| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
dm-integrity: fix NULL pointer dereference when the 'R' flag is used
If the dm-integrity device has the SB_FLAG_DIRTY_BITMAP flag set and the
user activates the device in the 'R' mode, a crash in dm_integrity_resume
happens because the function attempts to read the journal containing the
bitmap.
This patch makes dm-integrity skip any writes to the device in
dm_integrity_resume if the device is activated in the 'R' mode. |
| Heap-based buffer overflow in SQL Server allows an authorized attacker to execute code over a network. |
| In the Linux kernel, the following vulnerability has been resolved:
dm-integrity: validate the superblock on resume
dm_integrity_resume() re-reads the superblock from the device so that it
picks up the flags and the recalculate position. It performs no
validation on the result, while the constructor validates the superblock
it reads and sizes all the in-memory structures according to it. The user
may modify the superblock on the underlying device while the dm-integrity
device is suspended, so that the two no longer agree.
In particular, access_journal_data() shifts the journal entry index by
ic->sb->log2_sectors_per_block, while ic->journal_section_sectors and the
journal page list were computed with the value that was present at
constructor time. Increasing log2_sectors_per_block makes the index run
past the end of the journal, and dm-integrity then writes 512 bytes
through lowmem_page_address(NULL).
Snapshot the validated superblock in the constructor and refuse to resume
if any of the fields that describe the on-disk geometry changed. |
| Inefficient algorithmic complexity in ASP.NET Core allows an unauthorized attacker to deny service over a network. |
| In the Linux kernel, the following vulnerability has been resolved:
ipv6: udp: Create exceptions before socket matching
Currently, when ICMPv6 Packet Too Big and Redirect Message packets are
locally delivered and quote a UDP packet, an exception is only created
in the IPv6 exception cache if the kernel can match the UDP packet to an
existing socket.
This behavior allows off-path attackers to conduct a side-channel attack
on the exception cache in order to discover the ephemeral port used by a
connected UDP socket.
Commit 4785305c05b2 ("ipv6: use siphash in rt6_exception_hash()") and
commit a00df2caffed ("ipv6: make exception cache less predictible") tried
to mitigate such attacks by making it harder for attackers to discover
hash collisions in the exception cache and by randomizing the number of
exceptions a hash bucket can hold, respectively. Unfortunately, both of
the mitigations can be bypassed.
Instead, mitigate such attacks by always creating an exception, even
before trying to find a matching socket. Do that by calling
ip6_update_pmtu() and ip6_redirect(), the helpers used when the quoted
packet did not originate from a socket.
This means that guesses (right or wrong) from an off-path attacker will
always result in an exception being created or updated in the cache that
the attacker can observe.
Pass the ifindex of the ingress device and the default uid, in a similar
fashion to icmpv6_err(). Unlike IPv4, an oif of 0 would not match any
nexthop in ip6_redirect_nh_match() and no exception would be created in
response to a Redirect Message.
Note that this does not allow attackers to create exceptions that they
could not create before, as both helpers can already be reached with
little to no validation. For example, by sending an ICMPv6 error that
quotes an ICMPv6 Echo Reply or one that quotes a UDP source port that
matches a wildcard socket.
Also note that in the good case (matched socket) the above scheme comes
at the cost of an extra route lookup, as the no socket helpers perform
their own lookup before the one performed by ip6_sk_update_pmtu() /
ip6_sk_redirect(). When the two resolve to different nexthops, it also
results in two exceptions being created for the same destination IP. One
in the exception cache of the nexthop resolved by the no socket helpers
and another in the exception cache of the nexthop used by the socket. |
| Improper encoding or escaping of output in ASP.NET Core allows an unauthorized attacker to disclose information over a network. |
| In the Linux kernel, the following vulnerability has been resolved:
net: team: free skb when broadcast has no txable port
bc_transmit() sends the original skb through the last txable port.
When no port is txable, it returns false without consuming the skb, but
team_xmit() still returns NETDEV_TX_OK. AF_PACKET sends then retain the
skb and its socket write-memory charge.
A TX-enabled, link-down port keeps the broadcast transmit op installed.
Carrier can stay up through a second link-up, TX-disabled port or be
forced on by userspace. Free the original skb on the no-port path while
leaving the existing drop accounting intact.
Discovered by Sashiko during review of previous commit. |
| In the Linux kernel, the following vulnerability has been resolved:
regulator: of: fill in supply names in of_regulator_bulk_get_all()
of_regulator_bulk_get_all() returns an array it allocated itself, and
fills in only the consumer of each entry. Every other way of getting a
bulk array has the supply name set, because the caller provides it, and
the core expects it to be there: regulator_bulk_enable() prints it when
a supply fails to enable, so a caller that hands such an array to it
dereferences uninitialised memory on that path.
Copy each name into the array's own allocation, right behind the
entries, so that it shares the array's lifetime and callers still have
nothing extra to free. That also retires the fixed 64 byte stack buffer
the names were assembled in, which is_supply_name() never bounded the
copy against. |
| Stack-based buffer overflow in .NET allows an unauthorized attacker to deny service over a network. |
| In the Linux kernel, the following vulnerability has been resolved:
fs: don't silently unmount busy mounts
The propagate_mount_busy() helper exists to decide whether a synchronous
can suceed. It does that by looking at the copies of the victims in the
mounts its parent propagates to. One of those functions from propagation
hell.
It skips a copy that has child mounts except for the case of a single
child mount covering the copy's root. That algorithm used to be exact.
It isn't anymore.
Originally only childless copies were unmounted during propagation.
And 1064f874abc0 ("mnt: Tuck mounts under others instead of creating
shadow/side mounts.") ended up adding covered mounts to both sides in
one go.
Along came 99b19d16471e ("mnt: In propgate_umount handle visiting mounts
in any order") and that started to unmount more than before:
[1]: A copy of a mount is unmounted when each of its children is either
its overmount or a copy of the victim itself. In the former case
the overmount gets reparented. In the latter case the copy of the
victim will end up being unmounted.
So propagate_umount() ended up being rewritten without adjusting
propagate_mount_busy(). And currently propagate_mount_busy() doesn't
look at copies of the victim that have two or more children nor a single
child that is not the overmount.
So any such copy in [1] is not checked for
references. The result is that a synchronous unmount succeeds, pulling
out a mount that is still in use.
A container can run into this with its own mounts. The host shares a
tree with the container and the host has a mount in that tree:
mount -t tmpfs none /mnt
mount --make-shared /mnt
mount -t tmpfs none /mnt/a
The container's /mnt is a slave copy of the host mount. The container
now takes a detached copy of the tree and mounts that copy beneath its
propagated copy of /mnt/a and keeps the file descriptor to the tree
open:
unshare -m --propagation unchanged
mount --make-rslave /
fd = open_tree(AT_FDCWD, "/mnt", OPEN_TREE_CLONE | AT_RECURSIVE)
move_mount(fd, "", AT_FDCWD, "/mnt/a", MOVE_MOUNT_BENEATH)
grep /mnt /proc/self/mountinfo
775 355 ... /mnt ... master:646
777 775 ... /mnt/a ... master:646
776 777 ... /mnt/a ... master:647
778 777 ... /mnt/a/a ... master:647
The now attached tree (777) is now mounted beneath the container's copy
of /mnt/a. That's also where the host mount sits (as seen from the host
mount namespace). So the propagated copy (776) of the host's mount is on
top of its root and the copy of that (778) is inside it. Now the host
runs:
umount /mnt/a
and it succeeds. The moved tree tree and the copy inside of it are gone
from the container. Now only the propagated copy is left and it got
reparented to where it was before:
grep /mnt /proc/self/mountinfo
775 355 ... /mnt ... master:646
776 775 ... /mnt/a ...
But the file descriptor still refers to the moved tree. So a synchronous
umount detached a mount that is still in use.
Teach propagate_mount_busy() the same checks that propagate_umount()
does. It now checks for a single victim mount without children
for which it can assume that MNT_LOCKED is cleared on every copy like
propagate_umount() does. The copies form chains. IOW, when a copy itself
receives propagation from the victim's parent then the mount at the same
mountpoint inside that copy is a copy as well. Whether a mount receives
propagation from the victim's can simply be checked by walking its
masters. Walk each chain once from its top so we are sure that nested
copies stay linear. Check every reference once of every copy that gets
unmounted. |
| Out-of-bounds read in Remote Desktop Client allows an unauthorized attacker to disclose information over a network. |
| In the Linux kernel, the following vulnerability has been resolved:
namespace: check a recursive bind mount for mount namespace loops
do_loopback() refuses to bind mount the file of a mount namespace that
is as old as the caller's or older because a mount namespace that holds
a mount of its own file, or of an ancestor's, would cause a cycle. But
it only checks the dentry the bind mount starts from. With MS_REC
everything below it is copied and do_loopback() passes
CL_COPY_MNT_NS_FILE so mount namespace files below the source are
copied.
That's fine for a source in the caller's own mount namespace. Every
mount namespace file in there passed the same check when it was
bind-mounted. It isn't fine for a source in another mount namespace.
may_copy_tree() accepts a bind mount of any nsfs or pidfs file no matter
what mount namespace it lives in so that /proc/<pid>/ns/<ns> and pidfds
can be bind mounted. If we create a bind-mount stack of mount namespaces
file descriptors the kernel will copy them irrespective of their
ancestoral relationship to the mount namespace in question:
151 146 0:7 net:[4026531833] /tmp/nrc/y rw - nsfs nsfs rw
152 151 0:7 mnt:[4026532293] /tmp/nrc/y rw - nsfs nsfs rw
Mount 152 is a mount of the mount namespace file of mount namespace
4026532293 inside mount namespace 4026532293. That namespace and every
mount in it are leaked... All it takes is a process in an older mount
namespace that stacks the file and repeating it leaks without limit:
after control: Shmem: 380 kB
child: mount(/proc/self/fd/6, MS_BIND|MS_REC): ok
after cycle: Shmem: 65916 kB
do_move_mount() runs check_for_nsfs_mounts() over a detached tree for
exactly that reason. Do the same for the copy before it is grafted. |
| Out-of-bounds read in Remote Desktop Client allows an unauthorized attacker to disclose information over a network. |
| 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. |
| Missing release of memory after effective lifetime in Windows Layer-2 Bridge Network Driver allows an unauthorized attacker to deny service over an adjacent network. |
| In the Linux kernel, the following vulnerability has been resolved:
netfs: Fix missing alloc tagging of direct mempool allocations
Commit 1d78d56c43ef ("netfs: Fix folio_queue ENOMEM in writeback by
adding a mempool") added a mempool for the folio_queues and made the
request, subrequest and folio_queue allocations distinguish between
writeback and everything else. Writeback is part of memory reclaim
and must not fail due to ENOMEM, so it allocates under GFP_NOFS
through mempool_alloc(), which may dip into the pool's reserve and,
if that runs empty, wait for elements to be returned. The
GFP_KERNEL paths, which can return -ENOMEM to their callers, invoke
the pool's ->alloc() callback directly instead.
The direct call, however, skips the alloc_hooks() wrapper that the
mempool_alloc() macro provides. The pool callbacks, mempool_alloc_slab()
and mempool_kmalloc(), call kmem_cache_alloc_noprof() and kmalloc_noprof()
and rely on current->alloc_tag having been set by the caller. With
CONFIG_MEM_ALLOC_PROFILING_DEBUG=y this leads to
current->alloc_tag not set
WARNING: ./include/linux/alloc_tag.h:161 at __alloc_tagging_slab_alloc_hook
alloc_tag was not set
WARNING: ./include/linux/alloc_tag.h:166 at __alloc_tagging_slab_free_hook
at allocation and free time respectively, as reported when reading
files on a CIFS mount. The allocations are also missing from
/proc/allocinfo.
Wrap the direct ->alloc() invocations in alloc_hooks() with a new
mempool_alloc_noreserve() helper in include/linux/mempool.h, next to
the other alloc_hooks()-wrapped macros such as mempool_alloc(). The
GFP_KERNEL paths keep their failable allocation semantics, they just
get tagged now. |
| 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. |
| Incorrect permission assignment for critical resource in Azure Local allows an authorized attacker to elevate privileges locally. |
| In the Linux kernel, the following vulnerability has been resolved:
net: gso: limit recursive IP-in-IP segmentation
IP-in-IP GSO can re-enter inet_gso_segment() or ipv6_gso_segment()
for each nested IP header. encap_level tracks header bytes, not callback
depth, so a deep chain can exhaust the kernel stack. Making
inet_gso_segment() stackable introduced unbounded IPv4 nesting; IPIP
GSO/TSO later made the path reachable. The IPv6 stackable path was
introduced separately and uses the same guard.
Count IPv4 and IPv6 GSO handler entries in skb_gso_cb, initialized once
per top-level GSO operation and preserved across GRE/UDP context changes.
Use the existing IP_TUNNEL_RECURSION_LIMIT for both handlers. The first
five entries pass, and the sixth returns -EINVAL before dispatching
another GSO callback. |
| Improper check for certificate revocation in Azure Firmware allows an authorized attacker to bypass a security feature locally. |
