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| CVE | Vendors | Products | Updated | CVSS v3.1 |
|---|---|---|---|---|
| CVE-2026-98125 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: smb/client: fix stale page cache in insert/collapse range smb3_insert_range() and smb3_collapse_range() use truncate_pagecache_range() to invalidate the affected page cache. However, if off or old_eof is not page-aligned, the boundary pages are only partially zeroed and remain uptodate. As a result, the client may return stale data after a successful insert/collapse range operation. For example, with 4K pages: page 0 page 1 page 2 0------4K 4K------8K 8K------12K ^ ^ off=2K old_eof=10K Page 1 is removed from the page cache, while the boundary pages are only partially zeroed. After COPYCHUNK moves the data on the server, these cached pages may still return stale data. This can be reproduced on a CIFS mount: bash -c ' FILE=/mnt/scratch/repro # Use a 6 KiB file so EOF is not page-aligned. dd if=/dev/urandom of=/tmp/src bs=1K count=6 status=none # Expected: a 4 KiB hole followed by the original data. rm -f /tmp/expected truncate -s 4K /tmp/expected cat /tmp/src >> /tmp/expected cp /tmp/src "$FILE" # Prime the page cache before moving data on the server. cat "$FILE" > /dev/null fallocate --insert-range -o 0 -l 4K "$FILE" if cmp -s /tmp/expected "$FILE"; then echo "readback: OK" else echo "readback: STALE DATA" fi ' Fix this by writing back dirty data and discarding the page cache from the start of the page containing off to EOF before moving data on the server. | ||||
| CVE-2026-98126 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: smb/client: validate new EOF for zero range When FALLOC_FL_ZERO_RANGE is used without FALLOC_FL_KEEP_SIZE, smb3_zero_range() may extend EOF without checking RLIMIT_FSIZE, allowing the file to grow beyond the caller's file-size limit. Fix this by calling inode_newsize_ok() before sending the zero-range request when the operation would extend EOF. Reproducer, using a file on a CIFS mount: bash -c ' FILE=/mnt/cifs/repro trap "" SIGXFSZ ulimit -f 3072 truncate -s 2M "$FILE" fallocate --zero-range -o 0 -l 4M "$FILE" echo "fallocate rc=$?" stat -c "file size=%s" "$FILE" ' Before this change, the operation succeeds despite the 3 MiB limit: fallocate rc=0 file size=4194304 After this change, fallocate fails and leaves the file at 2 MiB. | ||||
| CVE-2026-98127 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: smb/client: validate new EOF for insert range smb3_insert_range() does not check if the new file size (i_size + len) is valid. This allows FALLOC_FL_INSERT_RANGE to bypass RLIMIT_FSIZE, exceed s_maxbytes, or produce a size outside the loff_t range. Use check_add_overflow() to calculate the new EOF. Validate it with inode_newsize_ok() before modifying the file. Reproducer, using a file on a CIFS mount: bash -c ' FILE=/mnt/cifs/repro trap "" SIGXFSZ ulimit -f 3072 # RLIMIT_FSIZE = 3 MiB # A regular write is stopped at 3 MiB. dd if=/dev/zero of="$FILE" bs=1M count=4 status=none stat -c "size after write: %s" "$FILE" # Insert 2 MiB into a 2 MiB file. truncate -s 2M "$FILE" fallocate -i -o 0 -l 2M "$FILE" stat -c "size after insert: %s" "$FILE" ' Before this change, the regular write stops at the 3 MiB limit, but insert range grows the file to 4 MiB: dd: error writing '/mnt/cifs/repro': File too large size after write: 3145728 size after insert: 4194304 After this change, insert range also fails at the limit and leaves the 2 MiB file unchanged: dd: error writing '/mnt/cifs/repro': File too large size after write: 3145728 fallocate: fallocate failed: File too large size after insert: 2097152 | ||||
| CVE-2026-98133 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: ntfs: leave HasEA flag untouched on setxattr failure In ntfs_set_ea(), the exit path unconditionally updates the HasEA flag based on ea_info_qsize. When an error occurs before ea_info_qsize is updated, NInoClearHasEA() hides existing on-disk EAs until the inode is evicted. Only update the flag on success. | ||||
| CVE-2026-98051 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: net: bcmasp: fix tx_spb_ring_full() checking same slot cnt times The loop initialised next_index from intf->tx_spb_index on every iteration, so incr_ring() always produced the same result and only one slot was ever tested. Move the initialisation before the loop so each iteration advances next_index and the function correctly checks that cnt consecutive descriptor slots are available before allowing a new transmission. | ||||
| CVE-2026-98052 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: net: bcmasp: clear txcb->last before writing each descriptor bcmasp_xmit() only wrote txcb->last = true for the final fragment of an SKB; non-final fragments left the field untouched. If a descriptor slot was reused while it still held a stale true from a previous SKB (possible when tx_spb_ring_full() underreported fullness), bcmasp_tx_reclaim() would see last == true mid-SKB and call dev_consume_skb_any() prematurely, freeing the sk_buff while its remaining fragments were still in flight. Unconditionally clear txcb->last before the conditional set so every descriptor slot starts from a known false state regardless of what a prior transmission left behind. | ||||
| CVE-2026-98062 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: bpf: Mark signal tracepoint siginfo arguments as scalar The signal_generate and signal_deliver tracepoints declare their info argument as a struct kernel_siginfo pointer. btf_ctx_access() therefore treats it as a trusted pointer for tp_btf programs. Signal delivery also uses SEND_SIG_NOINFO and SEND_SIG_PRIV as special values for this argument. Those values are zero and one respectively, and are not pointers. A tp_btf program can currently dereference either value and fault the kernel. In particular, signal_generate can run from timer interrupt context, turning the fault into a kernel panic. Record both tracepoints in raw_tp_null_args[] and mark argument one as a non-pointer. This preserves scalar access to the cookie while rejecting direct and helper-mediated pointer use. Merely marking it nullable would not suffice because SEND_SIG_PRIV is nonzero. | ||||
| CVE-2026-98067 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: erofs: disable LZ4 rolling decompression for now LZ4 rolling decompression [1] was introduced to reduce the memory footprint of temporary pages: For many cases, it is needed for users to read small data within a compressed extent (pcluster), either due to random small read, or since uptodate folios (typically order-0) cannot be reused for decompression again since decompression algorithm refills already-uptodate folios. Rolling decompression works because LZ4 is LZ77-based and only refers to the most recent 64 KiB of decompressed data, so in theory only a bounded rolling window of temporary pages is needed when decompressing. It can save a lot of temporary memory, e.g. 601,960-byte data can be compressed into a 256k LZ4 compressed extent, which means it needs 146 extra pages per request in the worst case if rolling decompression is disabled. However, the upstream LZ4 implementation is not under EROFS' control: For example, the literal copy memmove() may still **copy long literals backward** on x86 based on the address comparison even when the source and destination ranges do not overlap (IOWs, inline decompression doesn't need to be considered here). That breaks the rolling assumption and makes the optimization broken. Disable it for now to make sure the data correctness first since EROFS is used everywhere now: The rolling window approach can be revived once we either ensure that the official LZ4 code always copies forward for non-overlapping ranges or maintain our own LZ4 implementation in EROFS. The main impact is a higher runtime memory footprint; However, recent commit 0f6273ab4637 ("erofs: add a reserved buffer pool for lz4 decompression") helps mitigate this when enabled but it's still not perfect. [1] https://www.usenix.org/conference/atc19/presentation/gao § 3.3 Decompression | ||||
| CVE-2026-98068 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7.0 High |
| In the Linux kernel, the following vulnerability has been resolved: net/rds: don't let rds_conn_shutdown() consume a concurrent drop rds_conn_shutdown() finishes by moving the path from RDS_CONN_DISCONNECTING to RDS_CONN_DOWN, and also accepts RDS_CONN_ERROR as the starting state of that final transition, so that a FIN processed in softirq context during the teardown does not derail the shutdown into a noisy error path. But consuming that RDS_CONN_ERROR also consumes the shutdown pass that came with it: rds_conn_path_drop() sets RDS_CONN_ERROR and then queues cp_down_w, and a pass that starts on a path already in RDS_CONN_DOWN is a no-op. For the FIN case that is harmless - the socket the FIN arrived on is the very socket the teardown just released. It is not harmless for a dropper that attached something to the path first. rds_tcp_accept_one() is such a dropper. Its path claim in rds_tcp_accept_one_path() transitions RDS_CONN_DOWN -> RDS_CONN_CONNECTING, and a concurrent drop - a FIN on a previous socket in softirq context, an administrative reset - can put the path into RDS_CONN_ERROR between that claim and the state check that follows, which accepts RDS_CONN_ERROR. The accept then installs the freshly accepted socket with rds_tcp_set_callbacks() while the queued teardown - which sampled tc->t_sock before this socket existed - is still running. rds_connect_path_complete() fails its transition to RDS_CONN_UP and drops the path again, queueing the pass that should reap the socket it just installed. If the in-flight shutdown's final transition consumes that drop's RDS_CONN_ERROR, the queued pass finds the path in RDS_CONN_DOWN and does nothing. The installed socket is never torn down: it sits established with its callbacks armed and its rds_tcp_connection on rds_tcp_tc_list, the peer sees a connection that nothing ever reads, and the path is wedged in RDS_CONN_DOWN until some later event drops it again. Reproduced with widened race windows as an ever-growing receive queue on a socket owned by a path stuck in RDS_CONN_DOWN, with the peer's send path wedged behind it. Make the final transition only DISCONNECTING -> DOWN. If it fails because the path is in RDS_CONN_ERROR, a drop raced the teardown: cancel the reconnect timer and clear RDS_RECONNECT_PENDING - the one piece of the skipped tail that must not be left behind - and return, letting the pass the drop queued finish the job: it tears down whatever attached to the path in the meantime, completes the transition to RDS_CONN_DOWN, and re-arms the reconnect from its own tail. The timer quiesce in that branch matters because the racing drop does not always queue that pass: rds_conn_path_drop() returns without queueing when a destroy is pending - exactly the situation during a netns teardown or module unload, when a FIN on the dying socket is processed while rds_conn_path_destroy() flushes cp_down_w. If the flushed pass is the one that takes this return, no later pass exists, and rds_conn_path_destroy() would find cp_conn_w still armed (WARN_ON) and then free a path whose reconnect timer can still fire. With the cancel in the branch, every exit of a shutdown pass leaves the timer quiesced no matter which pass completes the transition. The FIN case keeps making progress, one pass later and still without noisy logging. Any other state keeps today's rds_conn_path_error() handling; no current cp_state writer can leave a DISCONNECTING path in anything but RDS_CONN_ERROR (every other writer is a cmpxchg from a non-DISCONNECTING state), so that branch is defensive. On kernels without the preceding patches the same hazard exists with the sample-based quiesce; the fix applies there equally. | ||||
| CVE-2026-98069 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 8.1 High |
| In the Linux kernel, the following vulnerability has been resolved: net/rds: acquire the fastpath locks in rds_conn_shutdown() rds_conn_shutdown() quiesces the transmit and receive-refill paths by waiting for RDS_IN_XMIT and RDS_RECV_REFILL to be sampled clear, and then runs the transport shutdown and rds_conn_path_reset(). Sampling the bits clear is not the same as owning them: the moment after the wait_event() returns, rds_send_xmit() can re-acquire RDS_IN_XMIT (or rds_ib_recv_refill() can re-acquire RDS_RECV_REFILL) and run concurrently with the teardown. The sender does recheck the connection state after taking the lock, but that recheck is a classic store-buffering pattern: teardown writes the state and reads the bit while the sender writes the bit and reads the state. 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 while the transport zeroes its rings (e.g. rds_ib_ring_init()) and rds_send_path_reset() rewrites the transmit state under it. Oracle UEK fixed the same class of crashes - a 14-year tail of BUG_ON()s in rds_ib_sub_signaled(), unexpected op-codes and NULL dereferences in rds_ib_send_cqe_handler() during failover testing - by making the teardown path *acquire* the fastpath bit locks instead of testing them ("rds: Make sure transmit path and connection tear-down does not run concurrently"). Ownership of a single word is decided by RMW atomicity, so no cross-variable ordering is needed. Do the same here: take both locks before calling the transport shutdown, hold them across rds_conn_path_reset(), and release them explicitly with a wake-up afterwards. Both are released with clear_bit_unlock(), so that the ring re-initialization done by the transport shutdown and the transmit state rewritten by rds_send_path_reset() are ordered before either bit is seen clear by the next acquire_in_xmit() or acquire_refill(). The fastpath users of these bits - rds_send_xmit() and rds_ib_recv_refill() - are trylock style and back off while teardown owns the locks, so no new lock dependency is introduced for them. rds_tcp_reset_callbacks() is different: since the previous patch it acquires RDS_IN_XMIT as well, and it blocks doing so, so its wait now spans the teardown instead of at most one send batch. That waiter runs from rds_tcp_accept_one() on the single-threaded krdsd workqueue and holds rds_tcp_accept_lock and t_conn_path_lock while it waits, so a duelling SYN accepted while its path is being torn down parks accept processing for the duration of the teardown - for TCP bounded by the (up to 5 s) drain loop in rds_tcp_conn_path_shutdown(). An IB path's drain in rds_ib_conn_path_shutdown() has no round cap, but no blocking waiter either: rds_tcp_reset_callbacks() is the only blocking acquirer of these bits and waits only on its own TCP path, and the fastpaths are trylock-and-back-off on both transports, so a long IB drain lengthens only that path's own quiesce. The window is narrow: the accept-side state check has to pass before the teardown moves the path to RDS_CONN_DISCONNECTING. Because krdsd is a single global workqueue, everything else queued there - accept processing for other connections and network namespaces, and the flush_workqueue(rds_wq) in rds_tcp_listen_stop() during namespace teardown - waits behind the parked accept worker for that time. It cannot deadlock, although the waits do point at each other: the teardown blocks until the bit's holder releases it, and the holder may be that krdsd accept worker. The holder finishes without needing anything the teardown owns: the sync cancels rds_tcp_reset_callbacks() issues target cp_send_w and cp_recv_w on the path's ordered cp_wq, whose only execution slot is occupied by the blocked cp_down_w itself, so they are pending at most and cancel without flushing - a reliance on cp_wq being ordered that is now noted next to those cancels (on ---truncated--- | ||||
| CVE-2026-98020 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: pds_core: fix cmd_regs access racing BAR unmap on reset pdsc_reset_prepare() and pdsc_reset_done()'s pdsc_map_bars() error path clear/iounmap cmd_regs without devcmd_lock, and pdsc_legacy_firmware_update()'s download loop derefs cmd_regs after dropping and retaking the lock without re-checking. An FLR concurrent with a devlink flash can unmap cmd_regs under an in-flight devcmd, causing a NULL deref or a write to unmapped MMIO. Take devcmd_lock across the BAR unmap/remap, and re-check cmd_regs in the download loop. Only the PF maps cmd_regs and runs devcmd, so skip the unmap on a VF, as pdsc_remove() and pdsc_reset_done() already do. A reset that completes entirely within the unlocked window is not a correctness problem for the image: the device clears its update session, so a resumed download is rejected, and it verifies the staged image before writing a flash slot, reporting PDS_RC_BAD_FW rather than activating it. pdsc_unmap_bars() also clears info_regs, intr_status and intr_ctrl. The interrupt and start/stop readers of those are quiesced before the unmap by pdsc_fw_down(), which frees the interrupts and tears down the queues. The debugfs readers are not, since those files outlive a reset; that is pre-existing and out of scope here. | ||||
| CVE-2026-98021 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: net: reject oversized tx_queue_len at netlink parse time rtnl_create_link() assigns IFLA_TXQLEN directly to dev->tx_queue_len without going through netif_change_tx_queue_len(), so a device created with "ip link add ... txqueuelen 500000" bypasses the S16_MAX cap and still triggers the oversized ring allocations in pfifo_fast, tun and tap. The veth peer nest (rtnl_nla_parse_ifinfomsg()) and the RTM_NEWLINK-on-existing-device path reach the same sinks. Enforce the cap in ifla_policy instead: IFLA_TXQLEN becomes NLA_POLICY_FULL_RANGE(NLA_U32, &txqlen_range) with txqlen_range = { .min = 0, .max = S16_MAX }. All netlink consumers parse against this policy - rtnl_setlink(), rtnl_newlink() (create and change), and the veth peer nest - so every netlink path is capped at parse time and rejects the attribute with -ERANGE plus a proper "integer out of range" extack message before any device state is modified (the RTM_SETLINK half-application wart is gone with it). Document the bound in the rt-link.yaml netlink spec. Conditions to recreate the bug: - CONFIG_NET_SCHED=y, CONFIG_VETH=y, CONFIG_USER_NS=y, CONFIG_NET_NS=y. - Unprivileged user in a fresh user+net namespace (unshare -Urn): ip link add v0 txqueuelen 500000 type veth peer name v1 -> on the fixed kernel this is rejected with -ERANGE ("integer out of range" extack) instead of installing an oversized tx_queue_len that later inflates pfifo_fast/tun/tap ring allocations. - ip link set v0 txqueuelen 500000 is likewise rejected at parse time. | ||||
| CVE-2026-98023 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: vxlan: reject dynamic fdb entries that reference a nexthop id The commit cited in the Fixes tag allowed VXLAN FDB entries to point to FDB nexthops so that overlay traffic could be load balanced across multiple VTEPs. Such entries can only be configured from user space, cannot be learned and cannot roam. They only make sense with a user space control plane such as E-VPN where data plane learning is disabled. Despite that, the VXLAN driver does not currently prevent such entries from being configured with the "dynamic" flag. The per-nexthop FDB list is only protected by the per-device hash lock, which is not sufficient when two VXLAN devices point to the same FDB nexthop and therefore share the list. Aging runs in softirq context without RTNL, so an entry deleted by one device can race with an addition or deletion from the other, leading to list corruption: list_del corruption. next->prev should be ffff8881069d9548, but was dead000000000122. (next=ffff8881069d9448) WARNING: CPU: 0 PID: 90 at lib/list_debug.c:65 __list_del_entry_valid_or_report+0x1aa/0x210 ... vxlan_fdb_destroy+0x5b8/0xad0 vxlan_cleanup+0x328/0x450 call_timer_fn+0x2a/0x1c0 run_timer_softirq+0x18c/0x210 BUG: KASAN: slab-use-after-free in vxlan_fdb_destroy Fix this by rejecting the bogus configuration of dynamic FDB entries that point to FDB nexthops, both when created and when an existing entry is updated. As such, the per-nexthop FDB list is only ever mutated under the RTNL lock. Add test cases to make sure that this does not regress in the future. | ||||
| CVE-2026-98024 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: s390/ism: folio_put() after error dmb->cpu_addr was allocated via folio_alloc(). Use folio_put() instead of kfree() in the error exit of ism_alloc_dmb() to avoid slab allocator corruption. While at it, reset dmb->cpu_addr after folio_put to avoid unintentional UAF by future callers. | ||||
| CVE-2026-98025 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7.0 High |
| In the Linux kernel, the following vulnerability has been resolved: net: usb: cx82310_eth: drop URB after 0xffff reboot sentinel to prevent partial_data heap overflow The 0xffff length sentinel detects a router reboot and schedules re-enabling of ethernet mode, but then falls through to the rest of the loop body. The next check is } else if (len > CX82310_MTU) { which is the else of the just-matched if -- it never fires for len == 0xffff. The MTU bound that normally caps the incomplete-packet save path is silently bypassed. With 0xffff > skb->len always true (rx_urb_size is 4096), the incomplete-packet branch saves dev->partial_len = skb->len bytes into dev->partial_data. partial_data is kmalloc(hard_mtu) = kmalloc(CX82310_MTU + 2) = 1516 bytes, but skb->len after the 2-byte header pull can be up to 4094. A device that sends a 4096-byte URB starting with [0xff 0xff] therefore copies 4094 device-provided bytes into a buffer allocated for 1516 bytes, exceeding its requested size by 2578 bytes. The next URB then reads dev->partial_len (4094) back from the same 1516-byte buffer and dev->partial_rem (65535 - 4094 = 61441) from the new URB's ~4KB skb, both well past their allocations, and delivers the spliced result as a 64KB "frame" to the network stack. Bail out of rx_fixup after scheduling the re-enable work; the remainder of a reboot-marker URB is not meaningful packet data. This restores the invariant that partial_len < CX82310_MTU + 2 on the save path, since every other route there has already passed the MTU check. | ||||
| CVE-2026-98026 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: net: bridge: mcast: properly convert mglist to rcu Sashiko reported a bug [1] that br_multicast_del_port_group unlists the port group not using proper rcu helper that preserves the next pointer and after that immediately frees the port group without waiting for rcu grace period. The only rcu walker of mglist is br_multicast_list_adjacent() and it turns out that function has always been buggy because mglist was never properly converted to RCU. Fix it by converting it to rcu and moving its initialization after eth_addr's. Initializing p->next can use RCU_INIT_POINTER because we have a barrier from the hlist_add_head_rcu call later, besides we're initializing an unpublished structure anyway. [1] https://netdev-ai.bots.linux.dev/sashiko/#/patchset/20260826014200.362304-1-littleddfu%40gmail.com | ||||
| CVE-2026-98030 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7 High |
| In the Linux kernel, the following vulnerability has been resolved: net: dsa: bcm_sf2: bound the CFP rule dump by the caller's buffer size bcm_sf2_cfp_rule_get_all() walks the whole cfp.unique bitmap into rule_locs[] without consulting nfc->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 CFP 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. | ||||
| CVE-2026-98031 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 5.5 Medium |
| In the Linux kernel, the following vulnerability has been resolved: nexthop: Initialize extack in remove_nh_grp_entry() remove_nh_grp_entry() prints the extack message when a listener fails to replace the reduced nexthop group. However, extack is not initialized and listeners are not required to set a message when returning an error. Neither netdevsim nor mlxsw do so when an allocation fails, resulting in the dereference of an uninitialized stack pointer. Fix by zero-initializing extack, as was done in commit 6347c5314cee ("nexthop: initialize extack in nh_res_bucket_migrate()"). | ||||
| CVE-2026-98032 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7.0 High |
| In the Linux kernel, the following vulnerability has been resolved: tracing: Fix subbuf resize races with trace_pipe_raw readers Concurrent subbuffer resizes may crash trace_pipe_raw readers or leak uninitialized memory to userspace due to stale size values. Modify ring_buffer_alloc_read_page() to handle the resizing of an existing buffer_data_read_page if necessary and add a new ring_buffer_read_page_size(). This new function enables ring-buffer buffer_data_read_page users to not call the racy ring_buffer_subbuf_size_get(). This makes the spare_size member of ftrace_buffer_info redundant. Finally, handle buffer_data_read_page/reader_page order discrepancy in ring_buffer_read_page(). On a mismatch simply copy manually the data to the buffer_data_read_page. | ||||
| CVE-2026-98033 | 1 Linux | 1 Linux Kernel | 2026-09-29 | 7.0 High |
| In the Linux kernel, the following vulnerability has been resolved: bpf: Preserve inner map identity in callback frames Callback frame constructors initialize map-typed argument registers with __mark_reg_known_zero() and then restore map_ptr. This clears map_uid, which is the only field distinguishing inner maps that share an inner_map_meta template. When a timer callback invokes bpf_for_each_map_elem() on a second inner map, both the saved first map and the second map value can reach the nested callback as the same template with map_uid zero. bpf_timer_init() then accepts pairing the timer from the second map with the first map. The runtime records the first map in the timer without taking a reference. Freeing that map does not find the timer stored in the second map, so a later timer callback dereferences the freed map. Copy map_uid from the same caller register as map_ptr when constructing for-each, timer/workqueue, and task-work callback arguments. The existing identity check can then reject mismatched inner maps while allowing a callback value to be paired with its actual map. | ||||