys_ioctl+0x90/0xd0
do_syscall_64+0x67/0x5f0
entry_SYSCALL_64_after_hwframe+0x4b/0x53
RIP: 0033:0x7f606d0b53bb
Opportunistically harden the shadow walks against fully invalid roots, but
WARN, as all callers are expected/required to pre-check for a valid root.
Don't WARN in the dummy root case as the whole point of using a dummy root
is to provide a root that's valid enough to enter the guest, i.e. it should
Just Work for all flows except those that *need* to know about dummy roots.
Alternatively, KVM could allocate a dedicated page and associated shadow
page structure for the dummy root, which is very tempting as it such an
approach should be more resilient against unexpected behavior. But that
would be a much larger and thus riskier change than simply terminating
walks of dummy roots. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: SEV: Do cache maintenance on the source VM *before* clearing SEV state
Explicitly flush caches after intra-host migration before clearing "SEV
active" on the source VM, as doing cache maintenance afterwards creates a
tiny window where memory reclaim could return memory to the host without
performing a cache flush, e.g. as pointed out by Sashiko:
CPU1 in sev_migrate_from():
src->active = false;
CPU2 running concurrent unmap:
Since active is false, the automatic cache flush in
sev_guest_memory_reclaimed is skipped.
The host frees and reallocates the page.
CPU1 in sev_migrate_from():
sev_writeback_caches(src_kvm);
Executes a hardware cache flush (wbnoinvd), which writes the guest's old
dirty ciphertext over the new page owner's data. |
| In the Linux kernel, the following vulnerability has been resolved:
KVM: nVMX: Force MSR bitmap refresh if runtime eVMCS controls are modified
Force a refresh of the vmcs02 MSR bitmap during nested VM-Enter if the
runtime eVMCS controls (pin, primary, secondary, etc.) are being updated.
If L1 isn't intercepting TPR writes, runs L2 with TPR virtualization, and
then runs the same L2 with TPR virtualization disabled, KVM will fail to
refresh msr_bitmap02 and leave TPR in passthrough mode even though TPR
virtualization is disabled. I.e. failure to refresh the bitmap lets L2 (or
L1 by proxy) read and write L0's TPR. |
| In the Linux kernel, the following vulnerability has been resolved:
mtd: rawnand: cadence: Initialize IRQ state before requesting IRQ
The Cadence NAND interrupt handler uses both the IRQ lock and completion
object. Registering the IRQ before initializing them leaves a window in
which a pending interrupt can access uninitialized synchronization state.
Initialize them before registering the handler. |
| In the Linux kernel, the following vulnerability has been resolved:
nvmet: pci-epf: reject too-short SGL segments
The host controls an SGL segment's length. A nonzero length smaller than
one SGL descriptor makes nr_descs zero, so
nvmet_pci_epf_get_sgl_segment() reads the type byte of sgls[-1] outside
the allocated buffer. If that byte contains a segment descriptor type,
the function also copies the preceding 16 bytes and returns a negative
descriptor count.
The read was reproduced in 3/3 KASAN boots with only the host transfer
mocked by a KUnit test:
BUG: KASAN: slab-out-of-bounds in nvmet_pci_epf_get_sgl_segment
Read of size 1 by task kunit_try_catch
Call Trace:
nvmet_pci_epf_get_sgl_segment
nvmet_pci_epf_short_sgl_test
Kernel panic - not syncing: KASAN: panic_on_warn set ...
Reject segments that cannot hold one descriptor before allocating or
transferring their contents. |
| In the Linux kernel, the following vulnerability has been resolved:
nvmet-auth: fix out-of-bounds write in nvmet_auth_challenge()
nvmet_auth_challenge() receives its output buffer as a void pointer.
sizeof(*d) therefore evaluates to one with GCC and undercounts the fixed
challenge header by 15 bytes. A short AUTH_RECEIVE buffer can pass the
check before the challenge is copied past the end of its allocation.
Use the typed challenge pointer when calculating the response size.
BUG: KASAN: slab-out-of-bounds in nvmet_execute_auth_receive (drivers/nvme/target/fabrics-cmd-auth.c:442 drivers/nvme/target/fabrics-cmd-auth.c:569)
Write of size 32 by task kworker/1:1H/64
Workqueue: nvmet_tcp_wq nvmet_tcp_io_work
Call Trace:
nvmet_execute_auth_receive (drivers/nvme/target/fabrics-cmd-auth.c:442 drivers/nvme/target/fabrics-cmd-auth.c:569)
nvmet_tcp_try_recv_pdu (drivers/nvme/target/tcp.c:1119 drivers/nvme/target/tcp.c:1243)
nvmet_tcp_io_work (drivers/nvme/target/tcp.c:1350 drivers/nvme/target/tcp.c:1382 drivers/nvme/target/tcp.c:1445)
process_one_work (kernel/workqueue.c:3322)
worker_thread (kernel/workqueue.c:3405 kernel/workqueue.c:3486)
kthread (kernel/kthread.c:436)
ret_from_fork (arch/x86/kernel/process.c:158)
ret_from_fork_asm (arch/x86/entry/entry_64.S:245)
The buggy address is located 16 bytes inside of
allocated 33-byte region [ffff888000554e80, ffff888000554ea1)
Kernel panic - not syncing: KASAN: panic_on_warn set ... |
| Out-of-bounds read in DNS Server allows an authorized attacker to disclose information over a network. |
| Out-of-bounds read in DNS Server allows an authorized attacker to disclose information over a network. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf, arm64: set up the frame pointer for the exception callback
A program acting as exception boundary saves all callee-saved registers,
so build_prologue() takes the exception_cb path and never calls
push_callee_regs(). That is the only place find_used_callee_regs() runs,
and with it the only place ctx->fp_used is set, so the callback prologue
does not emit the
mov x25, sp
that points BPF_REG_FP at the frame the callback runs on. x25 keeps
whatever it held when bpf_throw() was called. If the throw came from a
subprogram that uses its own BPF stack, that is the subprogram's frame
pointer, and since the subprogram never returns it never restores x25
either.
Stack accesses through BPF_REG_FP are rewritten to be stack pointer
relative, so those still land in the callback's own frame. Materializing
the register does not: a callback that passes the address of a local
variable to a helper hands over an address in the dead subprogram's
frame. That address is below the callback's stack pointer by then, and
the helper's own call chain covers it, so the helper can write over its
own return address. 0x1234 below is the value the helper was asked to
store:
pc : 0x1234
lr : 0x1234
Call trace:
0x1234 (P)
bpf_test_run+0x188/0x3e0
bpf_prog_test_run_skb+0x47c/0x998
__sys_bpf+0xbdc/0xdd8
Kernel panic - not syncing: Oops: Fatal exception in interrupt
Set ctx->fp_used on the exception callback path so that the existing code
further down sets x25 from the stack pointer. The epilogue restores it
from the main program's save area along with the other callee-saved
registers, as it already does. x86 sets the frame pointer for the
callback from the argument it is passed, and powerpc computes it from
the stack pointer. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Fix UAF due to concurrent consumption of ttrace lists in alloc_bulk
Syzkaller repeatedly triggered UAF splats related to nodes in
waiting_for_gp_ttrace within the bpf memalloc:
BUG: KASAN: slab-use-after-free in llist_del_first+0x85/0x110 lib/llist.c:61
Read of size 8 at addr ffff8881572cd080 by task syz.4.470/5112
...
llist_del_first+0x85/0x110 lib/llist.c:61
alloc_bulk+0x193/0x460 kernel/bpf/memalloc.c:229
bpf_mem_refill+0x386/0x560 kernel/bpf/memalloc.c:436
Freed by task 14:
...
__free_rcu kernel/bpf/memalloc.c:281 [inline]
__free_rcu_tasks_trace+0x48/0xd0 kernel/bpf/memalloc.c:291
rcu_tasks_invoke_cbs+0x1ec/0x3e0 kernel/rcu/tasks.h:571
rcu_tasks_one_gp+0x13d/0x220 kernel/rcu/tasks.h:621
rcu_tasks_kthread+0xf3/0x120 kernel/rcu/tasks.h:651
The reason is that the UAF occurs after the RCU Tasks Trace GP expires:
when the __free_rcu() callback runs, there is no synchronization
protecting llist_del_all() against concurrent alloc_bulk() operating on
waiting_for_gp_ttrace, leading to the race condition below:
CPU0 CPU1
__free_rcu (RCU Tasks Trace callback)
alloc_bulk
llist_del_first(&c->waiting_for_gp_ttrace)
entry = smp_load_acquire(&head->first);
do {
if (entry == NULL)
return NULL;
free_all(llist_del_all(&c->waiting_for_gp_ttrace))
llist_for_each_safe(pos, t, llnode)
free_one(pos);
next = READ_ONCE(entry->next); <-- trigger UAF
} while (!try_cmpxchg(&head->first, &entry, next));
In addition, there is also a theoretical race condition on the
free_by_rcu_ttrace list. This race requires two preconditions: an
in-flight Tasks Trace GP keeping c->call_rcu_ttrace_in_progress == 1,
and concurrent cross-CPU frees repopulating c->free_by_rcu_ttrace with
new nodes. Under these conditions, the following scenario triggers UAF:
// CPU0
// irq work is still busy (on PREEMPT_RT)
alloc_bulk()
llist_del_first(&c->free_by_rcu_ttrace)
entry = smp_load_acquire(&head->first);
do {
if (entry == NULL)
return NULL;
// CPU1
bpf_mem_alloc_destroy()
WRITE_ONCE(c->draining, true)
// wait for CPU0
irq_work_sync()
// CPU2
do_call_rcu_ttrace(tgt(CPU0))
if (c->draining) {
llist_del_all(&c->free_by_rcu_ttrace)
free_all()
}
// CPU0 continue
next = READ_ONCE(entry->next); <-- trigger UAF
while (!try_cmpxchg(&head->first, &entry, next));
Fix this by introducing a raw spinlock to synchronize the concurrent
consumption on waiting_for_gp_ttrace and free_by_rcu_ttrace. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Fix out-of-bounds read of rtt_min in sock_ops
A sockops prog reading skops->rtt_min never checks the sk type: on the
tcp_conn_request() path sock_ops->sk is a request_sock (non-full), and the
ctx rewrite casts it to a tcp_sock (full) and reads rtt_min past the end of
the request_sock, returning dirty adjacent memory.
SEC("sockops")
int prog(struct bpf_sock_ops *skops)
{
switch (skops->op) {
case BPF_SOCK_OPS_RWND_INIT:
leak = skops->rtt_min; /* reads the request_sock OOB */
...
}
}
For instance one such read returned rtt_min=0xffff8881, the high half of a
leaked kernel pointer.
Guarding that cast is exactly what SOCK_OPS_GET_FIELD() does -- it checks
is_locked_tcp_sock and returns 0 when sock_ops->sk is not a locked full
socket. Every other tcp_sock field in sock_ops goes through it; rtt_min is
the only one open-coded, so it skips the check.
Read rtt_min through SOCK_OPS_GET_FIELD() too. rtt_min is a bit special:
it is a struct minmax and we only want the current min, so pass
rtt_min.s[0].v. That is equivalent to the old hand-computed offset
offsetof(struct tcp_sock, rtt_min) + sizeof_field(struct minmax_sample, t)
(s[0] sits at rtt_min + 0 and .v at + sizeof(.t), i.e. what minmax_get()
returns), so the loaded field is unchanged and only the full-sock guard is
added. The two BUILD_BUG_ON()s that protected the hand-computed offset
are no longer needed.
Before patch:
0: r1 = *(u64 *)(r1 +0) ; r1 = skops->sk
1: r1 = *(u32 *)(r1 +2324) ; ((tcp_sock *)sk)->rtt_min.s[0].v
After patch:
0: *(u64 *)(r1 +56) = r9
1: r9 = *(u8 *)(r1 +50) ; is_locked_tcp_sock
2: if r9 == 0 goto pc+4 ; not a locked full sock -> 0
3: r9 = *(u64 *)(r1 +56)
4: r1 = *(u64 *)(r1 +0) ; r1 = skops->sk
5: r1 = *(u32 *)(r1 +2324) ; rtt_min.s[0].v
6: goto pc+2
7: r9 = *(u64 *)(r1 +56)
8: r1 = 0 |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Avoid soft lockup in __htab_map_lookup_and_delete_batch()
__htab_map_lookup_and_delete_batch() has no rescheduling point. The
batch count bounds how many entries are copied out, not how many
buckets are visited, so one BPF_MAP_LOOKUP_BATCH call can walk the
map end to end. The empty-bucket fast path is worse: it stays inside
a single rcu_read_lock() / bpf_disable_instrumentation() section for
any run of consecutive empty buckets.
That holds up on small maps, but it falls apart at scale. On a
144-CPU arm64 host running a CONFIG_PREEMPT_NONE kernel, periodic
BPF_MAP_LOOKUP_BATCH calls against an LRU hash map with 16,777,216
buckets held a CPU inside the batch op for 77+ seconds and triggered
the soft lockup watchdog.
Commit 75134f16e7dd ("bpf: Add schedule points in batch ops") fixed this
same problem in the generic batch ops, but not in this htab-native path,
which every htab-based hash map variant uses for its lookup[_and_delete]
batch ops.
Complete that fix here. Leave the critical section after 64 consecutive
empty buckets, call cond_resched_tasks_rcu_qs(), and resume at the saved
bucket cursor. No locks are held at that point, and resuming from the
cursor is already the function's behavior for non-empty buckets. Add the
same call to the per-bucket loop after copy_to_user(), where every lock
has been dropped. cond_resched_rcu() is not enough here: sleeping with
bpf_prog_active elevated makes tracing programs on that CPU silently
skip their invocations.
Plain cond_resched() is not enough either. It is a no-op under PREEMPT
and PREEMPT_LAZY, the only models arm64 and x86 have offered since
commit 7dadeaa6e851 ("sched: Further restrict the preemption modes").
It is also never a Tasks RCU quiescent state, in any model: the
reschedule counts as a preemption. The walking task stays a holdout and
stalls every synchronize_rcu_tasks() caller, ftrace and BPF trampoline
teardown included, until the syscall returns [1].
cond_resched_tasks_rcu_qs() is the usual tool for that [2]. It reports
the quiescent state at each yield and still reschedules as
cond_resched() does on PREEMPT_NONE and PREEMPT_VOLUNTARY kernels. |
