At OSDI'14 we published our paper on where we introduce two new
security properties that protect programs against control-flow hijack attacks
enabled by memory corruption vulnerabilities. The design space in this area is
already very cluttered and we use this blog post to highlight the differences
between the individual properties.
Some limitations of this post: we do not discuss actual implementations and
vulnerabilities in these implementations but compare the properties itself
(assuming the best implementation that is theoretically possible). Where we cite
overhead numbers we will refer to strongest prototype implementations of the
properties. This post is also limited in scope, for a complete overview of all
defense mechanisms we refer to the Eternal War in Memory paper and the related work
section of our paper.
Control-Flow Integrity (CFI)
The CFI property ensures that the control-flow of the application stays
within the control-flow graph, as determined ahead of execution. The original
CFI paper, as well as many its follow-ups determine this CFG using static
analysis. This analysis determines, for each indirect call site, a set of
functions that can be called from that call site on any program execution. All
indirect call sites are then instrumented with runtime checks to enforce that
the function being actually called is one of the functions from the
statically-determined set for that call site.
CFI does not prevent memory corruptions, instead it simply does a sanity check
on function pointers before they're actually used. The original CFI paper also
proposed to enforce the integrity of return addresses on the stack through the
use of a shadow stack.
CFI restricts the set of targets to which an attacker might redirect the
control-flow of a program, however, this set necessarily contains all functions
that might be called at a given call site on any program execution, not just
the current one.
To illustrate this, consider the classical callback function pattern. Say we
have a Linux driver that installs a callback for a read() operation on the
device file that it manages. The driver sets one of its functions as the
callback upon its initialization, and the kernel will call this function
whenever a read() operation on the corresponding device file is performed. Now,
the indirect call instruction in the kernel that does this dispatch is the same
for all drivers, so on different executions this instruction might call a
callback from any driver. CFI can, at most, enforce that this instruction
calls any callback for a read() operation from any driver. An attacker who can
overwrite the callback pointer might set it to any function from the above set.
If this call is followed by another indirect function call, an attacker might
even setup a chain of functions to execute.
Unfortunately, there are further limitations in practice. Doing the above
requires precise full-program static analysis: it must look at the entire kernel
and all possible drivers that the kernel can ever load. This is hard in
practice, and we are not aware of such analysis being ever done precisely on
non-trivial systems (not to mention that it's also a theoretically undecidable
problem). In practice, the CFI enforcement mechanism simply will not know which
functions the dispatch instruction in the above example might call, so it will
just enforce that it calls any valid function. This was the case for the
original CFI implementation and, although the follow-up work can reduce this set
in certain situations (such as virtual function calls), any CFI implementation
we are aware of would have cases where this set contains almost every
program function as well.
An potential attack against CFI mechanism was first introduced by Göktas et al.
in the paper that appeared at the Oakland'14 conference. The paper
introduces several new types of gadgets (i.e., fragments of existing application
code that might be reused by an attacker in a way that was not intended by the
application), including call gadgets. Such gadgets start at a regular function
prologue and end with an indirect function call, which enables to transfer
control to the next gadget in the chain. The program might have enough variety
of such linkable gadgets to perform meaningful computation: which is called a
call-oriented programming (analogous to return-oriented programming, which
similarly reuses gadgets that end with a return instruction). Göktas relied on
a combination of call gadgets and other types of gadgets to mount a practical
attack against a coarse-grained CFI implementation. This work can be
generalized, for some programs, to bypasses a fine-grained CFI enforcement
mechanism. We do have both simple and realistic examples of such programs. E.g.,
we where able to construct such an attack for the Perl interpreter (although, to
be fair, we did not automate or weaponize the attack yet).
Code-Pointer Separation (CPS)
CPS prevents an attacker from forging a code pointer out of thin air. At most,
an attacker can cause a program to call a function whose address was previously
taken during current execution. In the above example of a Linux driver, the
dispatch instruction can be potentially subverted to call, e.g., a valid
callback function of a registered driver, but not a callback from a driver that
is not registered on the current run of the kernel or a function that is never
used as a callback.
Causing a program to call wrong function under CPS requires to (i) find an
indirect function call that gets the pointer to be called through another level
of indirection, e.g., struct_ptr->func_ptr() but not just func_ptr();
(ii) use a memory corruption to overwrite the pointer that is used to get the
function pointer indirectly, e.g., struct_ptr, to a location that already
contains a valid function pointer at a proper offset, and (iii) ensure that the
program would not crash if the overwritten pointer, e.g., struct_ptr, is
used after it was overwritten but before the function is called.
CPS ensures precise protection of the return address using safe stack: this
address cannot be overwritten by the attacker in the first place. Moreover, the
safe stack also protects all register spills as well as most of the local
variables. The safe stack, as implemented for CPI and CPS, has much better
performance characteristics than a shadow stack. For details, see our CPI paper.
Code-Pointer Integrity (CPI)
CPI enforces memory safety for all sensitive pointers, which include all code
pointers (just as under CPS) but also all pointers that are used to access
other sensitive pointers indirectly. All such pointers are protected through
runtime memory safety checks. CPI guarantees that an attacker will not be able
to use a memory corruption vulnerability in order to directly change the value
of any sensitive pointer and cause the program to jump to a wrong location as a
result. The only attack vector that remains open is data-based attacks,
preventing those requires full memory safety.
In contrast to CFI, CPI does not execute any sanity checks on the code
pointers before using them, but instead prevents an attacker from corrupting
such pointers in the first place.
Let's compare the security CFI, CPS, and CPI security policies on several
- An attacker attempts to overwrite the return pointer on the stack:
- CFI: the strong CFI, as introduced by Abadi et al., includes optional return
address protection that is based on the shadow stack, which can prevent such
attack. To the best of our knowledge, the shadow stack causes around 5% extra
performance overhead on average, hence Abadi et al. made it optional. Many
modern CFI solutions provide weaker return pointer protection for performance
- CPS & CPI: the return address cannot be overwritten by a corrupt memory
access through the safe stack, which has nearly zero performance overhead
- An attacker attempts to overwrite a code pointer stored in memory using
memory corruption vulnerability:
- CFI: the attacker may redirect the control-flow to any target in the set of
targets permitted for the call site where the pointer is used. In practice,
this set often includes every function in the program (as was the case for
both the original CFI implementation and multiple follow-ups) or wide sets of
functions (e.g., all functions with the same number of arguments). The most
recent CFI implementation reduces the size
of these sets for virtual function calls in C++, but not generic indirect
- CPS & CPI: code pointers cannot be overwritten by corrupt data pointer
- An attacker attempts to overwrite a data pointer that is used to access the
code pointer indirectly (e.g., struct_ptr in struct_ptr->func_ptr()):
- CFI: same limitations as above.
- CPS: the attacker may redirect the control flow to any valid function whose
address has already been taken during current execution and is still stored in
the safe memory.
- CPI: all such indirect pointers are protected and cannot be overwritten by
a corrupt pointer dereferences.
- An attacker attempts to change program data and cause it to take different
execution path (e.g., by changing the decisions made in the if statements):
- CFI: no guarantees
- CPS: no guarantees
- CPI: no guarantees, but some of the pointers are protected.
Comparing CFI and CPS shows that CPS is stronger for scenario B, protecting the
application from attack. For scenario C, CFI may allow a smaller set of targets
then CPS, however, CPS makes it harder to make the program use wrong code
pointer in the first place. The security comparison between CFI and CPS in this
case depends on the program being protected.
Comparing CPS and CPI shows that CPI is stronger than CPS for scenario C.
Comparing CFI and CPI shows that CPI is stronger than CFI in scenarios B and C
due to the memory safety guarantees that it provides for all sensitive
pointers. CFI enforces a static property that was determined at compile time,
whereas CPI enforces actual integrity of all sensitive pointers.
Comparing performance is hard, especially as most CFI implementations are not
open-source. Also, we did not yet have time to evaluate the open-source VTV
implementation. As some indication at performance we list the performance
numbers as reported in the actual papers. Please be aware that these numbers
were measured on different systems under different configurations and serve --
at best -- as an indication of actual performance overhead.
Safe Stack / CPS / CPI:
||C/C++ SPEC2006 CPU
||C/C++ SPEC2006 CPU
||C/C++ SPEC2006 CPU
VTV and IFCC ():
The paper introduces VTV (virtual-table verification) and IFCC (indirect
function-call checks). The former protects virtual function calls, ensuring
that, at any given call site, only the virtual function that corresponds to a
static type of that call site or its subtype might be called. The IFCC
restricts indirect function calls to either any valid function, or any function
with the same number of arguments as the call site.
Note that neither VTV nor IFCC provide return addresses protection.
reports "1% to 8.4% overhead" for VTV only as evaluated on 7 (out
of 19) SPEC2006 CPU benchmarks and Google Chrome. The reported overhead for IFCC
only is around 4%.
Finest-grained CFI implementation
The overhead reported by the original CFI implementation is 21% when
return address protection is enabled. To our knowledge, one of the strongest
CFI implementations around is provided by WIT and has performance
overhead of 10% on average. This includes both function pointers and return