DOCSLIB.ORG

A Runtime Assertion Checker for the Java Modeling Language Yoonsik Cheon Iowa State University

Total Page：16

File Type：pdf, Size：1020Kb

Download full-text PDF Read full-text
A Runtime Assertion Checker for the Java Modeling Language Yoonsik Cheon Iowa State University View metadata, citation and similar papers at core.ac.uk brought to you by CORE provided by Digital Repository @ Iowa State University Computer Science Technical Reports Computer Science 4-2003 A Runtime Assertion Checker for the Java Modeling Language Yoonsik Cheon Iowa State University Follow this and additional works at: http://lib.dr.iastate.edu/cs_techreports Part of the Software Engineering Commons Recommended Citation Cheon, Yoonsik, "A Runtime Assertion Checker for the Java Modeling Language" (2003). Computer Science Technical Reports. 308. http://lib.dr.iastate.edu/cs_techreports/308 This Article is brought to you for free and open access by the Computer Science at Iowa State University Digital Repository. It has been accepted for inclusion in Computer Science Technical Reports by an authorized administrator of Iowa State University Digital Repository. For more information, please contact [email protected]. A Runtime Assertion Checker for the Java Modeling Language Abstract JML compiler to translate Java programs annotated with JML specifications into Java bytecode. The ompc iled bytecode transparently checks JML specifications at runtime. The JML ompc iler supports separate and modular compilation. The ppra oach brings programming benefits uchs as debugging and testing to BISLs and also helps programmers to use BISLs in their daily programming. A set of translation rules are defined from JML expressions into assertion checking code. The translation rules handle various kinds of undefinedness, such as runtime exceptions and non-executable constructs, in such a way as to both satisfy the standard rules of logic and detect as many assertion violations as possible. The rules also support various forms of quantifiers. Specification-only declarations such as model fields, ghost fields, and model methods are translated into access methods; e.g., an access method for a model field is an abstraction function that calculates an abstract value from the program state. The specification state of a stateful interface, due to specification-only fields such as ghost fields, is represented as a separate assertion class. Thus, an object's specification state is distributed over the object itself and one assertion object for each interface that its class implements. Assertion checking is also distributed in that a subtype delegates the responsibility of checking inherited specifications to its supertypes (or the assertion classes of its superinterfaces). The delegation approach supports multiple inheritance, and is modular. Finally, the effectiveness and practicality of runtime assertion checking is demonstrated by applying it to program testing. An approach is implemented that significantly automates unit testing. The eyk idea of the approach is to view interface specifications as test oracles and to use the runtime assertion checker as the decision procedure of the test oracles. The ppra oach also shows that the runtime assertion checker can be an effective framework for developing specification-based tools. Keywords documentation, formal methods, inheritance of specifications, programming by contract, runtime assertion checking, specification language, tools, unit testing, Java language, JML compiler, JML language Disciplines Software Engineering This article is available at Iowa State University Digital Repository: http://lib.dr.iastate.edu/cs_techreports/308 A Runtime Assertion Checker for the Java Modeling Language Yoonsik Cheon TR #03-09 April 2003 Keywords: documentation, formal methods, inheritance of speci¯cations, programming by con- tract, runtime assertion checking, speci¯cation language, tools, unit testing, Java language, JML compiler, JML language. 2001 CR Categories: D.2.1 [Software Engineering] Requirements/ Speci¯cations | languages, tools, JML; D.2.2 [Software Engineering] Design Tools and Techniques | modules and interfaces, object-oriented design methods; D.2.4 [Software Engineering] Software/Program Veri¯cation | As- sertion checkers, class invariants, correctness proofs, formal methods, programming by contract, reli- ability, tools, validation, JML; D.3.2 [Programming Languages] Language Classi¯cations | Object- oriented languages; F.3.1 [Logics and Meanings of Programs] Specifying and Verifying and Reasoning about Programs | Assertions, invariants, logics of programs, mechanical veri¯cation, pre- and post- conditions, speci¯cation techniques. Copyright °c 2003 by Yoonsik Cheon. All rights reserved. Department of Computer Science 226 Atanaso® Hall Iowa State University Ames, Iowa 50011-1041, USA Contents Acknowledgements vii Abstract viii 1 Introduction 1 1.1 Background . 1 1.1.1 Runtime Assertion Checking . 2 1.1.2 The Java Modeling Language . 3 1.2 Objectives . 4 1.3 The Problem . 5 1.3.1 Semantic Discrepancies between JML and Java . 6 1.3.2 Advanced Features of JML . 6 1.3.3 Research Scope . 7 1.4 Approach . 7 1.4.1 What, Where, and When to Check? . 7 1.4.2 Assertion Blocks, Methods, and Classes . 8 1.4.3 Local Contextual Interpretation . 9 1.4.4 Dynamic Delegation . 9 1.4.5 Access Methods . 10 1.5 Implementation . 10 1.5.1 The JML Compiler . 11 1.5.2 Reporting Assertion Violations . 12 1.6 Application | Speci¯cations as Test Oracles . 13 1.7 Contributions . 14 1.8 Outline . 14 2 An Overview of JML 16 2.1 Introduction . 16 2.2 Assertions and Expressions . 17 2.3 Method Speci¯cations . 18 2.3.1 Speci¯cation Clauses . 18 2.3.2 Heavyweight versus Lightweight . 19 2.3.3 Syntactic Sugars . 19 2.3.4 Privacy of Speci¯cations . 20 2.3.5 Semantics . 21 2.4 Type Speci¯cations . 21 2.4.1 Invariants . 21 2.4.2 Constraints . 21 2.4.3 Abstract Speci¯cations . 22 2.5 Inheritance of Speci¯cations . 22 i 2.6 An Example . 23 2.7 Discussion . 25 3 Expressions and Assertions 27 3.1 Introduction . 27 3.1.1 Notations . 28 3.2 The Unde¯nedness Problem . 29 3.2.1 An Example . 29 3.2.2 Demonic versus Angelic . 30 3.2.3 Contextual Interpretation . 31 3.2.4 Translation Rules . 32 3.2.5 Non-executable Constructs . 35 3.3 Quanti¯ed Expressions . 36 3.3.1 Abstract Syntax Extended . 36 3.3.2 Semantic Clari¯cation . 37 3.3.3 Pattern-based Static Analysis . 37 3.4 Related Work . 39 3.4.1 Unde¯nedness . 39 3.4.2 Quanti¯ed Expressions . 40 3.5 Discussion . 42 3.5.1 Anomaly of Contextual Interpretation . 42 3.5.2 Referring to Pre-state Expressions . 43 3.5.3 Other Approaches to Quanti¯ers . 43 3.6 Summary . 43 4 Method Speci¯cations 45 4.1 Introduction . 45 4.1.1 Reporting Assertion Violations . 46 4.1.2 Translation Scheme . 47 4.1.3 Wrapper Approach . 48 4.1.4 Outline . 49 4.2 Desugaring Speci¯cations . 49 4.2.1 Eliminating Old Variables . 51 4.2.2 Eliminating Old Expressions . 51 4.3 Wrapper Methods . 51 4.4 Precondition Methods . 53 4.5 Postcondition Methods . 54 4.5.1 Normal Postconditions . 54 4.5.2 Exceptional Postconditions . 55 4.6 Pre-state Expressions . 57 4.7 In-line Assertions . 59 4.7.1 Assertions, Assumptions, and Reasons . 59 4.7.2 Unreachable Statements . 60 4.7.3 Set Statements . 60 4.7.4 Loop Invariants and Variants . 61 4.8 Discussion . 64 4.8.1 Privacy of Method Speci¯cations . 64 4.8.2 Constructors and Finalizers . 65 4.8.3 Initializers . 65 ii 5 Type Speci¯cations 66 5.1 Introduction . 66 5.2 Invariants.
Load more

Recommended publications
  • Assertions, Pre/Post- Conditions and Invariants
    9/14/12 Assertions, pre/post- conditions and invariants Section 2.1 in Walls and Mirrors Section 4.5 Rosen Programming as a contract n Specifying what each method does q Specify it in a comment before method's header n Precondition q What is assumed to be true before the method is executed q Caller obligation n Postcondition q Specifies what will happen if the preconditions are met q Method obligation 1 9/14/12 Class Invariants n A class invariant is a condition that all objects of that class must satisfy while it can be observed by clients n What about Points in Cloud? q boundaries? q center? What is an assertion? n An assertion is a statement that says something about the state of your program n Should be true if there are no mistakes in the program //n == 1 while (n < limit) { n = 2 * n; } // what could you state here? 2 9/14/12 What is an assertion? n An assertion is a statement that says something about the state of your program n Should be true if there are no mistakes in the program //n == 1 while (n < limit) { n = 2 * n; } //n >= limit //more? What is an assertion? n An assertion is a statement that says something about the state of your program n Should be true if there are no mistakes in the program //n == 1 while (n < limit) { n = 2 * n; } //n >= limit //n is the smallest power of 2 >= limit 3 9/14/12 assert Using assert: assert n == 1; while (n < limit) { n = 2 * n; } assert n >= limit; When to use Assertions n We can use assertions to guarantee the behavior.
    [Show full text]
  • Sound Invariant Checking Using Type Modifiers and Object Capabilities
    Sound Invariant Checking Using Type Modifiers and Object Capabilities. Isaac Oscar Gariano Victoria University of Wellington [email protected] Marco Servetto Victoria University of Wellington [email protected] Alex Potanin Victoria University of Wellington [email protected] Abstract In this paper we use pre existing language support for type modifiers and object capabilities to enable a system for sound runtime verification of invariants. Our system guarantees that class invariants hold for all objects involved in execution. Invariants are specified simply as methods whose execution is statically guaranteed to be deterministic and not access any externally mutable state. We automatically call such invariant methods only when objects are created or the state they refer to may have been mutated. Our design restricts the range of expressible invariants but improves upon the usability and performance of our system compared to prior work. In addition, we soundly support mutation, dynamic dispatch, exceptions, and non determinism, while requiring only a modest amount of annotation. We present a case study showing that our system requires a lower annotation burden compared to Spec#, and performs orders of magnitude less runtime invariant checks compared to the widely used ‘visible state semantics’ protocols of D, Eiffel. We also formalise our approach and prove that such pre existing type modifier and object capability support is sufficient to ensure its soundness. 2012 ACM Subject Classification Theory of computation → Invariants, Theory of computation → Program verification, Software and its engineering → Object oriented languages Keywords and phrases type modifiers, object capabilities, runtime verification, class invariants Digital Object Identifier 10.4230/LIPIcs.CVIT.2016.23 1 Introduction Object oriented programming languages provide great flexibility through subtyping and arXiv:1902.10231v1 [cs.PL] 26 Feb 2019 dynamic dispatch: they allow code to be adapted and specialised to behave differently in different contexts.
    [Show full text]
  • The Contract Pattern Permission Granted to Copy for Plop ’97 Conference
    Copyright 1997, Michel de Champlain The Contract Pattern Permission granted to copy for PLoP ’97 Conference. All other rights reserved. Michel de Champlain Department of Computer Science University of Canterbury, Christchurch, New Zealand [email protected] http://www.cosc.canterbury.ac.nz/~michel 12 August 1997 Abstract This paper describes the Contract pattern, an idiom that lets you apply assertions to guarantee pre-conditions and post-conditions of methods and invariants on the state of objects. This pattern can be used to develop reliable classes by making the Design by Contract methodology—introduced by Meyer in the specific context of the Eiffel language—available in Java and possibly other object-oriented languages. Intent Provide an implementation of the Design by contract methodology [Meyer88] for developing reliable classes and create robust objects in Java. This programming pattern lets you apply assertions to guarantee the pre-conditions and post-conditions of methods and invariants on the state of objects. Also Known As Design by contract Motivation Sometimes it’s necessary to restrict a class interface because you don’t want to provide all the possible input combinations for all collaborators. One way to achieve this goal is to make minimal and consistent methods that impose pre-conditions on input parameters. Moreover, you might want to ensure the class behavior by imposing post-conditions before methods return. You might also need to make sure that the class is always in a stable state. Such an approach can result in reliable classes to create robust objects. Consider for example a bounded counter that has both a lower and upper bound.
    [Show full text]
  • Safely Creating Correct Subclasses Without Superclass Code Clyde Dwain Ruby Iowa State University
    Iowa State University Capstones, Theses and Retrospective Theses and Dissertations Dissertations 2006 Modular subclass verification: safely creating correct subclasses without superclass code Clyde Dwain Ruby Iowa State University Follow this and additional works at: https://lib.dr.iastate.edu/rtd Part of the Computer Sciences Commons Recommended Citation Ruby, Clyde Dwain, "Modular subclass verification: safely creating correct subclasses without superclass code " (2006). Retrospective Theses and Dissertations. 1877. https://lib.dr.iastate.edu/rtd/1877 This Dissertation is brought to you for free and open access by the Iowa State University Capstones, Theses and Dissertations at Iowa State University Digital Repository. It has been accepted for inclusion in Retrospective Theses and Dissertations by an authorized administrator of Iowa State University Digital Repository. For more information, please contact [email protected]. Modular subclass verification: Safely creating correct subclasses without superclass code by Clyde Dwain Ruby A dissertation submitted to the graduate faculty in partial fulfillment of the requirements for the degree of DOCTOR OF PHILOSOPHY Major: Computer Science Program of Study Committee Gary T. Leavens, Major Professor Samik Basu Clifford Bergman Shashi K. Gadia Jonathan D. H. Smith Iowa State University Ames, Iowa 2006 Copyright © Clyde Dwain Ruby, 2006. All rights reserved. UMI Number: 3243833 Copyright 2006 by Ruby, Clyde Dwain All rights reserved. UMI Microform 3243833 Copyright 2007 by ProQuest Information and Learning Company. All rights reserved. This microform edition is protected against unauthorized copying under Title 17, United States Code. ProQuest Information and Learning Company 300 North Zeeb Road P.O. Box 1346 Ann Arbor, MI 48106-1346 ii TABLE OF CONTENTS ACKNOWLEDGEMENTS .
    [Show full text]
  • 3. Design by Contract
    3. Design by Contract Oscar Nierstrasz Design by Contract Bertrand Meyer, Touch of Class — Learning to Program Well with Objects and Contracts, Springer, 2009. 2 Bertrand Meyer is a French computer scientist who was a Professor at ETH Zürich (successor of Niklaus Wirth) from 2001-2015. He is best known as the inventor of “Design by Contract”, and as the designer of the Eiffel programming language, which provides built-in for DbC. DbC was first described in a technical report by Meyer in 1986: https://en.wikipedia.org/wiki/Design_by_contract Who’s to blame? The components fit but the system does not work. Who’s to blame? The component developer or the system integrator? 3 DbC makes clear the “contract” between a supplier (an object or “component”) and its client. When something goes wrong, the contract states whose fault it is. This simplifies both design and debugging. Why DbC? > Design by Contract —documents assumptions (what do objects expect?) —simplifies code (no special actions for failure) —aids debugging (identifies who’s to blame) 4 As we shall see, DbC improves your OO design in several ways. First, contracts make explicit the assumptions under which an object (supplier) will work correctly. Second, they simplify your code, since no special action is required when things go wrong — the exception handling framework provides the necessary tools. Third, contracts help in debugging since errors are caught earlier, when contracts are violated, not when your program crashes because of an invalid state, and it is clear where to lay the blame for the violation (i.e., in the object or its client).
    [Show full text]
  • FORTH-ICS / TR 381 July 2006 'An Informal Proof on the Contradictory
    FORTH-ICS / TR 381 July 2006 ‘An Informal Proof on the Contradictory Role of Finalizers in a Garbage Collection Context’ Anthony Savidis ABSTRACT Garbage collection in OOP languages provides facilities to hook code that is executed upon object finalization. Semantically, the point in time that finalizer code is executed is not determined by the application logic, but by the garbage collection system. This fact renders a potential mismatch, since application-specific code, i.e. the finalizer implementation, normally affecting program state and control flow, is called automatically at a point in time that is indifferent to the application semantics. Although an analogous situation is observed in event-based systems, since event processors are called-back asynchronously by the underlying system, there is a fundamental difference: while event generation is essentially nondeterministic, originated from external event sources, object destruction is a semantic action that is always decided by applications. In summary, the mismatch is that although applications decide if and when destruction occurs, the garbage collector is responsible to commit the relevant code. We prove that this mismatch is due to the contradictory presence of finalizers in a garbage collection context. Page 1 Introduction Preface This technical report has been also motivated by requests from colleagues to report the informal proof for the contradictory role of finalizers in garbage collection, simply relying on informal common software engineering principles. Additionally, the reporting has been also motivated by arguments that came to my attention via personal communication that either the claim cannot be proved (I assume that “cannot prove” arguments are subject to formal proofs, so such arguments are more weak than the claim they attack), or that the claim is trivial requiring essentially no proof (intuitively true, but does not explain why Java, C#, Python, just to name a few languages, still adopt finalizers).
    [Show full text]
  • A Machine-Checked Proof for a Translation of Event-B Machines to JML
    A Machine-Checked Proof for a Translation of Event-B Machines to JML N´estorCata~no1,Camilo Rueda2, and Tim Wahls3 1 EAFIT Unviersity [email protected] 2 Camilo Rueda Pontificia Universidad Javeriana [email protected] 3 Tim Wahls Dickinson College [email protected] Abstract. We present a machine-checked soundness proof of a trans- lation of Event-B to the Java Modeling Language (JML). The transla- tion is based on an operator EB2Jml that maps Event-B events to JML method specifications, and deterministic and non-deterministic assign- ments to JML method post-conditions. This translation has previously been implemented as the EventB2Jml tool. We adopted a taking our own medicine approach in the formalisation of our proof so that Event-B as well as JML are formalised in Event-B and the proof is discharged with the Rodin platform. Hence, for any Event-B substitution (whether an event or an assignment) and for the JML method specification obtained by applying EventB2Jml to the substitution, we prove that the semantics of the JML method specification is simulated by the semantics of the sub- stitution. Therefore, the JML specification obtained as translation from the Event-B substitution is a refinement of the substitution. Our proof includes invariants and the standard Event-B initialising event, but it does not include full machines or Event-B contexts. We assume that the semantics of JML and Event-B operate both on the same initial and final states, and we justify our assumption. Keywords: Event-B, JML, EventB2Jml, Formal Proof, Translation, Rodin 1 Introduction arXiv:1309.2339v1 [cs.SE] 9 Sep 2013 In an earlier work [14], we have proposed a design methodology that combines the use of design-by-contract techniques with JML [36,29,33] and correctness-by- construction techniques with Event-B [1].
    [Show full text]
  • Verification of Object Oriented Programs Using Class Invariants
    Verification of Object Oriented Programs Using Class Invariants Kees Huizing and Ruurd Kuiper and SOOP?? Eindhoven University of Technology, PO Box 513, 5600 MB Eindhoven, The Netherlands, [email protected], [email protected] Abstract A proof system is presented for the verification and derivation of object oriented pro- grams with as main features strong typing, dynamic binding, and inheritance. The proof system is inspired on Meyer’s system of class invariants [12] and remedies its unsound- ness, which is already recognized by Meyer. Dynamic binding is treated in a flexible way: when throughout the class hierarchy overriding methods respect the pre- and post- conditions of the overridden methods, very simple proof rules for method calls suffice; more powerful proof rules are supplied for cases where one cannot or does not want to follow this restriction. The proof system is complete relative to proofs for properties of pointers and the data domain. 1 Introduction Although formal verification is not very common in the discipline of object oriented programming, the importance of formal specification is generally ac- knowledged ([12]). With the increased interest in component based develop- ment, it becomes even more important that components are specified in an un- ambiguous manner, since users or buyers of components often have no other knowledge about a component than its specification and at the same time rely heavily on its correct functioning in their framework. The specification of a class, sometimes called contract, usually contains at least pre- and postcondi- tions for the public mehtods and a class invariant. A class invariant expresses which states of the objects of the class are consis- tent, or “legal”.
    [Show full text]
  • The Fields and the Class Invariant /** B[0..Size-1] Is a Min-Heap, I.E
    Heaps with priorities We previously studied the max-heap, which implements a bag of integers with insertion of a value and extraction (polling) of a maximum value in O(log n) time for a heap of size n. Min-heaps are similar. We now extend the idea to implement a min-heap of a set of distinct values each with a priority. The priority will be a double value. The separation of value from priority is needed in several different situations. Here are two: • Consider finding the shortest route from Ithaca, NY, to Washington, D.C. A heap will contain points along the way, and the priorities will be the shortest distances found thus far from Ithaca to those points. • In a discrete event simulation, pending events can be kept in a heap, with the priority being the time at which an event is to occur. Keeping the set in a min-heap makes it easy to extract the next event to process. We will define all the methods needed, but we won’t show many of the method bodies, since they are fairly easy to write based on the earlier discussion of max- and min-heaps. Class Heap and inner class Item /** A heap of elements of type E. */ We will implement a generic min-heap with priorities. The public class Heap<E> { start of class Heap appears to the right. An object of this class implements a heap of elements of type E. For example, E could private class Item { be a point on a map or an event to be simulated.
    [Show full text]
  • Weaving Aspects Into C++ Applications for Validation of Temporal Invariants
    Weaving Aspects into C++ Applications for Validation of Temporal Invariants Tanton H. Gibbs and Brian A. Malloy Computer Science Department Clemson University Clemson, SC 29634, USA thgibbs, malloy @cs.clemson.edu f g Abstract [5]. During maintenance, invariants have proven invaluable for disclosing locations in the program In this paper, we describe temporal invariants, where the maintainer may have introduced a fault which are class invariants that are qualified by the [10]. operators eventually, always, never, or already. Tem- The problem with class invariants is that some poral invariants can capture assertions that may not important assertions about a class are initially in- be valid initially but, as the program continues, must valid, due to circular dependencies or the monoton- eventually become valid. Moreover, temporal invari- ically increasing nature of the system under con- ants can indicate references to memory that should struction [8]. For example, during compilation of eventually be deallocated. To facilitate incorporation a program, a name object is constructed to facil- of temporal invariants as a maintenance or reengi- itate name lookup. However, certain fields in the neering activity, we weave invariants into the sys- name object, such as the corresponding scope of the tem as aspects. In our case study of a C++ sys- name, may not be known until after name lookup tem, the aspects are woven into join points using has occurred [24]. Moreover, assertions about a policies. We investigate the effectiveness of tem- class may have an important impact on memory poral invariants and we compare the performance management. For example, some heap-based ob- of our aspect-oriented implementation with several jects of a class must be eventually deleted or system other approaches.
    [Show full text]
  • Modular Verification of Static Class Invariants
    Modular verification of static class invariants K. Rustan M. Leino 1 and Peter Muller¨ 2 1 Microsoft Research, Redmond, WA, USA, [email protected] 2 ETH Zurich,¨ Switzerland, [email protected] Abstract. Object invariants describe the consistency of object-oriented data struc- tures and are central to reasoning about the correctness of object-oriented soft- ware. But object invariants are not the only consistency conditions on which a program may depend. The data in object-oriented programs consists not just of object fields, but also of static fields, which hold data that is shared among ob- jects. The consistency of static fields is described by static class invariants, which are enforced at the class level. Static class invariants can also mention instance fields, describing the consistency of dynamic data structures rooted in static fields. Sometimes there are even consistency conditions that relate the instance fields of many or all objects of a class; static class invariants describe these relations, too, since they cannot be enforced by any one object in isolation. This paper presents a systematic way (a methodology) for specifying and verify- ing static class invariants in object-oriented programs. The methodology supports the three major uses of static fields and invariants in the Java library. The method- ology is amenable to static, modular verification and is sound. Keywords: Static class invariant, verification, object-oriented programming, sta- tic field 1 Introduction A central problem in reasoning about a computer program’s correctness comes down to reasoning about which program states the program can ever reach. Programmers rely on that only some states are reached, most prominently by assuming that the pro- gram’s data structures satisfy certain properties.
    [Show full text]
  • Software Engineering Concepts: Some Principles Managing Resources in a Class
    Software Engineering Concepts: Some Principles Managing Resources in a Class CS 311 Data Structures and Algorithms Lecture Slides Wednesday, September 16, 2009 Glenn G. Chappell Department of Computer Science University of Alaska Fairbanks [email protected] © 2005–2009 Glenn G. Chappell Unit Overview Advanced C++ & Software Engineering Concepts Major Topics: Advanced C++ Major Topics: S.E. Concepts The structure of a package Abstraction Parameter passing Invariants Operator overloading Testing Silently written & called Some principles functions Pointers & dynamic allocation Managing resources in a class Templates Containers & iterators Error handling Introduction to exceptions Introduction to Linked Lists 16 Sep 2009 CS 311 Fall 2009 2 Review Software Engineering Concepts: Invariants An invariant is a condition that is always true at a particular point in an algorithm. Special kinds Precondition . An invariant at the beginning of a function. The responsibility for making sure the preconditions are true rests with the calling code. What must be true for the function to execute properly. Postcondition . An invariant at the end of a function. Tells what services the function has performed for the caller. Describe the function’s effect using statements about objects & values. Class invariant . An invariant that holds whenever an object of the class exists, and execution is not in the middle of a public member function call. Statements about data members that indicate what it means for an object to be valid or usable. 16 Sep 2009 CS 311 Fall 2009 3 Review Software Engineering Concepts: Testing Recommended development process: Step 1. Get the code to compile . Write dummy versions of all modules.
    [Show full text]