Opaque Predicate Detection by Abstract Interpretation

Code obfuscation and software watermarking are well known techniques designed to prevent the illegal reuse of software. Code obfuscation prevents malicious reverse engineering, while software watermarking protects code from piracy. An interesting class of algorithms for code obfuscation and software watermarking relies on the insertion of opaque predicates. It turns out that attackers based on a dynamic or an hybrid static-dynamic approach are either not precise or time consuming in eliminating opaque predicates. We present an abstract interpretation-based methodology for removing opaque predicates from programs. Abstract interpretation provides the right framework for proving the correctness of our approach, together with a general methodology for designing efficient attackers for a relevant class of opaque predicates. Experimental evaluations show that abstract interpretation based attacks significantly reduce the time needed to eliminate opaque predicates

Opaque control-flow integrity.

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,

Evaluation Methodologies in Software Protection Research

Man-at-the-end (MATE) attackers have full control over the system on which the attacked software runs, and try to break the confidentiality or integrity of assets embedded in the software. Both companies and malware authors want to prevent such attacks. This has driven an arms race between attackers and defenders, resulting in a plethora of different protection and analysis methods. However, it remains difficult to measure the strength of protections because MATE attackers can reach their goals in many different ways and a universally accepted evaluation methodology does not exist. This survey systematically reviews the evaluation methodologies of papers on obfuscation, a major class of protections against MATE attacks. For 572 papers, we collected 113 aspects of their evaluation methodologies, ranging from sample set types and sizes, over sample treatment, to performed measurements. We provide detailed insights into how the academic state of the art evaluates both the protections and analyses thereon. In summary, there is a clear need for better evaluation methodologies. We identify nine challenges for software protection evaluations, which represent threats to the validity, reproducibility, and interpretation of research results in the context of MATE attacks

Security and trust in cloud computing and IoT through applying obfuscation, diversification, and trusted computing technologies

Cloud computing and Internet of Things (IoT) are very widely spread and commonly used technologies nowadays. The advanced services offered by cloud computing have made it a highly demanded technology. Enterprises and businesses are more and more relying on the cloud to deliver services to their customers. The prevalent use of cloud means that more data is stored outside the organization’s premises, which raises concerns about the security and privacy of the stored and processed data. This highlights the significance of effective security practices to secure the cloud infrastructure. The number of IoT devices is growing rapidly and the technology is being employed in a wide range of sectors including smart healthcare, industry automation, and smart environments. These devices collect and exchange a great deal of information, some of which may contain critical and personal data of the users of the device. Hence, it is highly significant to protect the collected and shared data over the network; notwithstanding, the studies signify that attacks on these devices are increasing, while a high percentage of IoT devices lack proper security measures to protect the devices, the data, and the privacy of the users. In this dissertation, we study the security of cloud computing and IoT and propose software-based security approaches supported by the hardware-based technologies to provide robust measures for enhancing the security of these environments. To achieve this goal, we use obfuscation and diversification as the potential software security techniques. Code obfuscation protects the software from malicious reverse engineering and diversification mitigates the risk of large-scale exploits. We study trusted computing and Trusted Execution Environments (TEE) as the hardware-based security solutions. Trusted Platform Module (TPM) provides security and trust through a hardware root of trust, and assures the integrity of a platform. We also study Intel SGX which is a TEE solution that guarantees the integrity and confidentiality of the code and data loaded onto its protected container, enclave. More precisely, through obfuscation and diversification of the operating systems and APIs of the IoT devices, we secure them at the application level, and by obfuscation and diversification of the communication protocols, we protect the communication of data between them at the network level. For securing the cloud computing, we employ obfuscation and diversification techniques for securing the cloud computing software at the client-side. For an enhanced level of security, we employ hardware-based security solutions, TPM and SGX. These solutions, in addition to security, ensure layered trust in various layers from hardware to the application. As the result of this PhD research, this dissertation addresses a number of security risks targeting IoT and cloud computing through the delivered publications and presents a brief outlook on the future research directions.Pilvilaskenta ja esineiden internet ovat nykyään hyvin tavallisia ja laajasti sovellettuja tekniikkoja. Pilvilaskennan pitkälle kehittyneet palvelut ovat tehneet siitä hyvin kysytyn teknologian. Yritykset enenevässä määrin nojaavat pilviteknologiaan toteuttaessaan palveluita asiakkailleen. Vallitsevassa pilviteknologian soveltamistilanteessa yritykset ulkoistavat tietojensa käsittelyä yrityksen ulkopuolelle, minkä voidaan nähdä nostavan esiin huolia taltioitavan ja käsiteltävän tiedon turvallisuudesta ja yksityisyydestä. Tämä korostaa tehokkaiden turvallisuusratkaisujen merkitystä osana pilvi-infrastruktuurin turvaamista. Esineiden internet -laitteiden lukumäärä on nopeasti kasvanut. Teknologiana sitä sovelletaan laajasti monilla sektoreilla, kuten älykkäässä terveydenhuollossa, teollisuusautomaatiossa ja älytiloissa. Sellaiset laitteet keräävät ja välittävät suuria määriä informaatiota, joka voi sisältää laitteiden käyttäjien kannalta kriittistä ja yksityistä tietoa. Tästä syystä johtuen on erittäin merkityksellistä suojata verkon yli kerättävää ja jaettavaa tietoa. Monet tutkimukset osoittavat esineiden internet -laitteisiin kohdistuvien tietoturvahyökkäysten määrän olevan nousussa, ja samaan aikaan suuri osuus näistä laitteista ei omaa kunnollisia teknisiä ominaisuuksia itse laitteiden tai niiden käyttäjien yksityisen tiedon suojaamiseksi. Tässä väitöskirjassa tutkitaan pilvilaskennan sekä esineiden internetin tietoturvaa ja esitetään ohjelmistopohjaisia tietoturvalähestymistapoja turvautumalla osittain laitteistopohjaisiin teknologioihin. Esitetyt lähestymistavat tarjoavat vankkoja keinoja tietoturvallisuuden kohentamiseksi näissä konteksteissa. Tämän saavuttamiseksi työssä sovelletaan obfuskaatiota ja diversifiointia potentiaalisiana ohjelmistopohjaisina tietoturvatekniikkoina. Suoritettavan koodin obfuskointi suojaa pahantahtoiselta ohjelmiston takaisinmallinnukselta ja diversifiointi torjuu tietoturva-aukkojen laaja-alaisen hyödyntämisen riskiä. Väitöskirjatyössä tutkitaan luotettua laskentaa ja luotettavan laskennan suoritusalustoja laitteistopohjaisina tietoturvaratkaisuina. TPM (Trusted Platform Module) tarjoaa turvallisuutta ja luottamuksellisuutta rakentuen laitteistopohjaiseen luottamukseen. Pyrkimyksenä on taata suoritusalustan eheys. Työssä tutkitaan myös Intel SGX:ää yhtenä luotettavan suorituksen suoritusalustana, joka takaa suoritettavan koodin ja datan eheyden sekä luottamuksellisuuden pohjautuen suojatun säiliön, saarekkeen, tekniseen toteutukseen. Tarkemmin ilmaistuna työssä turvataan käyttöjärjestelmä- ja sovellusrajapintatasojen obfuskaation ja diversifioinnin kautta esineiden internet -laitteiden ohjelmistokerrosta. Soveltamalla samoja tekniikoita protokollakerrokseen, työssä suojataan laitteiden välistä tiedonvaihtoa verkkotasolla. Pilvilaskennan turvaamiseksi työssä sovelletaan obfuskaatio ja diversifiointitekniikoita asiakaspuolen ohjelmistoratkaisuihin. Vankemman tietoturvallisuuden saavuttamiseksi työssä hyödynnetään laitteistopohjaisia TPM- ja SGX-ratkaisuja. Tietoturvallisuuden lisäksi nämä ratkaisut tarjoavat monikerroksisen luottamuksen rakentuen laitteistotasolta ohjelmistokerrokseen asti. Tämän väitöskirjatutkimustyön tuloksena, osajulkaisuiden kautta, vastataan moniin esineiden internet -laitteisiin ja pilvilaskentaan kohdistuviin tietoturvauhkiin. Työssä esitetään myös näkemyksiä jatkotutkimusaiheista

Code obfuscation and malware detection by abstract interpretation

Non disponibileAn obfuscating transformation aims at confusing a program in order to make it more difficult to understand while preserving its functionality. Software protection and malware detection are two major applications of code obfuscation. Software developers use code obfuscation in order to defend their programs against attacks to the intellectual property, usually called malicious host attacks. In fact, by making the programs more difficult to understand it is possible to obstruct malicious reverse engineering \u2013 a typical attack to the intellectual property of programs. On the other side, malware writers usually obfuscate their malicious code in order to avoid detection. In this setting, the ability of code obfuscation to foil most of the existing detection techniques, such as misuse detection algorithms, relies in their purely syntactic nature that makes malware detection sensitive to slight modifications of programs syntax. In the software protection scenario, researchers try to develop sophisticated obfuscating techniques that are able to resist as many attacks as possible. In the malware detection scenario, on the other hand, it is important to design advanced detection algorithms in order to detect as many variants of obfuscated malware as possible. It is clear how both malicious host and malicious code attacks represent harmful threats to the security of computer networks. In this dissertation, we are interested in both security issues described above. In particular, we describe a formal approach to code obfuscation and malware detection based on program semantics and abstract interpretation. This theoretical framework is useful in contrasting some well known drawbacks of software protection through code obfuscation, as well as for improving existing malware detection schemes. In fact, the lack of rigorous theoretical bases for code obfuscation prevents any possibility to formally study and certify their effectiveness in protecting proprietary programs. Moreover, in order to design malware detection schemes that are resilient to obfuscation we have to focus on program semantics rather than on program syntax. Our formal framework for code obfuscation relies on a semantics-based definition of code obfuscation that characterizes each program transformation T as a potential obfuscation in terms of the most concrete property preserved by T on program semantics. Deobfuscating techniques, and reverse engineering in general, usually begin with some sort of static program analysis, which can be specified as an abstraction of program semantics. In the software protection scenario, this observation naturally leads to model attackers as abstractions of program semantics. In fact, the abstraction modeling the attacker expresses the amount of information, namely the semantic properties, that the attacker is able to observe. It follows that, comparing the degree of abstraction of an attacker A with the one of the most concrete property preserved by an obfuscating transformation T, it is possible to understand whether obfuscation T defeats attack A. Following the same reasoning it is possible to compare the efficiency of different obfuscating transformations, as well as the ability of different attackers in contrasting a given obfuscation. We apply our semantics-based framework to a known control code obfuscation technique that aims at confusing the control flow of the original program by inserting opaque predicates. As argued above, an obfuscating transformation modifies a program while preserving an abstraction of its semantics. This means that different obfuscated versions of the same malware have to share (at least) the malicious intent, namely the maliciousness of their semantics, even if they may express it through different syntactic forms. The basic idea of our formal approach to malware detection is to use program semantics to model both malware and program behaviour, and semantic abstractions to hide the details changed by the obfuscation. Thus, given an obfuscation T, we are interested in defining an abstraction of program semantics that does not distinguish between the semantics of malware M and the semantics of its obfuscated version T(M). In particular, we provide this suitable abstraction for an interesting class of commonly used obfuscating transformations. It is clear that, given a malware detector D, it is always possible to define its semantic counterpart by analyzing how D works on program semantics. At this point, by translating both malware detectors and obfuscating transformations in the semantic world, we are able to certify which obfuscations a detector is able to handle. This means that our semanticsbased framework provides a formal setting where malware detectors designers can prove the efficiency of their algorithms