Return Oriented Programming (ROP) has become the exploitation technique of choice for modern memory-safety vulnerability attacks. Recently, there have been multiple attempts at defenses to prevent ROP attacks. In this paper, we introduce three new attack methods that break many existing ROP defenses. Then we show how to break kBouncer and ROPecker, two recent low-overhead defenses that can be applied to legacy software on existing hardware. We examine several recent ROP attacks seen in the wild and demonstrate that our techniques successfully cloak them so they are not detected by these defenses. Our attacks apply to many CFI-based defenses which we argue are weaker than previously thought. Future defenses will need to take our attacks into account. 1

Code-reuse attacks based on return oriented program-ming are among the most popular exploitation tech-niques used by attackers today. Few practical defenses are able to stop such attacks on arbitrary binaries with-out access to source code. A notable exception are the techniques that employ new hardware, such as Intel’s Last Branch Record (LBR) registers, to track all indirect branches and raise an alert when a sensitive system call is reached by means of too many indirect branches to short gadgets—under the assumption that such gadget chains would be indicative of a ROP attack. In this paper, we evaluate the implications. What is “too many ” and how short is “short”? Getting the thresholds wrong has seri-ous consequences. In this paper, we show by means of an attack on Internet Explorer that while current defenses based on these techniques raise the bar for exploitation, they can be bypassed. Conversely, tuning the thresholds to make the defenses more aggressive, may flag legit-imate program behavior as an attack. We analyze the problem in detail and show that determining the right val-ues is difficult. 1

Abstract-Code-reuse attacks such as return-oriented programming (ROP) pose a severe threat to modern software. Designing practical and effective defenses against code-reuse attacks is highly challenging. One line of defense builds upon fine-grained code diversification to prevent the adversary from constructing a reliable code-reuse attack. However, all solutions proposed so far are either vulnerable to memory disclosure or are impractical for deployment on commodity systems. In this paper, we address the deficiencies of existing solutions and present the first practical, fine-grained code randomization defense, called Readactor, resilient to both static and dynamic ROP attacks. We distinguish between direct memory disclosure, where the attacker reads code pages, and indirect memory disclosure, where attackers use code pointers on data pages to infer the code layout without reading code pages. Unlike previous work, Readactor resists both types of memory disclosure. Moreover, our technique protects both statically and dynamically generated code. We use a new compiler-based code generation paradigm that uses hardware features provided by modern CPUs to enable execute-only memory and hide code pointers from leakage to the adversary. Finally, our extensive evaluation shows that our approach is practical-we protect the entire Google Chromium browser and its V8 JIT compiler-and efficient with an average SPEC CPU2006 performance overhead of only 6.4%.

Abstract-A new binary software randomization and ControlFlow Integrity (CFI) enforcement system is presented, which is the first to efficiently resist code-reuse attacks launched by informed adversaries who possess full knowledge of the inmemory code layout of victim programs. The defense mitigates a recent wave of implementation disclosure attacks, by which adversaries can exfiltrate in-memory code details in order to prepare code-reuse attacks (e.g., Return-Oriented Programming (ROP) attacks) that bypass fine-grained randomization defenses. Such implementation-aware attacks defeat traditional fine-grained randomization by undermining its assumption that the randomized locations of abusable code gadgets remain secret. Opaque CFI (O-CFI) overcomes this weakness through a novel combination of fine-grained code-randomization and coarsegrained control-flow integrity checking. It conceals the graph of hijackable control-flow edges even from attackers who can view the complete stack, heap, and binary code of the victim process. For maximal efficiency, the integrity checks are implemented using instructions that will soon be hardware-accelerated on commodity x86-x64 processors. The approach is highly practical since it does not require a modified compiler and can protect legacy binaries without access to source code. Experiments using our fully functional prototype implementation show that O-CFI provides significant probabilistic protection against ROP attacks launched by adversaries with complete code layout knowledge, and exhibits only 4.7% mean performance overhead on current hardware (with further overhead reductions to follow on forthcoming Intel processors). I. MOTIVATION Code-reuse attacks (cf., Permission to freely reproduce all or part of this paper for noncommercial purposes is granted provided that copies bear this notice and the full citation on the first page. Reproduction for commercial purposes is strictly prohibited without the prior written consent of the Internet Society, the first-named author (for reproduction of an entire paper only), and the author's employer if the paper was prepared within the scope of employment. This has motivated copious work on defenses against codereuse threats. Prior defenses can generally be categorized into: CFI [1] and artificial software diversity CFI restricts all of a program's runtime control-flows to a graph of whitelisted control-flow edges. Usually the graph is derived from the semantics of the program source code or a conservative disassembly of its binary code. As a result, CFIprotected programs reject control-flow hijacks that attempt to traverse edges not supported by the original program's semantics. Fine-grained CFI monitors indirect control-flows precisely; for example, function callees must return to their exact callers. Although such precision provides the highest security, it also tends to incur high performance overheads (e.g., 21% for precise caller-callee return-matching [1]). Because this overhead is often too high for industry adoption, researchers have proposed many optimized, coarser-grained variants of CFI. Coarse-grained CFI trades some security for better performance by reducing the precision of the checks. For example, functions must return to valid call sites (but not necessarily to the particular site that invoked the callee). Unfortunately, such relaxations have proved dangerous-a number of recent proof-of-concept exploits have shown how even minor relaxations of the control-flow policy can be exploited to effect attacks Artificial software diversity offers a different but complementary approach that randomizes programs in such a way that attacks succeeding against one program instance have a very low probability of success against other (independently randomized) instances of the same program. Probabilistic defenses rely on memory secrecy-i.e., the effects of randomization must remain hidden from attackers. One of the simplest and most widely adopted forms of artificial diversity is Address Space Layout Randomization (ASLR), which randomizes the base addresses of program segments at loadtime. Unfortunately, merely randomizing the base addresses does not yield sufficient entropy to preserve memory secrecy in many cases; there are numerous successful derandomization attacks against ASLR Recently, a new wave of implementation disclosure attacks Experiments show that O-CFI enjoys performance overheads comparable to standard fine-grained diversity and non-opaque, coarse-grained CFI. Moreover, O-CFI's control-flow checking logic is implemented using Intel x86/x64 memory-protection extensions (MPX) that are expected to be hardware-accelerated in commodity CPUs from 2015 onwards. We therefore expect even better performance for O-CFI in the near future. Our contributions are as follows: • We introduce O-CFI, the first low-overhead code-reuse defense that tolerates implementation disclosures. • We describe our implementation of a fully functional prototype that protects stripped, x86 legacy binaries without source code. II. THREAT MODEL Our work is motivated by the emergence of attacks against fine-grained diversity and coarse-grained control-flow integrity. We therefore introduce these attacks and distill them into a single, unified threat model. A. Bypassing Coarse-Grained CFI Ideally, CFI permits only programmer-intended control-flow transfers during a program's execution. The typical approach is to assign a unique ID to each permissible indirect controlflow target, and check the IDs at runtime. Unfortunately, this introduces performance overhead proportional to the degree of the graph-the more overlaps between valid target sets of indirect branch instructions, the more IDs must be stored and checked at each branch. Moreover, perfect CFI cannot be realized with a purely static control-flow graph; for example, the permissible destinations of function returns depend on the calling context, which is only known at runtime. Fine-grained CFI therefore implements a dynamically computed shadow stack, incurring high overheads To avoid this, coarse-grained CFI implementations resort to a reduced-degree, static approximation of the control-flow graph, and merge identifiers at the cost of reduced security. Typically, attackers need more than a single 4K page worth of code to find enough gadgets to mount a code-reuse attack. To discourage brute-force searches for more code pages, artificial diversity defenses routinely mine the address space with unmapped pages that abort the process if accessed B. Assumptions Given these sobering realities, we adopt a conservative threat model that assumes that attackers will eventually find and disassemble all code pages in victim processes. Our threat model therefore assumes that the adversary knows the complete in-memory code layout-including the locations of any gadgets required to launch a ROP attack. We also assume that the attacker can read and write the full contents of the heap and stack, as well as any data structures used by the dynamic loader. In keeping with common practice, we assume that data execution protection is activated, so that code page permissions can be maintained as either writable or executable but not both. However, we assume that attackers cannot safely perform a comprehensive, linear scan of virtual memory, since defenders may place unmapped guard pages at random locations. Instead, attackers must follow references from one disclosed memory page to another III. O-CFI OVERVIEW O-CFI combines insights from CFI and automated software diversity. It extends CFI with a new, coarse-grained CFI enforcement strategy inspired by bounds-checking, that validates control-flow transfers without divulging the bounds against which their destinations are checked. Bounds-checking is fast, the bounds are easier to conceal than arbitrary gadget locations, and the bounds are randomizable. This imbues CFI and fine-grained software diversity with an additional layer of protection against code-reuse attacks aided by implementation disclosures. As a result, O-CFI enjoys performance similar to coarse-grained CFI, with probabilistic security guarantees similar to fine-grained artificial diversity in the absence of implementation disclosures. Following traditional CFI, an O-CFI policy assigns to each indirect branch site a destination set that captures its set of permissible destination addresses. Such a graph can be derived from the program's source code or (with lesser precision) a conservative disassembly of its object code. We next reformulate this policy as a bounds-checking problem by reducing each destination set to only its minimal and maximal members. This policy approximation can be efficiently enforced by confining each branch to the memory-aligned addresses within its destination set range. All intended destination addresses are aligned within these bounds, so the enforcement conservatively preserves intended control-flows. Code layout is optimized to tighten the bounds, so that the set of unintended, aligned destinations within the bounds remains minimal. These few remaining unintended but reachable destinations are protected by the artificial diversity half of our approach. Our artificial diversity approach probabilistically protects the aligned, in-bounds, but policy-violating control-flows by applying fine-grained randomization to the binary code at load-time. While the overall strategy for implementing this randomization step is based on prior works Reformulating CFI in this way forces attackers to change their plan of attack. The recent attacks against coarse-grained CFI succeed by finding exploitable code that is reachable due to policy-relaxations needed for acceptable performance. These relaxations admit an alarming array of false-positives: instead of identifying the actual caller, all call-preceded instructions are incorrectly identified as permitted branch destinations. Such instructions saturate a typical address space, giving attackers too much wiggle room to build attacks. O-CFI counters this by changing the approximation approach: each branch destination is restricted to a relatively short span of aligned addresses, with all the bounds chosen pseudo-randomly at load-time. This greatly narrows the field of possible hijacks, and it removes the opportunity for attackers to analyze programs ahead of time for viable ROP gadget chains. In O-CFI, no two program instances admit the same set of ROP payloads, since the bounds are all randomized every time the program is loaded. Since the security of coarse-grained CFI depends in part on the precision of its policy approximation, it is worthwhile to improve the precision by tightening the bounds imposed upon each branch. This effectively reduces the space of attacker guesses that might succeed in hijacking any given branch. To reduce this space as much as possible, we introduce a novel binary code optimization, called portals, that minimizes the distance covered by the lowest and greatest element of each indirect branch's destination set. Our fine-grained artificial diversity implementation is an adaptation and extension of binary stirring To protect against information leaks that might disclose bounds information, our implementation is carefully designed to keep all bounds opaque to external threats. They are randomly chosen at load-time (as a side-effect of binary stirring) and stored in a bounds lookup table (BLT) located at a randomly chosen base address. The table size is very small relative to the virtual address space, and attackers cannot safely perform bruteforce scans of the full address space (see §II-B), so guessing the BLT's location is probabilistically infeasible for attackers. No code or data sections contain any pointer references to BLT addresses; all references are computed dynamically at load-time and stored henceforth exclusively in protected registers. A. Bounding the Control Flow For each indirect branch site with (non-empty) destination set D, O-CFI guards the branch instruction with a bounds-check that continues execution only if the impending target t satisfies t ∈ [min D, max D]. Indirect branch instructions include all control-flow transfer instructions that target computed destinations, including return instructions. Failure of the boundscheck solicits immediate process termination with an error code (for easier debugging). Termination could be replaced with a different intervention if desired, such as an automated attack analysis or alarm, followed by restart and re-randomization. The bounds-check implementation first loads the pair (min D, max D) from the BLT into registers via an indirect, indexed memory reference. The load instruction's arguments and syntax are independent of the BLT's location, concealing its address from attackers who can read the checking code. The impending branch target t is then checked against the loaded bounds. If the check succeeds, execution continues; otherwise the process immediately terminates with a bounds range (#BR) exception. The #BR exception helps distinguish between crashes and guessing attacks. To resist guessing attacks (e.g., BROP), web servers and other services should use this exception to trigger re-randomization as they restart. Following the approaches of PittSFIeld To bypass these checks, an attacker must craft a payload whose every gadget is properly aligned and falls within the bounds of the preceding gadget's conclusory indirect branch. The odds of guessing a reachable series of such gadgets decrease exponentially with the number of gadgets in the desired payload. B. Opacifying Control-flow Bounds Diversifying bounds. The bounds introduced by O-CFI constitute a coarse-grained CFI policy. Section II warns that such coarse granularity can lead to vulnerabilities. However, to exploit such vulnerabilities, attackers must discover which control-flows adhere to the CFI policy and which do not. To make the impermissible flows opaque to attackers, we use diversity. Our prototype uses a modified version of the technique outlined by Wartell et al. Performing fine-grain code randomization at load-time indirectly randomizes the ranges used to bound the control-flow. In contrast to other CFI techniques, attackers therefore do not have a priori knowledge of the control-flow bounds. Preventing Information Leaks. Attackers bypass fine grained diversity using information leaks, such as those described in §II-A. Were O-CFI's control-flow bounds expressed as constants in the instruction stream, attackers could bypass our defense via information leaks. To avoid this, we instead confine this sensitive information to an isolated data page, the BLT. The BLT is initialized at a random virtual memory address at load-time, and there are no pointer references (obfuscated or otherwise) to any BLT address in any code or data page in the process. This keeps its location hidden from attackers. Furthermore, we take additional steps to prevent accidental BLT disclosure via pointer leaks. Our prototype stores BLT base addresses in segment selectors-a legacy feature of all x86 processors. In particular, each load from the BLT uses the gs segment selector and a unique index to read the correct bounds. We only use the gs selector for instructions that implement bounds checks, so there are no other instructions that adversaries can reuse to learn its value. Attackers are also prevented from executing instructions that reveal the contents of the segment registers, since such instructions are privileged. To succeed, attackers must therefore (i) guess branch ranges, or (ii) guess the base address of the BLT. The odds of correctly guessing the location of the BLT are low enough to provide probabilistic protection. On 32-bit Windows Systems, for instance, the chances of guessing the base address are or less than one in two billion. Incorrect guesses alert defenders and trigger re-randomization with high probability (by accessing an unallocated memory page). The likelihood of successfully guessing a reachable gadget chain is a function of the length of the chain and the span of the bounds. The next section therefore focuses on reducing the average bounds span. C. Tightening Control-flow Check Bounds The distance between the lowest and highest intended destinations of any given indirect control-flow transfer instruction depends on the code layout. Placing indirect branches close to their targets both reduces bounds and improves locality, elevating both security and efficiency. Therefore we organize the code segment into clusters-one per indirect branch-each containing the basic blocks targeted by a particular branch. To accommodate blocks that are destinations of multiple distinct branch instructions, we consider three options: (i) put the block in one cluster and expand the bounds of other branches to include its address, (ii) create duplicate copies of the block in multiple clusters, or (iii) add a portal block to each cluster, which unconditionally jumps to the block. Each solution incurs a trade-off: expanding bounds reduces security, creating duplicates increases code size, and portals introduce runtime overhead. The options are not mutually exclusive, affording optimizers a range of strategies. Our experiments indicate that portals are often the best choice (see below). The capacity of the portal system limits the number of portals per nexus. Varying nexus capacity allows O-CFI to be tuned to different requirements. Setting it to zero prevents the creation of any portals, forcing the optimizer to choose alternative options. At the other extreme, setting no upper limit allows a portal to be created for every target, reducing all bounds ranges to wt, where w is the alignment width (usually 16 bytes; see §V-A) and t is the number of targets of the branch. At this setting, all indirect branches can only branch into a nexus, and through them, only to exactly those addresses that have been statically identified as targets. Thus, O-CFI with unbounded nexus capacity enforces fine-grained, static CFI. The extra layer of indirection imposed by a portal has a minor impact on runtime; there is thus a trade-off between security and performance. Users may opt for full CFI enforcement with O-CFI for security-critical components, and lower the nexus capacity to a desired performance level for less critical software. In our experiments, we found that a nexus capacity of 12 results in a significant reduction in bounds sizes with imperceptible performance effects. All of our experiments in §V use this nexus capacity. Section V-D details how different nexus capacities affect bound ranges. D. Example Defense against JIT-ROP The following example illustrates how O-CFI secures binaries against disclosure attacks. Consider a binary whose code segment contains five useful gadgets g 1 , . . . , g 5 . Each gadget terminates in an indirect branch protected by a bounds check. Under appropriate conditions, a disclosure attack such as JIT-ROP is able to recover a large portion of the runtime layout of the binary In our example, if g 1 is selected to be part of the payload, it can only be chained with gadget g 4 or g 5 . Attempting to jump from g 1 to any other gadget triggers a bounds violation that stops the attack. Similarly, an attack that hijacks a controlflow to c 1 can only redirect it to gadgets g 1 , g 2 , or g 3 ; all other gadgets are outside cluster c 1 and are therefore detected as impermissible destinations of the hijacked branch. Broadly speaking, all links in a payload's chain must traverse edges in the Cartesian product of the (aligned) gadget sets within the corresponding clusters. A successful attack must therefore limit itself to an extremely sparse graph of available edges. Our experiments (see §V-C) indicate that in practice the probability of successfully chaining gadgets in such a sparse graph is very low-just 0.01% for a four-gadget payload. The entropy of our procedure is further analyzed in §VI-A. IV. O-CFI IMPLEMENTATION We have implemented a fully functional prototype of O-CFI for the Intel x86 architecture. Our implementation uses a binary rewriting framework that secures COTS x86 binaries without source, debug, or relocation information. Like traditional CFI, however, we emphasize that O-CFI is equally suitable for inclusion in a compiler. Our rewriter generates a transformed version of the binary that leverages 1) a coarse-grained CFI policy that bounds control-flows, 2) fine-grained randomization to thwart traditional ROP attacks and diversify control-flow bounds so they become unknown and unreliable for attackers, 3) x86 segmentation registers to prevent accidental leakage of the bounds lookup table (BLT), and 4) an SFI framework similar to PittSFIeld [29] to enforce instruction alignment, denying attackers access to misaligned instructions that bypass bounds checks. The architecture of O-CFI is shown in A. Static Binary Rewriting 1) Conservative Disassembly: We first disassemble the code section using a conservative disassembler. Similar to the approach outlined by Wartell et al. [45], the code section is duplicated, with the old copy (renamed to .told) serving as a read-only data segment and the new copy (called .tnew) containing the rewritten executable code. The .told section is set non-executable, and all code blocks identified as possible targets of indirect jumps are overwritten with a five-byte tagged pointer. The tagged pointer consists of a tag byte (0xF4) followed by the four-byte address of that block in the .tnew section. The tag byte facilitates efficient runtime redirection of stale pointers to their correct targets, as explained below. Since the .told section preserves all static data at its original addresses, data pointers in the rewritten code section continue to behave correctly. This makes the rewriting system resilient to disassembly errors that misclassify data as code. Disassembly errors that misclassify code as data could omit such code from the .tnew section, resulting in a crash at runtime. To avoid this, we use settings that encourage the disassembler to interpret all bytes with valid instruction encodings as code. In our experiments, these settings suffice to avoid all disassembly errors that affect proper code translation. 2) SFI and Randomization Framework: To prevent attacks from jumping over the guards that constrain branch ranges, the new code segment is split into power-of-two sized basicblocks called chunks 6 Direct branches are statically rewritten to reference their new target addresses. Indirect branches require extra effort, since their exact targets are only known at runtime. At runtime, there are two common cases: (a) the impending target is already within the .tnew section (e.g., it was pushed by a call), or (b) the impending target is a stale pointer that points into the .told section (e.g., it was loaded from a method dispatch table in the heap, which the static rewriter does not modify). The first case requires no special treatment; the second solicits an efficient dynamic lookup and redirection of the stale pointer to its new location The stale pointer redirection mechanism is not relied upon for security. Like all indirect branch targets, redirected pointers undergo a mask and bounds-check before becoming controlflow destinations. Thus, corrupting or defeating the redirection mechanism does not circumvent the security policy. The ability to redirect code pointers lays the foundation for load-time randomization. Once the new randomized locations for basic blocks have been finalized, updating the values in the .told section allows our redirection mechanism to correctly redirect all indirect branches to the new, randomized block locations. Direct branches are simply modified in-place. 3) Branch Instrumentation: The above techniques enforce SFI and fine-grained randomization. This protects against traditional ROP attacks, but not against implementation-aware attacks, which require the additional hardening implemented by O-CFI's bounds-checking. Bounds-checking is applied after stale pointer redirection alongside masking, to further limit the set of accessible gadgets. SFI enforcement prevents attack payloads from circumventing these bounds checks. Algorithm 1 CreateClusters(S): Cluster basic blocks to place the targets of indirect branches as close together as possible. Input: S {the set of the basic blocks in the code segment} Output: C {a set of clusters, one per indirect branch. Each c ∈ C is a block set containing all targets of a specific branch, plus an empty nexus for later portal insertion.} C ← ∅ for all b ∈ Branches(S) do c ← ∅ for all t ∈ Targets(b) do b ← GetBasicBlock (t) if b / ∈ C then c ← c ∪ {b } end for {The nexus is an empty basic block to hold portals.} C ← C ∪ {(c ∪ CreateNexus())} end for {Add unclaimed basic-blocks into a single final cluster.} C ← C ∪ {(S − C)} Furthermore, due to randomization, the bounds remain unknown to implementation-aware attackers, and vary from program instance to program instance. Attacks cannot statically pre-compute bounds ranges because the runtime randomization phase changes bounds values on each execution. They also cannot dynamically leak the bounds, all of which are stored securely in the BLT and never leaked to the stack or heap. Attackers must therefore hazard guesses as to which gadget chains are safely accessible for any given program instance. Our bounds-checking logic is detailed in 4) Accurate Target Identification: To ensure that we identify all intended targets of indirect branches, we employ disassembly heuristics that identify a superset of potential targets. As an example, we follow the following sequence of steps to identify the set of potential targets for a return instruction: 1) Identify all code references to the function that contains the return. This includes direct and indirect branches to the function entry point, as well as to any basic block within the function. 2) For each identified branch that is not a call, find all code references that flow into it. 3) Recursively traverse all non-call references until a fixed point is reached (i.e., a set with only calls). 4) The instruction immediately after each call forms the target set for that return. Our heuristics are tuned to prefer false positives (non-targets treated as possibly valid destinations), since such errors do not significantly affect the operation of our system. In particular, each such error only marginally weakens the system's security (by admitting an unnecessary control-flow link that remains guarded by randomization) and slightly increases generated code size. A compiler-side solution could be more precise, at the cost of requiring source code and recompilation of programs. 5) Bounds Range Minimization: As discussed in §III-C, we use a combination of clustering and portals to reduce bounds ranges. While the portals themselves are created only at binary load-time, it is in the static phase that branch targets are clustered together and empty nexuses created. Algorithm 1 gives 7

Abstract—StackArmor is a comprehensive protection tech-nique for stack-based memory error vulnerabilities in binaries. It relies on binary analysis and rewriting strategies to drastically re-duce the uniquely high spatial and temporal memory predictabil-ity of traditional call stack organizations. Unlike prior solutions, StackArmor can protect against arbitrary stack-based attacks, requires no access to the source code, and offers a policy-driven protection strategy that allows end users to tune the security-performance tradeoff according to their needs. We present an implementation of StackArmor for x86 64 Linux and provide a detailed experimental analysis of our prototype on popular server programs and standard benchmarks (SPEC CPU2006). Our results demonstrate that StackArmor offers better security than prior binary- and source-level approaches, at the cost of only mod-est performance and memory overhead even with full protection. I.

Abstract—Malicious software are usually obfuscated to avoid detection and resist analysis. When new malware is encountered, such obfuscations have to be penetrated or removed (“deobfus-cated”) in order to understand the internal logic of the code and devise countermeasures. This paper discusses a generic approach for deobfuscation of obfuscated executable code. Our approach does not make any assumptions about the nature of the obfuscations used, but instead uses semantics-preserving program transformations to simplify away obfuscation code. We have applied a prototype implementation of our ideas to a variety of different kinds of obfuscation, including emulation-based obfuscation, emulation-based obfuscation with runtime code unpacking, and return-oriented programming. Our experimental results are encouraging and suggest that this approach can be effective in extracting the internal logic from code obfuscated using a variety of obfuscation techniques, including tools such as Themida that previous approaches could not handle.

Return-oriented programming (ROP) has become the dominant form of vulnerability exploitation in both user and kernel space. Many defenses against ROP exploits exist, which can significantly raise the bar against attackers. Although protecting existing code, such as applications and the kernel, might be possible, taking countermeasures against dynamic code, i.e., code that is generated only at run-time, is much harder. Attackers have already started exploiting Just-in-Time (JIT) engines, available in all modern browsers, to introduce their (shell)code (either native code or re-usable gadgets) during JIT compilation, and then taking advantage of it. Recognizing this immediate threat, browser vendors started employing defenses for hardening their JIT engines. In this paper, we show that—no matter the employed defenses—JIT engines are still exploitable using solely dynamically generated gadgets. We

We propose a novel dynamic software watermarking design based on Return-Oriented Programming (ROP). Our design formats watermarking code into well-crafted data arrange-ments that look like normal data but could be triggered to execute. Once triggered, the pre-constructed ROP execution will recover the hidden watermark message. The proposed ROP-based watermarking technique is more stealthy and re-silient over existing techniques since the watermarking code

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ABSTRACT Code randomization is an effective defense against code reuse attacks. It scrambles program code to prevent attackers from locating useful functions or gadgets. The key to secure code randomization is achieving high entropy. A practical approach to boost entropy is on-demand live randomization that works on running processes. However, enabling live randomization is challenging in that it often requires manual efforts to solve ambiguity in identifying function pointers. In this paper, we propose Remix, an efficient and practical live randomization system for both user processes and kernel modules. Remix randomly shuffles basic blocks within their respective functions. By doing so, it avoids the complexity of migrating stale function pointers, and allows mixing randomized and non-randomized code to strike a balance between performance and security. Remix randomizes a running process in two steps: it first randomly reorders its basic blocks, and then comprehensively migrates live pointers to basic blocks. Our experiments show that Remix can significantly increase randomness with low performance overhead on both CPU and I/O intensive benchmarks and kernel modules, even at very short randomization intervals.