| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| Issue summary: A malicious remote peer may flood the local QUIC
stack with NEW_CONNECTION_ID frames by avoiding a limit check on
how many connection IDs the remote QUIC stack can use.
Impact summary: The local QUIC stack sends a RETIRE_CONN_ID frame
for every NEW_CONNECTION_ID frame it receives. The RETIRE_CONN_ID
frame is dispatched via the Control Frame Queue (CFQ). If the remote
peer also withholds ACKs, then it can force the local stack
to allocate ~400MB (depending on ACK delay).
CWE: CWE-770: Allocation of Resources Without Limits or Throttling
Description: RFC 9000 sections 5.1.1 and 5.1.2 [1] describe the mechanism
by which a remote peer can notify the local QUIC stack to change the
destination connection ID (a.k.a. CID) the local stack uses to
identify the connection at the remote peer. Each CID is associated
with a sequence number. The sequence number is transmitted
in NEW_CONNECTION_ID and RETIRE_CONNECTION_ID frames to identify the CID
which is being either associated with a connection or retired.
The remote peer sends a NEW_CONNECTION_ID frame to let the local stack know
a new CID is being associated with an existing connection. The
NEW_CONNECTION_ID frame carries the new CID, its sequence number, and the
retire-prior-to number. The retire-prior-to identifies existing
CIDs that are to be retired. The local QUIC stack must send a
RETIRE_CONNECTION_ID for every destination CID whose sequence number
is less than retire-prior-to. The CID becomes retired after the
local stack receives an ACK for its RETIRE_CONNECTION_ID frame.
Although the OpenSSL QUIC stack supports at most one destination CID
for every connection, it can be tricked into processing more than
one RETIRE_CONNECTION_ID frame per connection. The OpenSSL QUIC
stack currently retires the destination CID as soon as it receives
the NEW_CONNECTION_ID, while in fact the destination CID must
be retired after an ACK for the RETIRE_CONNECTION_ID frame is received.
Correcting the flawed logic also fixes the backlog growth.
[1] https://datatracker.ietf.org/doc/html/rfc9000#name-issuing-connection-ids
FIPS impact: no
The FIPS module is not affected as the QUIC implementation is outside of
the OpenSSL FIPS module boundary. |
| This CVE ID has been rejected or withdrawn by its CVE Numbering Authority. |
| Improper neutralization of input during web page generation ('cross-site scripting') vulnerability in Mediawiki - Cargo extension allows Reflected XSS.
This issue affects Mediawiki - Cargo extension: through 3.9.4. |
| Improper neutralization of input during web page generation ('cross-site scripting') vulnerability in Mediawiki - Cargo extension allows Reflected XSS.
This issue affects Mediawiki - Cargo extension: through 3.9.4. |
| Improper neutralization of input during web page generation ('cross-site scripting') vulnerability in the Mediawiki - Cargo extension allows Stored XSS.
This issue affects Mediawiki - Cargo extension: through 3.9.4. |
| Renovate versions from 42.68.1 before 42.96.3 and from 43.0.0 before 43.4.4, including the renovate/renovate Docker images, and Mend Renovate CE/EE images (renovate-ce, renovate-ee-server, renovate-ee-worker) from 13.3.0 before 13.6.0, fail to restrict environment variables to an allowlist when spawning child processes. As a result, child processes (e.g. npm install, postUpgradeTasks, postUpdateOptions) gain full access to all environment variables of the Renovate process, allowing insider or outside attackers to exfiltrate secrets accessible to the Renovate deployment. |
| Renovate versions from 43.65.0 before 43.102.11 contain a remote code execution vulnerability in bazel-module and bazelisk managers when using lockFileMaintenance. Attackers can execute arbitrary code by providing malicious dependencies that are referenced in bazel mod deps calls, such as within ctx.execute statements. |
| In the Linux kernel, the following vulnerability has been resolved:
net/mlx5e: Fix use-after-free race in sample_restore_put()
Concurrent teardown of TC sample rules sharing the same restore
context may re-read restore->count after dropping restore_lock.
At that point another thread may already have completed cleanup and
freed the restore object.
Use the result of the refcount decrement while holding restore_lock to
determine whether cleanup is needed. |
| In the Linux kernel, the following vulnerability has been resolved:
eth: nfp: bound the ntuple rule dump by the caller's buffer size
nfp_net_get_fs_loc() dumps every entry of nn->fs.list into rule_locs[]
without consulting cmd->rule_cnt, which is how many entries the caller
had room for. ETHTOOL_GRXCLSRLALL requires no CAP_NET_ADMIN and the
ioctl sizes the buffer from the rule_cnt userspace passes in, so once an
admin has installed flow steering rules any user can ask for fewer slots
than there are rules and run off the end of the allocation. A rule_cnt
of 0 leaves the buffer pointer NULL and the walk dereferences it.
Bail out with -EMSGSIZE when the buffer fills up, the way the other
ntuple capable drivers do, and report how many locations were filled so
a shrinking rule list does not leave the caller reading stale slots. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Reject untrusted allocated-object pointers
When the final RCU read-side critical section ends, a local kptr is demoted
to PTR_UNTRUSTED but retains MEM_ALLOC. The pointer may be NULL or may refer
to an object whose lifetime is no longer protected.
type_is_ptr_alloc_obj() nevertheless recognizes any PTR_TO_BTF_ID with
MEM_ALLOC as a live allocated object. In particular, a refcount-only local
kptr never carries NON_OWN_REF, so it still passes the
bpf_refcount_acquire() argument check after RCU protection ends. The kfunc
can then dereference NULL or stale memory.
Make type_is_ptr_alloc_obj() reject PTR_UNTRUSTED pointers. Since
type_is_non_owning_ref() is based on the same predicate, graph kfunc
arguments obey the same live-object requirement. Fault-protected reads of
the demoted pointer remain valid: writes are already rejected, and read
fixups use bpf_may_fault_on_deref() rather than this predicate.
[ kkd: Rewrote commit log ] |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Require MEM_PERCPU for percpu kptr stores
map_kptr_match_type() treats perm_flags as the set of register type flags
that a kptr field permits. Adding MEM_PERCPU to that set for
BPF_KPTR_PERCPU does not require the source register to carry it, however.
The subset test consequently accepts both a plain bpf_obj_new() allocation
and a referenced kernel pointer into a __percpu_kptr map field.
Loads from the field are always marked MEM_PERCPU. Consumers then treat the
stored value as the cookie returned by bpf_percpu_obj_new(): per-CPU pointer
helpers relocate it, and map teardown selects the per-CPU free path. A plain
allocation can therefore provide an arbitrary kernel read/write, while a
kernel pointer can be relocated into an invalid address or sent through a
missing destructor.
Require the source MEM_PERCPU flag to match the destination field kind.
This preserves valid bpf_percpu_obj_new() stores and rejects both the
program-BTF and kernel-BTF variants. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Mark the zero register precise for a register-form NULL check
check_cond_jmp_op() accepts "if rA <op> rB" as a NULL check for a
nullable pointer rA when rB is a scalar known to be zero,
lifts PTR_MAYBE_NULL from rA in the corresponding branch and does not
mark rB precise. Consider the following program:
r0 = bpf_get_prandom_u32();
r6 = 1; /* the r6 == 0 path is explored first */
if (r0 == 0) goto 1f;
r6 = 0;
1:
r0 = bpf_map_lookup_elem(map, &0); /* absent, NULL at runtime */
if (r0 == r6) goto 2f; /* taken as a NULL check for r0 */
*(u8 *)(r0 + 0); /* verifier: map value; runtime: zero */
2:
return 0;
The r6 == 0 path is explored first and the dereference is accepted.
The r6 == 1 path is pruned at the checkpoint recorded for (1),
so the comparison is never verified with a non-zero r6. At runtime a
failed lookup returns NULL, NULL != 1 takes the non-NULL edge and the
program dereferences a pointer that is zero. |
| ZigBee ZCL protocol dissector crash in 4.6.0 to 4.6.8 allows denial of service |
| MBIM protocol dissector crash in 4.6.0 to 4.6.8 and 4.4.0 to 4.4.18 allows denial of service |
| SCTP protocol dissector crash in 4.6.0 to 4.6.8 and 4.4.0 to 4.4.18 allows denial of service |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Don't resurrect a scalar id dropped by collect_linked_regs()
check_cond_jmp_op() copies the compared registers into
env->{false,true}_reg{1,2} before collect_linked_regs() runs and copies
those snapshots back into both branch states afterwards.
collect_linked_regs() records at most LINKED_REGS_MAX members of a
linked registers group in the jump history and calls clear_scalar_id()
for every member that does not fit. The compared register is not exempt
from that.
As a consequence, sync_linked_regs() might adjust ranges for more
registers than bpf_bt_sync_linked_regs() can propagate precision to.
Collect the linked registers before the snapshots are taken instead.
This might lead to some unnecessary clear_scalar_id's, but from
previous testing situations with many linked registers are
extremely rare. |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Don't infer non-NULL from a pointer with an unbounded offset
reg_not_null() decides that a register holds a non-NULL value by
looking at its type alone. For pointer types that allow arithmetic the
type only guarantees a non-NULL base, in case of an unbound offset
the runtime offset value might still add up to NULL.
Consider the followng program:
r6 = bpf_map_lookup_elem(map, &0); /* present */
if (r6 == 0) return 0;
r7 = bpf_map_lookup_elem(map, &1); /* absent, NULL at runtime */
r8 = r7;
r8 -= r6; /* pointer - pointer: unknown scalar, -r6 */
r8 <<= 1;
r8 >>= 1; /* any non-negative offset is accepted by */
/* check_reg_sane_offset_ptr() */
r6 += r8; /* verifier: map value; runtime: zero */
if (r7 != r6) return 0;
*(u8 *)(r7 + 0); /* r7 is inferred non-NULL, both are zero */
At runtime both registers are zero, the comparison is true and the
load faults with NULL pointer dereference.
Require the offset to be within +-BPF_MAX_VAR_OFF in reg_not_null(). |
| In the Linux kernel, the following vulnerability has been resolved:
bpf: Mark syscall helpers as sleepable
bpf_sys_bpf() executes the bpf(2) syscall body, which can take mutexes,
allocate with GFP_KERNEL, and wait for an RCU grace period.
bpf_sys_close() reaches close_fd() and filp_close(), which can sleep as
well.
Both helpers are limited to BPF_PROG_TYPE_SYSCALL, whose main program is
sleepable. That does not make every callback sleepable: a syscall program
can register a bpf_timer callback, and the verifier checks that callback
in a non-sleepable context while retaining the syscall helper set.
Without .might_sleep on the prototypes, such a callback can invoke
bpf_sys_bpf() from hrtimer softirq context and trigger a
scheduling-while-atomic failure. bpf_sys_close() is exposed through the
same missing context check.
Set .might_sleep on both prototypes so the existing helper-context check
rejects them from timer callbacks and other atomic regions. Calls from the
sleepable main body remain valid. |
| In the Linux kernel, the following vulnerability has been resolved:
net/rds: acquire RDS_IN_XMIT in rds_tcp_reset_callbacks()
rds_tcp_reset_callbacks() quiesces the transmit path by setting the
path state to RDS_CONN_RESETTING and then waiting for RDS_IN_XMIT to
be sampled clear before swapping the underlying socket and calling
rds_send_path_reset().
Sampling the bit clear is not the same as owning it: rds_send_xmit()
can re-acquire RDS_IN_XMIT right after the wait_event() returns. Its
state recheck after taking the lock is a store-buffering pattern (the
resetter writes the state and reads the bit, the sender writes the
bit and reads the state) and acquire_in_xmit() is only an acquire
operation, so on weakly ordered architectures both sides can miss
each other's write and the transmit path then runs concurrently with
rds_send_path_reset() rewriting cp_xmit_* state - which is exactly
what the comment above rds_send_path_reset() tells its callers to
prevent.
Take the lock instead, hold it across the socket swap and
rds_send_path_reset(), and release it with a wake-up at the end. The
lock-ordering constraint documented above the wait still holds: the
lock is acquired before lock_sock(), so a sender inside tcp_sendmsg()
can never be waited on while we hold the socket lock.
Two details of the old code go away with the same change:
- t_sock is now read only after the lock is acquired. The old code
cached it before waiting; the teardown in rds_conn_shutdown()
releases that socket and clears t_sock, so a pointer cached before
the wait can be stale by the time the accept path resumes. Reading
it under RDS_IN_XMIT is what makes the exclusion complete once the
teardown owns the same lock, which the next patch arranges; until
then the teardown still only samples the bit, and the two paths
remain as exposed to each other as they are today.
- The old !osock early path called rds_send_path_reset() with no
serialization at all. It now runs under the lock like the normal
path. The conditional RDS_CONN_RESETTING transition of the
previous patch happens before the socket check either way: a path
found without a socket is either still connecting (its reconnect
worker blocked on t_conn_path_lock) and legitimately goes
RESETTING -> UP on the new socket, or it has been torn down
meanwhile and is dropped.
The in-function comment describing the old wait-based quiesce is
rewritten to describe the lock-based one, and the stale block comment
above the function (which still described a return value and an
incomplete list of t_sock writers) is refreshed to name all four
writers - the connect, accept, teardown and swap paths - and what
serializes each of them. |
| In the Linux kernel, the following vulnerability has been resolved:
net/rds: use wq_has_sleeper() in release_in_xmit()
release_in_xmit() clears RDS_IN_XMIT with clear_bit_unlock() and then
checks waitqueue_active() to decide whether anyone needs waking.
clear_bit_unlock() is only a release operation: it orders the
critical section before the bit clear, but does not order the
subsequent plain load of the wait queue head after it. The waiter
side does the mirror image - it adds itself to the wait queue and
then tests the bit. That is the classic store-buffering pattern: the
releasing CPU can read the wait queue as empty while the waiting CPU
still reads the bit as set, so the sleeper is never woken.
The waiters are rds_conn_shutdown() and rds_tcp_reset_callbacks(),
both in uninterruptible wait_event() with no timeout. A lost wake-up
strands the shutdown worker on its single-threaded workqueue until
some other sender releases the bit again - and on a connection that
is being torn down precisely because it failed, there may never be
another sender.
The barrier used to be there: release_in_xmit() did clear_bit()
followed by smp_mb__after_atomic() until commit 1422f28826d2 ("rds:
introduce acquire/release ordering in acquire/release_in_xmit()")
folded both into clear_bit_unlock(), which strengthened the lock
hand-off but silently dropped the full barrier the wake-up check
depends on. The refill counterpart, release_refill() in
net/rds/ib_recv.c, still carries its smp_mb__after_atomic() for
exactly this reason.
Use wq_has_sleeper(), which is waitqueue_active() preceded by the
required full barrier. |
