The historical researches of Louis Couturat saved the logical work of Leibniz from the oblivion of neglect and forgetfulness. They revealed that Leibniz developed in succession several versions of a “logical calculus” (calculus ratiocinator or calculus universalis). In consequence of Couturat's investigations it has become well known that Leibniz's development of these logical calculi adumbrated the notion of a logistic system; and for these foreshadowings of the logistic treatment of formal logic Leibniz is rightly regarded as the father of symbolic logic.

It is clear from what has been said that it is scarcely possible to overestimate the debt which the contemporary student of Leibniz's logic owes to Couturat. This gratitude must, however, be accompanied by the realization that Couturat's own theory of logic is gravely defective. Couturat was persuaded that the extensional point of view in logic is the only one which is correct, an opinion now quite antiquated, and shared by no one. This prejudice of Couturat's marred his exposition of Leibniz's logic. It led him to battle with windmills: he viewed the logic of Leibniz as rife with shortcomings stemming from an intensional approach.

The task of this paper is a re-examination of Leibniz's logic. It will consider without prejudgment how Leibniz conceived of the major formal systems he developed as logical calculi – that is, these systems will be studied with a view to the interpretation or interpretations which Leibniz himself intends for them. The aim is to undo some of the damage which Couturat's preconception has done to the just understanding of Leibniz's logic and to the proper evaluation of his contribution.

The reader of Part VI will have noticed that among the set-theoretic models considered there some models were missing which were announced in Part II for certain proofs of independence. These models will be supplied now.

Mainly two models have to be constructed: one with the property that there exists a set which is its own only element, and another in which the axioms I–III and VII, but not Va, are satisfied. In either case we need not satisfy the axiom of infinity. Thereby it becomes possible to set up the models on the basis of only I–III, and either VII or Va, a basis from which number theory can be obtained as we saw in Part II.

On both these bases the Π0-system of Part VI, which satisfies the axioms I–V and VII, but not VI, can be constructed, as we stated there. An isomorphic model can also be obtained on that basis, by first setting up number theory as in Part II, and then proceeding as Ackermann did.

Let us recall the main points of this procedure.

For the sake of clarity in the discussion of this and the subsequent models, it will be necessary to distinguish precisely between the concepts which are relative to the basic set-theoretic system, and those which are relative to the model to be defined.

This paper reviews various symbolisms for the two-valued propositional calculus, and introduces a new set of signs which embodies the principles of Leśniewski's symbolism, yet resembles better known signs. This type of symbolism, serving as a diagram, may be used either in place of or as auxiliary to the usual symbolisms.

Zest for both system and objectivity is the formal logician's original sin. He pays for it by constant frustrations and by living ofttimes the life of an intellectual outcaste. The task of squeezing a large body of stubborn facts into a more or less rigid system can be a painful one, especially since the facts of mathematics are among the most stubborn of all facts. Moreover, the more general and abstract we get, the farther removed we are from the raw mathematical experience. As intuition ceases to operate effectively, we fall into many unexpected traps. The formal logician gets little sympathy for his frustrations. He is regarded as too rigid by his philosophical colleagues and too speculative by his mathematical friends. The life of an intellectual outcaste may be a result partly of temperament and partly of the youthfulness of the logic profession. The unfortunate lack of wide appeal of logic may, however, be prolonged partly on account of the fact that very little of the well-established techniques of mathematics seems applicable to the treatment of serious problems of logic.

The axiomatic method is well suited to provide results which are both exact and systematic. How attractive would it be if we could get an axiom system in which all the axioms and deductions were intuitively clear and all theorems of mathematics were provable? Such a system would undoubtedly satisfy Descartes who admits solely intuition and deduction, which are, for him, the only “mental operations by which we are able, wholly without fear of illusion, to arrive at the knowledge of things.” Indeed, according to Descartes, intuition and deduction “are the most certain routes to knowledge, and the mind should admit no others. All the rest should be rejected as suspect of error and dangerous.”

This paper will be concerned with an adaptation of Parry's trapezoid symbolism to the solving of problems.

In Parry's notation ‘pq’, in particular, sets forth the relationship that obtains when p is true, whether q be true or false. The symbol may also be regarded as manifesting the matrix analysis of p. It is this aspect which offers promise for manipulation. We propose to write the symbol ‘’ for p itself. It is to be kept in mind, however, that ‘’ is the symbol of p when there are but two variables. How to proceed when there are more than two will be seen later.

The symbol for q will in like manner be ‘⌋’.

Let ‘’ and ‘’, then, be regarded as elemental symbols setting forth the matrix analysis of the variables. Suppose further that in transcribing from conventional to the trapezoid symbolism we substitute for the variables only, retaining the constants (antithetically to Parry's procedure). Thus ‘p ∨ q’ becomes, not ‘pq’ as with Parry, but ‘ ∨ ’. Problem-solving consists of merging the trapezoidal symbols according to procedures specified by the constants. As symbols are merged, the constants are eliminated.

The resultant symbol constitutes the solution.

The rules for merging symbols are derived from the elementary processes of forming truth-tables:

Disjunction of two symbols consists of superimposing them.

Conjunction consists of preserving only such sides as are common to both the conjoined symbols.

Hilbert and Ackermann ([1], p. 107) define a first order axiom system as one in which the axioms contain one or more predicate constants, but no predicate variables. Here “axiom” refers to the specific subject-matter axioms and not to the rules of the restricted predicate calculus (quantification theory), which rules are presupposed for each first-order system. It is pointed out by them that an exception could be made for the predicate of identity; for the axiom scheme for this predicate, namely

which has in (b) the variable predicate F, could nevertheless be replaced, in any given first-order system, by a finite set of axioms without predicate variables. Thus, for example, if Φ[x, y) is the one constant predicate of such a system then PId(b) could be replaced by

Thus one postulates, in addition to the reflexivity, symmetry, and transitivity of identity, the substitutivity of identical entities in each of the possible “atomic” contexts of a variable (occurrences in the primitive predicates). In this method of introducing identity it has to be taken as an additional primitive predicate and further axioms are consequently needed. In such a system having PId(a), (b1)−(b3) as axioms, the scheme PId(b) can be derived as a meta-theorem of the system, F(x) then being any formula of the system.

The celebrated theorem of Löwenheim and Skolem tells us that every consistent set S of quantificational schemata (i.e., every set of well-formed formulas of the lower predicate calculus admitting of a true interpretation in some non-empty universe) admits of a true numerical interpretation (i.e., an interpretation of predicate letters such that all schemata of S come out true when the variables of quantification are construed as ranging over just the positive integers).

Later literature goes farther, and shows how, given S, actually to produce a numerical interpretation which will fit S in case S is consistent. The general case is covered by Kleene (see Bibliography). The special case where S contains just one schema (or any finite number, since we can form their conjunction) had been dealt with by Hilbert and Bernays. Certain extensions, along lines not to be embarked on here, have been made by Kleene, Kreisel, Hasenjäger, and Wang.

My present purpose is expository: to make the construction of the numerical interpretation, and the proof of its adequacy, more easily intelligible than they hitherto have been. The reasoning is mainly Kleene's, though closer in some ways to earlier reasoning of Gödel.

A primitive-recursive sequence of rational numbers sn is said to be primitive-recursively irrational, if there are primitive recursive functions n(k), i(p, q) > 0 and N(p, q) such that:

1. (k)(n ≥ n(k) → ∣sn – sn(k)∣ < 2−k).

2. (p)(q)(q > 0 & n ≥ N(p, q) → ∣sn ± p/q∣ > 1/i(p, q)).

The object of the present note is to establish the primitive-recursive irrationality of a sequence which converges to π. In a previous paper we proved the primitive-recursive irrationality of the exponential series Σxn/n!, for all rational values of x, and showed that a primitive-(general-) recursively irrational sequence sn is strongly primitive-(general-)recursive convergent in any scale, where a recursive sequence sn is said to be strongly primitive-(general-)recursive convergent in the scale r (r ≥ 2), if there is a non-decreasing primitive-(general-) recursive function r(k) such that

,

where [x] is the greatest integer contained in x, i.e. [x] = i if i ≤ x < i + 1, [x] = —i if i ≤ —x < i+1, where i is a non-negative integer.

A rational recursive sequence sn is said to be recursive convergent, if there is a recursive function n(k) such that

.

If a sequence sn is strongly recursive convergent in a scale r, then it is recursive convergent and its limit is the recursive real number where, for any k ≥ 0,

.

In this paper a theorem will be established which states that a particular axiom of choice is equivalent to complete distributivity of union and intersection. The theorem will be formulated and proved in the system of logic of [4]. In addition to definitions of [4], the following will be used.

In terms of these definitions, the theorem can be formulated as follows.

The dual of this statement, obtained by interchanging I and U, is also a theorem and has a similar proof.

Let C be an axiom system formalized within the first order functional calculus, and let C′ be related to C as the Bernays-Gödel set theory is related to the Zermelo-Fraenkel set theory. (An exact definition of C′ will be given later.) Ilse Novak [5] and Mostowski [8] have shown that, if C is consistent, then C′ is consistent. (The converse is obvious.) Mostowski has also proved the stronger result that any theorem of C′ which can be formalized in C is a theorem of C.

The proofs of Novak and Mostowski do not provide a direct method for obtaining a contradiction in C from a contradiction in C′. We could, of course, obtain such a contradiction by proving the theorems of C one by one; the above result assures us that we must eventually obtain a contradiction. A similar process is necessary to obtain the proof of a theorem in C from its proof in C′. The purpose of this paper is to give a new proof of these theorems which provides a direct method of obtaining the desired contradiction or proof.

The advantage of the proof may be stated more specifically by arithmetizing the syntax of C and C′.

In a previous paper I have described a decision method for testing whether or not two arbitrary finite truth-tables are “equal”, i.e., determine the same set of tautologies. The same problem for the case of infinite truth-tables remained open.

In the present note we shall show that the answer to the problem in the case of infinite truth-tables is negative; in fact we shall prove that there exists neither a decision method for testing the equality of arbitrary truth-tables, nor even one for testing the equality of what we call “recursive” truth-tables.

The terminology of Kalicki [3] will be used; for brevity's sake the discussion involving the notions of recursiveness and recursive enumerability uses the informal terminology and mode of argumentation employed by Post in [9].

Many properties of curves, surfaces, or other varieties in Algebraic Geometry can be formulated in the lower functional calculus as predicates of the coefficients of the polynomial or polynomials which define the variety (curve, surface) in question. For example, the property of a plane curve of order n to possess exactly m double points, or the property to be of genus p — where m and p are specified integers — can be formulated in this way. Similarly many statements on the relation between two or more varieties, e.g., concerning the number and type of their intersection points, can be expressed in the lower functional calculus. It is usual to study the properties of a variety in an algebraically closed field. Accordingly, it is of considerable interest to investigate the general structure of the class of predicates mentioned above in relation to algebraically closed fields. The following result will be proved in the present paper.

Main Theorem. Let F be a commutative algebraic field of arbitrary characteristic, and let F′ = F[x1, …, xn] be the ring of polynomials of n variables with coefficients in F. With every predicate Q{x1, …, xn) which is formulated in the lower functional calculus in terms of the relations of equality, addition, and multiplication and (possibly) in terms of some of the elements of F, there can be associated an ascending chain of ideals in F′,

such that

for every extension F* of F which is algebraically closed. In this formula, V0, …, V2k+ 1 are the varieties of the ideals J1, …, J2k+i in the coordinate space Sn: (x1, … xn) over F*, and VQ is the set of points of Sn which satisfy Q.

In a previous paper it was shown that the system K of basic logic could be formulated in a simpler way owing to the fact that the proper ancestral could be defined in terms of the other concepts of that system. In the present paper analogous but more far-reaching results will be obtained for the system K′ of extended basic logic. In particular we will show that negation and the dual of the proper ancestral, as well as the proper ancestral itself, are definable in terms of the other concepts of K′. Hence, in order to define K′, we need to add only a single non-finitary rule to the rules used to define K. This rule was already among the rules originally used to define K′. It asserts that ‘Aa’ is in K′ if (and only if) every ‘b’ is such that ‘ab’ is in K′.

We will also show that a large class of non-finitary classes and relations are represented in K′, among which is K′ itself, just as all finitary syntactical classes and all finitary two- and three-place syntactical relations are represented in K, one of which is the finitary syntactical class K itself. The point is that K′ is adequate to handle the sort of transfinite induction that is essential in formulating K′, just as K is adequate to handle the ordinary finitary mathematical induction required in defining K′.

Quantification theory, or the first-order predicate calculus, is ordinarily so formulated as to provide as theorems all and only those formulas which come out true under all interpretations in all non-empty domains. There are two strong reasons for thus leaving aside the empty domain.

(i) Where D is any non-empty domain, any quantificational formula which comes out true under all interpretations in all domains larger than D will come out true also under all interpretations in D. (Cf. [4], p. 92.) Thus, though any small domain has a certain triviality, all but one of them, the empty domain, can be included without cost. To include the empty one, on the other hand, would mean surrendering some formulas which are valid everywhere else and thus generally useful.

(ii) An easy supplementary test enables us anyway, when we please, to decide whether a formula holds for the empty domain. We have only to mark the universal quantifications as true and the existential ones as false, and apply truth-table considerations.

Incidentally, the existence of that supplementary test shows that there is no difficulty in framing an inclusive quantification theory (i.e., inclusive of the empty domain) if we so desire. A proof in this theory can be made to consist simply of a proof in the exclusive (or usual) theory followed by a check by the method of (ii). We may, however, be curious to see a more direct or autonomous formulation: one which does not consist, like the above, of the exclusive theory plus a rule of expurgation. And, in fact, such formulations have of late been forthcoming: Mostowski [5], Hailperin [3], and, as part of a broader context, Church [2]. I shall not presuppose acquaintance with these papers, except in my final paragraph (and then only with [3]).

Consider any interpreted theory Θ, formulated in the notation of quantification theory (or lower predicate calculus) with interpreted predicate letters. It will be proved that Θ is translatable into a theory, likewise formulated in the notation of quantification theory, in which there is only one predicate letter, and it a dyadic one.

Let us assume a fragment of set theory, adequate to assure the existence, for all x and y without regard to logical type, of the set {x, y) whose members are x and y, and to assure the distinctness of x from {x, y} and {{x}}. ({x} is explained as {x, x}.) Let us construe the ordered pair x; y in Kuratowski's fashion, viz. as {{x}, {x, y}}, and then construe x;y;z as x;(y;z), and x;y;z;w as x;(y;z;w), and so on. Let us refer to w, w;w, w;w;w, etc. as 1w, 2w, 3w, etc.

Suppose the predicates of Θ are ‘F1’, ‘F2’, …, finite or infinite in number, and respectively d1-adic, d2-adic, …. Now let Θ′ be a theory whose notation consists of that of quantification theory with just the single dyadic predicate ‘F’, interpreted thus:

The universe of Θ′ is to comprise all objects of the universe of Θ and, in addition, {x, y) for every x and y in the universe of Θ′. (Of course the universe of Θ may happen already to comprise all this.)

Now I shall show how the familiar notations ‘x = y’, ‘x = {y, z}’, etc., and ultimately the desired ‘’, ‘’, etc. themselves can all be defined within Θ′.

As we know, the system of material implication is led into an inconsistency by the Russellian class, defined as λx. Nϵxx. This class, however, does no harm to many other systems, for example, the three-valued system L3 given by J. Łukasiewicz. A natural question is whether or not there exist classes which affect some of these other logical systems. The main result of the present paper is to answer this question affirmatively. At the end of this paper we point out that the interpretation given by J. Łukasiewicz for the system L3 is not satisfactory, and propose a new interpretation.

B. Russell deduced the mentioned inconsistency by the aid of the notion of negation. Later on, H. B. Curry pointed out that we could get the same result without the aid of that notion. None of these results affects the system L3 and other similar systems. But these systems may be involved.

To show this, we need the following definitions.

A function of two variables Cpq will be called an implication when the following “implication rule” is valid:

Under this definition we should note that material equivalence Epq, for example, is an implication.

Let C be such an implication. Then the symbol “(Cp)iq” is defined recursively by

The class an is defined as λx. (Cϵxx)np, where p is a propositional variable. The rule of absorption of order n, denoted by (An), is: