A mechanism is required to binarily patch the running hypervisor with new opcodes that have come about due to primarily security updates.
This document describes the design of the API that would allow us to upload to the hypervisor binary patches.
The document is split in four sections:
The document has gone under various reviews and only covers v1 design.
The end of the document has a section titled Not Yet Done which
outlines ideas and design for the future version of this work.
The mechanism needs to be flexible to patch the hypervisor in multiple ways and be as simple as possible. The compiled code is contiguous in memory with no gaps - so we have no luxury of 'moving' existing code and must either insert a trampoline to the new code to be executed - or only modify in-place the code if there is sufficient space. The placement of new code has to be done by hypervisor and the virtual address for the new code is allocated dynamically.
This implies that the hypervisor must compute the new offsets when splicing in the new trampoline code. Where the trampoline is added (inside the function we are patching or just the callers?) is also important.
To lessen the amount of code in hypervisor, the consumer of the API is responsible for identifying which mechanism to employ and how many locations to patch. Combinations of modifying in-place code, adding trampoline, etc has to be supported. The API should allow read/write any memory within the hypervisor virtual address space.
We must also have a mechanism to query what has been applied and a mechanism to revert it if needed.
The expected workflows of higher-level tools that manage multiple patches on production machines would be:
The first mechanism to patch that comes in mind is in-place replacement. That is replace the affected code with new code. Unfortunately the x86 ISA is variable size which places limits on how much space we have available to replace the instructions. That is not a problem if the change is smaller than the original opcode and we can fill it with nops. Problems will appear if the replacement code is longer.
The second mechanism is by ti replace the call or jump to the old function with the address of the new function.
A third mechanism is to add a jump to the new function at the
start of the old function. N.B. The Xen hypervisor implements the third
mechanism. See Trampoline (e9 opcode) section for more details.
As example we will assume the hypervisor does not have XSA-132 (see domctl/sysctl: don't leak hypervisor stack to toolstacks) and we would like to binary patch the hypervisor with it. The original code looks as so:
48 89 e0 mov %rsp,%rax
48 25 00 80 ff ff and $0xffffffffffff8000,%rax
while the new patched hypervisor would be:
48 c7 45 b8 00 00 00 00 movq $0x0,-0x48(%rbp)
48 c7 45 c0 00 00 00 00 movq $0x0,-0x40(%rbp)
48 c7 45 c8 00 00 00 00 movq $0x0,-0x38(%rbp)
48 89 e0 mov %rsp,%rax
48 25 00 80 ff ff and $0xffffffffffff8000,%rax
This is inside the arch_do_domctl. This new change adds 21 extra bytes of code which alters all the offsets inside the function. To alter these offsets and add the extra 21 bytes of code we might not have enough space in .text to squeeze this in.
As such we could simplify this problem by only patching the site which calls arch_do_domctl:
do_domctl:
e8 4b b1 05 00 callq ffff82d08015fbb9 <arch_do_domctl>
with a new address for where the new arch_do_domctl would be (this
area would be allocated dynamically).
Astute readers will wonder what we need to do if we were to patch do_domctl
- which is not called directly by hypervisor but on behalf of the guests via
the compat_hypercall_table and hypercall_table. Patching the offset in
hypercall_table for do_domctl:
ffff82d08024d490: 79 30
ffff82d08024d492: 10 80 d0 82 ff ff
with the new address where the new do_domctl is possible. The other
place where it is used is in hvm_hypercall64_table which would need
to be patched in a similar way. This would require an in-place splicing
of the new virtual address of arch_do_domctl.
In summary this example patched the callee of the affected function by
However it has drawbacks - the safety checks which have to make sure the function is not on the stack - must also check every caller. For some patches this could mean - if there were an sufficient large amount of callers - that we would never be able to apply the update.
Having the patching done at predetermined instances where the stacks are not deep mostly solves this problem.
An alternative mechanism exists where we can insert a trampoline in the existing function to be patched to jump directly to the new code. This lessens the locations to be patched to one but it puts pressure on the CPU branching logic (I-cache, but it is just one unconditional jump).
For this example we will assume that the hypervisor has not been compiled with
XSA-125 (see
pre-fill structures for certain HYPERVISOR_xen_version sub-ops)
which mem-sets an structure in xen_version hypercall. This function is not
called anywhere in the hypervisor (it is called by the guest) but
referenced in the compat_hypercall_table and hypercall_table (and
indirectly called from that). Patching the offset in hypercall_table for the
old do_xen_version:
ffff82d08024b270 <hypercall_table>:
...
ffff82d08024b2f8: 9e 2f 11 80 d0 82 ff ff
with the new address where the new do_xen_version is possible. The other
place where it is used is in hvm_hypercall64_table which would need
to be patched in a similar way. This would require an in-place splicing
of the new virtual address of do_xen_version.
An alternative solution would be to patch insert a trampoline in the
old do_xen_version function to directly jump to the new do_xen_version:
ffff82d080112f9e do_xen_version:
ffff82d080112f9e: 48 c7 c0 da ff ff ff mov $0xffffffffffffffda,%rax
ffff82d080112fa5: 83 ff 09 cmp $0x9,%edi
ffff82d080112fa8: 0f 87 24 05 00 00 ja ffff82d0801134d2 ; do_xen_version+0x534
with:
ffff82d080112f9e do_xen_version:
ffff82d080112f9e: e9 XX YY ZZ QQ jmpq [new do_xen_version]
which would lessen the amount of patching to just one location.
In summary this example patched the affected function to jump to the new replacement function which required:
The disadvantage of this are that the unconditional jump will consume a small I-cache penalty. However the simplicity of the patching and higher chance of passing safety checks make this a worthwhile option.
This patching has a similar drawback as inline patching - the safety checks have to make sure the function is not on the stack. However since we are replacing at a higher level (a full function as opposed to various offsets within functions) the checks are simpler.
Having the patching done at predetermined instances where the stacks are not deep mostly solves this problem as well.
With this method we can re-write the hypervisor - and as such we MUST be diligent in only allowing certain guests to perform this operation.
Furthermore with SecureBoot or tboot, we MUST also verify the signature of the payload to be certain it came from a trusted source and integrity was intact.
As such the hypercall MUST support an XSM policy to limit what the guest is allowed to invoke. If the system is booted with signature checking the signature checking will be enforced.
The payload MUST contain enough data to allow us to apply the update and also safely reverse it. As such we MUST know:
This binary format can be constructed using an custom binary format but there are severe disadvantages of it:
As such having the payload in an ELF file is the sensible way. We would be carrying the various sets of structures (and data) in the ELF sections under different names and with definitions.
Note that every structure has padding. This is added so that the hypervisor can re-use those fields as it sees fit.
Earlier design attempted to ineptly explain the relations of the ELF sections to each other without using proper ELF mechanism (sh_info, sh_link, data structures using Elf types, etc). This design will explain the structures and how they are used together and not dig in the ELF format - except mention that the section names should match the structure names.
The Xen Live Patch payload is a relocatable ELF binary. A typical binary would have:
It may also have some architecture-specific sections. For example:
The Xen Live Patch core code loads the payload as a standard ELF binary, relocates it and handles the architecture-specifc sections as needed. This process is much like what the Linux kernel module loader does.
The payload contains at least three sections:
.livepatch.funcs - which is an array of livepatch_func structures..livepatch.depends - which is an ELF Note that describes what the payload
depends on. MUST have one..note.gnu.build-id - the build-id of this payload. MUST have one.The .livepatch.funcs contains an array of livepatch_func structures
which describe the functions to be patched:
struct livepatch_func {
const char *name;
void *new_addr;
void *old_addr;
uint32_t new_size;
uint32_t old_size;
uint8_t version;
uint8_t opaque[31];
};
The size of the structure is 64 bytes on 64-bit hypervisors. It will be 52 on 32-bit hypervisors.
name is the symbol name of the old function. Only used if old_addr is
zero, otherwise will be used during dynamic linking (when hypervisor loads
the payload).old_addr is the address of the function to be patched and is filled in at
payload generation time if hypervisor function address is known. If unknown,
the value MUST be zero and the hypervisor will attempt to resolve the
address.new_addr can either have a non-zero value or be zero.
old_addr location
new_size bytes.old_size contains the sizes of the respective old_addr function in
bytes. The value of old_size MUST not be zero.new_size depends on what new_addr contains:
new_addr contains an non-zero value, then new_size has the size of
the new function (which will replace the one at old_addr) in bytes.new_addr is zero then new_size determines how many
instruction bytes to NOP (up to opaque size modulo smallest platform
instruction - 1 byte x86 and 4 bytes on ARM).version is to be one.opaque MUST be zero.The size of the livepatch_func array is determined from the ELF section
size.
When applying the patch the hypervisor iterates over each livepatch_func
structure and the core code inserts a trampoline at old_addr to new_addr.
The new_addr is altered when the ELF payload is loaded.
When reverting a patch, the hypervisor iterates over each livepatch_func
and the core code copies the data from the undo buffer (private internal copy)
to old_addr.
It optionally may contain the address of functions to be called right before being applied and after being reverted:
.livepatch.hooks.load - an array of function pointers..livepatch.hooks.unload - an array of function pointers.A simple example of what a payload file can be:
/* MUST be in sync with hypervisor. */
struct livepatch_func {
const char *name;
void *new_addr;
void *old_addr;
uint32_t new_size;
uint32_t old_size;
uint8_t version;
uint8_t pad[31];
};
/* Our replacement function for xen_extra_version. */
const char *xen_hello_world(void)
{
return "Hello World";
}
static unsigned char patch_this_fnc[] = "xen_extra_version";
struct livepatch_func livepatch_hello_world = {
.version = LIVEPATCH_PAYLOAD_VERSION,
.name = patch_this_fnc,
.new_addr = xen_hello_world,
.old_addr = (void *)0xffff82d08013963c, /* Extracted from xen-syms. */
.new_size = 13, /* To be be computed by scripts. */
.old_size = 13, /* -----------""--------------- */
} __attribute__((__section__(".livepatch.funcs")));
Code must be compiled with -fPIC.
This section contains an array of function pointers to be executed before payload is being applied (.livepatch.funcs) or after reverting the payload. This is useful to prepare data structures that need to be modified patching.
Each entry in this array is eight bytes.
The type definition of the function are as follow:
typedef void (*livepatch_loadcall_t)(void);
typedef void (*livepatch_unloadcall_t)(void);
To support dependencies checking and safe loading (to load the appropiate payload against the right hypervisor) there is a need to embbed an build-id dependency.
This is done by the payload containing an section .livepatch.depends
which follows the format of an ELF Note. The contents of this
(name, and description) are specific to the linker utilized to
build the hypevisor and payload.
If GNU linker is used then the name is GNU and the description
is a NT_GNU_BUILD_ID type ID. The description can be an SHA1
checksum, MD5 checksum or any unique value.
The size of these structures varies with the --build-id linker option.
We will employ the sub operations of the system management hypercall (sysctl). There are to be four sub-operations:
Most of the actions are asynchronous therefore the caller is responsible to verify that it has been applied properly by retrieving the summary of it and verifying that there are no error codes associated with the payload.
We MUST make some of them asynchronous due to the nature of patching it requires every physical CPU to be lock-step with each other. The patching mechanism while an implementation detail, is not an short operation and as such the design MUST assume it will be an long-running operation.
The sub-operations will spell out how preemption is to be handled (if at all).
Furthermore it is possible to have multiple different payloads for the same function. As such an unique name per payload has to be visible to allow proper manipulation.
The hypercall is part of the xen_sysctl. The top level structure contains
one uint32_t to determine the sub-operations and one padding field which
MUST always be zero.
struct xen_sysctl_livepatch_op {
uint32_t cmd; /* IN: XEN_SYSCTL_LIVEPATCH_*. */
uint32_t pad; /* IN: Always zero. */
union {
... see below ...
} u;
};
while the rest of hypercall specific structures are part of the this structure.
Most of the hypercalls employ an shared structure called struct xen_livepatch_name
which contains:
name - pointer where the string for the name is located.size - the size of the stringpad - padding - to be zero.The structure is as follow:
/*
* Uniquely identifies the payload. Should be human readable.
* Includes the NUL terminator
*/
#define XEN_LIVEPATCH_NAME_SIZE 128
struct xen_livepatch_name {
XEN_GUEST_HANDLE_64(char) name; /* IN, pointer to name. */
uint16_t size; /* IN, size of name. May be upto
XEN_LIVEPATCH_NAME_SIZE. */
uint16_t pad[3]; /* IN: MUST be zero. */
};
Upload a payload to the hypervisor. The payload is verified against basic checks and if there are any issues the proper return code will be returned. The payload is not applied at this time - that is controlled by XEN_SYSCTL_LIVEPATCH_ACTION.
The caller provides:
struct xen_livepatch_name called name which has the unique name.size the size of the ELF payload (in bytes).payload the virtual address of where the ELF payload is.The name could be an UUID that stays fixed forever for a given
payload. It can be embedded into the ELF payload at creation time
and extracted by tools.
The return value is zero if the payload was succesfully uploaded.
Otherwise an -XEN_EXX return value is provided. Duplicate name are not supported.
The payload is the ELF payload as mentioned in the Payload format section.
The structure is as follow:
struct xen_sysctl_livepatch_upload {
xen_livepatch_name_t name; /* IN, name of the patch. */
uint64_t size; /* IN, size of the ELF file. */
XEN_GUEST_HANDLE_64(uint8) payload; /* IN: ELF file. */
};
Retrieve an status of an specific payload. This caller provides:
struct xen_livepatch_name called name which has the unique name.struct xen_livepatch_status structure. The member values will
be over-written upon completion.Upon completion the struct xen_livepatch_status is updated.
status - indicates the current status of the payload:
rc - -XEN_EXX type errors encountered while performing the last
LIVEPATCH_ACTION_* operation. The normal values can be zero or -XEN_EAGAIN which
respectively mean: success or operation in progress. Other values
imply an error occurred. If there is an error in rc, status will NOT
have changed.The return value of the hypercall is zero on success and -XEN_EXX on failure.
(Note that the rc value can be different from the return value, as in
rc=-XEN_EAGAIN and return value can be 0).
For example, supposing there is an payload:
status: LIVEPATCH_STATUS_CHECKED
rc: 0
We apply an action - LIVEPATCH_ACTION_REVERT - to revert it (which won't work as we have not even applied it. Afterwards we will have:
status: LIVEPATCH_STATUS_CHECKED
rc: -XEN_EINVAL
It has failed but it remains loaded.
This operation is synchronous and does not require preemption.
The structure is as follow:
struct xen_livepatch_status {
#define LIVEPATCH_STATUS_CHECKED 1
#define LIVEPATCH_STATUS_APPLIED 2
uint32_t state; /* OUT: LIVEPATCH_STATE_*. */
int32_t rc; /* OUT: 0 if no error, otherwise -XEN_EXX. */
};
struct xen_sysctl_livepatch_get {
xen_livepatch_name_t name; /* IN, the name of the payload. */
xen_livepatch_status_t status; /* IN/OUT: status of the payload. */
};
Retrieve an array of abbreviated status and names of payloads that are loaded in the hypervisor.
The caller provides:
version. Version of the payload. Caller should re-use the field provided by
the hypervisor. If the value differs the data is stale.idx Index iterator. The index into the hypervisor's payload count. It is
recommended that on first invocation zero be used so that nr (which the
hypervisor will update with the remaining payload count) be provided.
Also the hypervisor will provide version with the most current value.nr The max number of entries to populate. Can be zero which will result
in the hypercall being a probing one and return the number of payloads
(and update the version).pad - MUST be zero.status Virtual address of where to write struct xen_livepatch_status
structures. Caller MUST allocate up to nr of them.name - Virtual address of where to write the unique name of the payload.
Caller MUST allocate up to nr of them. Each MUST be of
XEN_LIVEPATCH_NAME_SIZE size. Note that XEN_LIVEPATCH_NAME_SIZE includes
the NUL terminator.len - Virtual address of where to write the length of each unique name
of the payload. Caller MUST allocate up to nr of them. Each MUST be
of sizeof(uint32_t) (4 bytes).If the hypercall returns an positive number, it is the number (upto nr
provided to the hypercall) of the payloads returned, along with nr updated
with the number of remaining payloads, version updated (it may be the same
across hypercalls - if it varies the data is stale and further calls could
fail). The status, name, and len are updated at their designed index
value (idx) with the returned value of data.
If the hypercall returns -XEN_E2BIG the nr is too big and should be
lowered.
If the hypercall returns an zero value there are no more payloads.
Note that due to the asynchronous nature of hypercalls the control domain might
have added or removed a number of payloads making this information stale. It is
the responsibility of the toolstack to use the version field to check
between each invocation. if the version differs it should discard the stale
data and start from scratch. It is OK for the toolstack to use the new
version field.
The struct xen_livepatch_status structure contains an status of payload which includes:
status - indicates the current status of the payload:
rc - -XEN_EXX type errors encountered while performing the last
LIVEPATCH_ACTION_* operation. The normal values can be zero or -XEN_EAGAIN which
respectively mean: success or operation in progress. Other values
imply an error occurred. If there is an error in rc, status will NOT
have changed.The structure is as follow:
struct xen_sysctl_livepatch_list {
uint32_t version; /* OUT: Hypervisor stamps value.
If varies between calls, we are
getting stale data. */
uint32_t idx; /* IN: Index into hypervisor list. */
uint32_t nr; /* IN: How many status, names, and len
should be filled out. Can be zero to get
amount of payloads and version.
OUT: How many payloads left. */
uint32_t pad; /* IN: Must be zero. */
XEN_GUEST_HANDLE_64(xen_livepatch_status_t) status; /* OUT. Must have enough
space allocate for nr of them. */
XEN_GUEST_HANDLE_64(char) id; /* OUT: Array of names. Each member
MUST XEN_LIVEPATCH_NAME_SIZE in size.
Must have nr of them. */
XEN_GUEST_HANDLE_64(uint32) len; /* OUT: Array of lengths of name's.
Must have nr of them. */
};
Perform an operation on the payload structure referenced by the name field.
The operation request is asynchronous and the status should be retrieved
by using either XEN_SYSCTL_LIVEPATCH_GET or XEN_SYSCTL_LIVEPATCH_LIST hypercall.
The caller provides:
struct xen_livepatch_name name containing the unique name.cmd The command requested:
name will result in failure unless
XEN_SYSCTL_LIVEPATCH_UPLOAD hypercall is perfomed with same name.rc in xen_livepatch_status
retrieved via XEN_SYSCTL_LIVEPATCH_GET will be -XEN_EBUSY.rc in xen_livepatch_status
retrieved via XEN_SYSCTL_LIVEPATCH_GET will be -XEN_EBUSY.rc in xen_livepatch_status retrieved via XEN_SYSCTL_LIVEPATCH_GET
will be -XEN_EBUSY.time The upper bound of time (ns) the cmd should take. Zero means to use
the hypervisor default. If within the time the operation does not succeed
the operation would go in error state.pad - MUST be zero.The return value will be zero unless the provided fields are incorrect.
The structure is as follow:
#define LIVEPATCH_ACTION_UNLOAD 1
#define LIVEPATCH_ACTION_REVERT 2
#define LIVEPATCH_ACTION_APPLY 3
#define LIVEPATCH_ACTION_REPLACE 4
struct xen_sysctl_livepatch_action {
xen_livepatch_name_t name; /* IN, name of the patch. */
uint32_t cmd; /* IN: LIVEPATCH_ACTION_* */
uint32_t time; /* IN: If zero then uses */
/* hypervisor default. */
/* Or upper bound of time (ns) */
/* for operation to take. */
};
There is a strict ordering state of what the commands can be. The LIVEPATCH_ACTION prefix has been dropped to easy reading and does not include the LIVEPATCH_STATES:
/->\
\ /
UNLOAD <--- CHECK ---> REPLACE|APPLY --> REVERT --\
\ |
\-------------------<-------------/
Note that:
The state transition table of valid states and action states:
+---------+---------+--------------------------------+-------+--------+
| ACTION | Current | Result | Next STATE: |
| ACTION | STATE | |CHECKED|APPLIED |
+---------+----------+-------------------------------+-------+--------+
| UNLOAD | CHECKED | Unload payload. Always works. | | |
| | | No next states. | | |
+---------+---------+--------------------------------+-------+--------+
| APPLY | CHECKED | Apply payload (success). | | x |
+---------+---------+--------------------------------+-------+--------+
| APPLY | CHECKED | Apply payload (error|timeout) | x | |
+---------+---------+--------------------------------+-------+--------+
| REPLACE | CHECKED | Revert payloads and apply new | | x |
| | | payload with success. | | |
+---------+---------+--------------------------------+-------+--------+
| REPLACE | CHECKED | Revert payloads and apply new | x | |
| | | payload with error. | | |
+---------+---------+--------------------------------+-------+--------+
| REVERT | APPLIED | Revert payload (success). | x | |
+---------+---------+--------------------------------+-------+--------+
| REVERT | APPLIED | Revert payload (error|timeout) | | x |
+---------+---------+--------------------------------+-------+--------+
All the other state transitions are invalid.
The normal sequence of events is to:
->rc. If -XEN_EAGAIN spin. If zero go to next step.->rc. If in -XEN_EAGAIN spin. If zero exit with success.Implementation quirks should not be discussed in a design document.
However these observations can provide aid when developing against this document.
Alternative assembler is a mechanism to use different instructions depending
on what the CPU supports. This is done by providing multiple streams of code
that can be patched in - or if the CPU does not support it - padded with
nop operations. The alternative assembler macros cause the compiler to
expand the code to place a most generic code in place - emit a special
ELF .section header to tag this location. During run-time the hypervisor
can leave the areas alone or patch them with an better suited opcodes.
Note that patching functions that copy to or from guest memory requires to support alternative support. For example this can be due to SMAP (specifically stac and clac operations) which is enabled on Broadwell and later architectures. It may be related to other alternative instructions.
During the discussion on the design two candidates bubbled where the call stack for each CPU would be deterministic. This would minimize the chance of the patch not being applied due to safety checks failing. Safety checks such as not patching code which is on the stack - which can lead to corruption.
The hypervisor's time rendezvous code runs synchronously across all CPUs
every second. Using the stop_machine to patch can stall the time rendezvous
code and result in NMI. As such having the patching be done at the tail
of rendezvous code should avoid this problem.
However the entrance point for that code is do_softirq ->
timer_softirq_action -> time_calibration which ends up calling
on_selected_cpus on remote CPUs.
The remote CPUs receive CALL_FUNCTION_VECTOR IPI and execute the desired function.
Before we call VMXResume we check whether any soft IRQs need to be executed. This is a good spot because all Xen stacks are effectively empty at that point.
To randezvous all the CPUs an barrier with an maximum timeout (which could be adjusted), combined with forcing all other CPUs through the hypervisor with IPIs, can be utilized to execute lockstep instructions on all CPUs.
The approach is similar in concept to stop_machine and the time rendezvous
but is time-bound. However the local CPU stack is much shorter and
a lot more deterministic.
This is implemented in the Xen hypervisor.
Hotpatch generation often requires support for compiling the target
with -ffunction-sections / -fdata-sections. Changes would have to
be done to the linker scripts to support this.
The design of that is not discussed in this design.
This is implemented in a seperate tool which lives in a seperate GIT repo.
Currently it resides at git://xenbits.xen.org/livepatch-build-tools.git
We may need support for adapting or augmenting exception tables if patching such code. Hotpatches may need to bring their own small exception tables (similar to how Linux modules support this).
If supporting hotpatches that introduce additional exception-locations is not important, one could also change the exception table in-place and reorder it afterwards.
As found almost every patch (XSA) to a non-trivial function requires additional entries in the exception table and/or the bug frames.
This is implemented in the Xen hypervisor.
The patching might require strings to be updated as well. As such we must be also able to patch the strings as needed. This sounds simple - but the compiler has a habit of coalescing strings that are the same - which means if we in-place alter the strings - other users will be inadvertently affected as well.
This is also where pointers to functions live - and we may need to patch this as well. And switch-style jump tables.
To guard against that we must be prepared to do patching similar to trampoline patching or in-line depending on the flavour. If we can do in-line patching we would need to:
.rodata to be writeable..rodata to be read-only.If are doing trampoline patching we would need to:
The trampoline patching is implemented in the Xen hypervisor.
In place patching writable data is not suitable as it is unclear what should be done depending on the current state of data. As such it should not be attempted.
However, functions which are being patched can bring in changes to strings (.data or .rodata section changes), or even to .bss sections.
As such the ELF payload can introduce new .rodata, .bss, and .data sections. Patching in the new function will end up also patching in the new .rodata section and the new function will reference the new string in the new .rodata section.
This is implemented in the Xen hypervisor.
Only the privileged domain should be allowed to do this operation.
Live patch patches interdependencies are tricky.
There are the ways this can be addressed: * A single large patch that subsumes and replaces all previous ones. Over the life-time of patching the hypervisor this large patch grows to accumulate all the code changes. * Hotpatch stack - where an mechanism exists that loads the hotpatches in the same order they were built in. We would need an build-id of the hypevisor to make sure the hot-patches are build against the correct build. * Payload containing the old code to check against that. That allows the hotpatches to be loaded indepedently (if they don't overlap) - or if the old code also containst previously patched code - even if they overlap.
The disadvantage of the first large patch is that it can grow over time and not provide an bisection mechanism to identify faulty patches.
The hot-patch stack puts stricts requirements on the order of the patches being loaded and requires an hypervisor build-id to match against.
The old code allows much more flexibility and an additional guard, but is more complex to implement.
The second option which requires an build-id of the hypervisor is implemented in the Xen hypervisor.
Specifically each payload has two build-id ELF notes: * The build-id of the payload itself (generated via --build-id). * The build-id of the payload it depends on (extracted from the the previous payload or hypervisor during build time).
This means that the very first payload depends on the hypervisor build-id.
This is for further development of live patching.
The implementation must also have a mechanism for (in no particular order):
symbol or symbol+offset).new_size is zero.This problem is related to hotpatch construction and potentially has influence on the design of the hotpatching infrastructure in Xen.
For example:
We have file1.c with functions f1 and f2 (in that order). f2 contains a BUG() (or WARN()) macro and at that point embeds the source line number into the generated code for f2.
Now we want to hotpatch f1 and the hotpatch source-code patch adds 2 lines to f1 and as a consequence shifts out f2 by two lines. The newly constructed file1.o will now contain differences in both binary functions f1 (because we actually changed it with the applied patch) and f2 (because the contained BUG macro embeds the new line number).
Without additional information, an algorithm comparing file1.o before and after hotpatch application will determine both functions to be changed and will have to include both into the binary hotpatch.
Options:
Transform source code patches for hotpatches to be line-neutral for each chunk. This can be done in almost all cases with either reformatting of the source code or by introducing artificial preprocessor "#line n" directives to adjust for the introduced differences.
This approach is low-tech and simple. Potentially generated backtraces and existing debug information refers to the original build and does not reflect hotpatching state except for actually hotpatched functions but should be mostly correct.
Ignoring the problem and living with artificially large hotpatches that unnecessarily patch many functions.
This approach might lead to some very large hotpatches depending on content of specific source file. It may also trigger pulling in functions into the hotpatch that cannot reasonable be hotpatched due to limitations of a hotpatching framework (init-sections, parts of the hotpatching framework itself, ...) and may thereby prevent us from patching a specific problem.
The decision between 1. and 2. can be made on a patch--by-patch basis.
Introducing an indirection table for storing line numbers and treating that specially for binary diffing. Linux may follow this approach.
We might either use this indirection table for runtime use and patch that with each hotpatch (similarly to exception tables) or we might purely use it when building hotpatches to ignore functions that only differ at exactly the location where a line-number is embedded.
For BUG(), WARN(), etc., the line number is embedded into the bug frame, not the function itself.
Similar considerations are true to a lesser extent for __FILE__, but it could be argued that file renaming should be done outside of hotpatches.
The signature checking requires that the layout of the data in memory MUST be same for signature to be verified. This means that the payload data layout in ELF format MUST match what the hypervisor would be expecting such that it can properly do signature verification.
The signature is based on the all of the payloads continuously laid out
in memory. The signature is to be appended at the end of the ELF payload
prefixed with the string '~Module signature appended~\n', followed by
an signature header then followed by the signature, key identifier, and signers
name.
Specifically the signature header would be:
#define PKEY_ALGO_DSA 0
#define PKEY_ALGO_RSA 1
#define PKEY_ID_PGP 0 /* OpenPGP generated key ID */
#define PKEY_ID_X509 1 /* X.509 arbitrary subjectKeyIdentifier */
#define HASH_ALGO_MD4 0
#define HASH_ALGO_MD5 1
#define HASH_ALGO_SHA1 2
#define HASH_ALGO_RIPE_MD_160 3
#define HASH_ALGO_SHA256 4
#define HASH_ALGO_SHA384 5
#define HASH_ALGO_SHA512 6
#define HASH_ALGO_SHA224 7
#define HASH_ALGO_RIPE_MD_128 8
#define HASH_ALGO_RIPE_MD_256 9
#define HASH_ALGO_RIPE_MD_320 10
#define HASH_ALGO_WP_256 11
#define HASH_ALGO_WP_384 12
#define HASH_ALGO_WP_512 13
#define HASH_ALGO_TGR_128 14
#define HASH_ALGO_TGR_160 15
#define HASH_ALGO_TGR_192 16
struct elf_payload_signature {
u8 algo; /* Public-key crypto algorithm PKEY_ALGO_*. */
u8 hash; /* Digest algorithm: HASH_ALGO_*. */
u8 id_type; /* Key identifier type PKEY_ID*. */
u8 signer_len; /* Length of signer's name */
u8 key_id_len; /* Length of key identifier */
u8 __pad[3];
__be32 sig_len; /* Length of signature data */
};
(Note that this has been borrowed from Linux module signature code.).
In place patching writable data is not suitable as it is unclear what should be done depending on the current state of data. As such it should not be attempted.
That said we should provide hook functions so that the existing data can be changed during payload application.
To guarantee safety we disallow re-applying an payload after it has been reverted. This is because we cannot guarantee that the state of .bss and .data to be exactly as it was during loading. Hence the administrator MUST unload the payload and upload it again to apply it.
There is an exception to this: if the payload only has .livepatch.funcs; and the .data or .bss sections are of zero length.
The hypervisor should verify that the in-place patching would fit within the code or data.
The e9 opcode used for jmpq uses a 32-bit signed displacement. That means we are limited to up to 2GB of virtual address to place the new code from the old code. That should not be a problem since Xen hypervisor has a very small footprint.
However if we need - we can always add two trampolines. One at the 2GB limit that calls the next trampoline.
Please note there is a small limitation for trampolines in
function entries: The target function (+ trailing padding) must be able
to accomodate the trampoline. On x86 with +-2 GB relative jumps,
this means 5 bytes are required which means that old_size MUST be
at least five bytes if patching in trampoline.
Depending on compiler settings, there are several functions in Xen that are smaller (without inter-function padding).
readelf -sW xen-syms | grep " FUNC " | \
awk '{ if ($3 < 5) print $3, $4, $5, $8 }'
...
3 FUNC LOCAL wbinvd_ipi
3 FUNC LOCAL shadow_l1_index
...
A compile-time check for, e.g., a minimum alignment of functions or a runtime check that verifies symbol size (+ padding to next symbols) for that in the hypervisor is advised.
The tool for generating payloads currently does perform a compile-time check to ensure that the function to be replaced is large enough.
The unconditional branch instruction (for the encoding see the
DDI 0406C.c and DDI 0487A.j Architecture Reference Manual's).
with proper offset is used for an unconditional branch to the new code.
This means that that old_size MUST be at least four bytes if patching
in trampoline.
The instruction offset is limited on ARM32 to +/- 32MB to displacement and on ARM64 to +/- 128MB displacement.
The new code is placed in the 8M - 10M virtual address space while the Xen code is in 2M - 4M. That gives us enough space.
The hypervisor also checks the displacement during loading of the payload.